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linux/mm/Kconfig

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# SPDX-License-Identifier: GPL-2.0-only
menu "Memory Management options"
#
# For some reason microblaze and nios2 hard code SWAP=n. Hopefully we can
# add proper SWAP support to them, in which case this can be remove.
#
config ARCH_NO_SWAP
bool
config ZPOOL
bool
menuconfig SWAP
bool "Support for paging of anonymous memory (swap)"
depends on MMU && BLOCK && !ARCH_NO_SWAP
default y
help
This option allows you to choose whether you want to have support
for so called swap devices or swap files in your kernel that are
used to provide more virtual memory than the actual RAM present
in your computer. If unsure say Y.
config ZSWAP
bool "Compressed cache for swap pages"
depends on SWAP
select CRYPTO
select ZPOOL
help
A lightweight compressed cache for swap pages. It takes
pages that are in the process of being swapped out and attempts to
compress them into a dynamically allocated RAM-based memory pool.
This can result in a significant I/O reduction on swap device and,
in the case where decompressing from RAM is faster than swap device
reads, can also improve workload performance.
config ZSWAP_DEFAULT_ON
bool "Enable the compressed cache for swap pages by default"
depends on ZSWAP
help
If selected, the compressed cache for swap pages will be enabled
at boot, otherwise it will be disabled.
The selection made here can be overridden by using the kernel
command line 'zswap.enabled=' option.
zswap: shrink zswap pool based on memory pressure Currently, we only shrink the zswap pool when the user-defined limit is hit. This means that if we set the limit too high, cold data that are unlikely to be used again will reside in the pool, wasting precious memory. It is hard to predict how much zswap space will be needed ahead of time, as this depends on the workload (specifically, on factors such as memory access patterns and compressibility of the memory pages). This patch implements a memcg- and NUMA-aware shrinker for zswap, that is initiated when there is memory pressure. The shrinker does not have any parameter that must be tuned by the user, and can be opted in or out on a per-memcg basis. Furthermore, to make it more robust for many workloads and prevent overshrinking (i.e evicting warm pages that might be refaulted into memory), we build in the following heuristics: * Estimate the number of warm pages residing in zswap, and attempt to protect this region of the zswap LRU. * Scale the number of freeable objects by an estimate of the memory saving factor. The better zswap compresses the data, the fewer pages we will evict to swap (as we will otherwise incur IO for relatively small memory saving). * During reclaim, if the shrinker encounters a page that is also being brought into memory, the shrinker will cautiously terminate its shrinking action, as this is a sign that it is touching the warmer region of the zswap LRU. As a proof of concept, we ran the following synthetic benchmark: build the linux kernel in a memory-limited cgroup, and allocate some cold data in tmpfs to see if the shrinker could write them out and improved the overall performance. Depending on the amount of cold data generated, we observe from 14% to 35% reduction in kernel CPU time used in the kernel builds. [nphamcs@gmail.com: check shrinker enablement early, use less costly stat flushing] Link: https://lkml.kernel.org/r/20231206194456.3234203-1-nphamcs@gmail.com Link: https://lkml.kernel.org/r/20231130194023.4102148-7-nphamcs@gmail.com Signed-off-by: Nhat Pham <nphamcs@gmail.com> Acked-by: Johannes Weiner <hannes@cmpxchg.org> Tested-by: Bagas Sanjaya <bagasdotme@gmail.com> Cc: Chris Li <chrisl@kernel.org> Cc: Dan Streetman <ddstreet@ieee.org> Cc: Domenico Cerasuolo <cerasuolodomenico@gmail.com> Cc: Michal Hocko <mhocko@kernel.org> Cc: Muchun Song <muchun.song@linux.dev> Cc: Roman Gushchin <roman.gushchin@linux.dev> Cc: Seth Jennings <sjenning@redhat.com> Cc: Shakeel Butt <shakeelb@google.com> Cc: Shuah Khan <shuah@kernel.org> Cc: Vitaly Wool <vitaly.wool@konsulko.com> Cc: Yosry Ahmed <yosryahmed@google.com> Cc: Chengming Zhou <chengming.zhou@linux.dev> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2023-11-30 12:40:23 -07:00
config ZSWAP_SHRINKER_DEFAULT_ON
bool "Shrink the zswap pool on memory pressure"
depends on ZSWAP
default n
help
If selected, the zswap shrinker will be enabled, and the pages
stored in the zswap pool will become available for reclaim (i.e
written back to the backing swap device) on memory pressure.
This means that zswap writeback could happen even if the pool is
not yet full, or the cgroup zswap limit has not been reached,
reducing the chance that cold pages will reside in the zswap pool
and consume memory indefinitely.
choice
prompt "Default compressor"
depends on ZSWAP
default ZSWAP_COMPRESSOR_DEFAULT_LZO
help
Selects the default compression algorithm for the compressed cache
for swap pages.
For an overview what kind of performance can be expected from
a particular compression algorithm please refer to the benchmarks
available at the following LWN page:
https://lwn.net/Articles/751795/
If in doubt, select 'LZO'.
The selection made here can be overridden by using the kernel
command line 'zswap.compressor=' option.
config ZSWAP_COMPRESSOR_DEFAULT_DEFLATE
bool "Deflate"
select CRYPTO_DEFLATE
help
Use the Deflate algorithm as the default compression algorithm.
config ZSWAP_COMPRESSOR_DEFAULT_LZO
bool "LZO"
select CRYPTO_LZO
help
Use the LZO algorithm as the default compression algorithm.
config ZSWAP_COMPRESSOR_DEFAULT_842
bool "842"
select CRYPTO_842
help
Use the 842 algorithm as the default compression algorithm.
config ZSWAP_COMPRESSOR_DEFAULT_LZ4
bool "LZ4"
select CRYPTO_LZ4
help
Use the LZ4 algorithm as the default compression algorithm.
config ZSWAP_COMPRESSOR_DEFAULT_LZ4HC
bool "LZ4HC"
select CRYPTO_LZ4HC
help
Use the LZ4HC algorithm as the default compression algorithm.
config ZSWAP_COMPRESSOR_DEFAULT_ZSTD
bool "zstd"
select CRYPTO_ZSTD
help
Use the zstd algorithm as the default compression algorithm.
endchoice
config ZSWAP_COMPRESSOR_DEFAULT
string
depends on ZSWAP
default "deflate" if ZSWAP_COMPRESSOR_DEFAULT_DEFLATE
default "lzo" if ZSWAP_COMPRESSOR_DEFAULT_LZO
default "842" if ZSWAP_COMPRESSOR_DEFAULT_842
default "lz4" if ZSWAP_COMPRESSOR_DEFAULT_LZ4
default "lz4hc" if ZSWAP_COMPRESSOR_DEFAULT_LZ4HC
default "zstd" if ZSWAP_COMPRESSOR_DEFAULT_ZSTD
default ""
choice
prompt "Default allocator"
depends on ZSWAP
default ZSWAP_ZPOOL_DEFAULT_ZSMALLOC if MMU
default ZSWAP_ZPOOL_DEFAULT_ZBUD
help
Selects the default allocator for the compressed cache for
swap pages.
The default is 'zbud' for compatibility, however please do
read the description of each of the allocators below before
making a right choice.
The selection made here can be overridden by using the kernel
command line 'zswap.zpool=' option.
config ZSWAP_ZPOOL_DEFAULT_ZBUD
bool "zbud"
select ZBUD
help
Use the zbud allocator as the default allocator.
mm: z3fold: deprecate CONFIG_Z3FOLD The z3fold compressed pages allocator is rarely used, most users use zsmalloc. The only disadvantage of zsmalloc in comparison is the dependency on MMU, and zbud is a more common option for !MMU as it was the default zswap allocator for a long time. Historically, zsmalloc had worse latency than zbud and z3fold but offered better memory savings. This is no longer the case as shown by a simple recent analysis [1]. That analysis showed that z3fold does not have any advantage over zsmalloc or zbud considering both performance and memory usage. In a kernel build test on tmpfs in a limited cgroup, z3fold took 3% more time and used 1.8% more memory. The latency of zswap_load() was 7% higher, and that of zswap_store() was 10% higher. Zsmalloc is better in all metrics. Moreover, z3fold apparently has latent bugs, which was made noticeable by a recent soft lockup bug report with z3fold [2]. Switching to zsmalloc not only fixed the problem, but also reduced the swap usage from 6~8G to 1~2G. Other users have also reported being bitten by mistakenly enabling z3fold. Other than hurting users, z3fold is repeatedly causing wasted engineering effort. Apart from investigating the above bug, it came up in multiple development discussions (e.g. [3]) as something we need to handle, when there aren't any legit users (at least not intentionally). The natural course of action is to deprecate z3fold, and remove in a few cycles if no objections are raised from active users. Next on the list should be zbud, as it offers marginal latency gains at the cost of huge memory waste when compared to zsmalloc. That one will need to wait until zsmalloc does not depend on MMU. Rename the user-visible config option from CONFIG_Z3FOLD to CONFIG_Z3FOLD_DEPRECATED so that users with CONFIG_Z3FOLD=y get a new prompt with explanation during make oldconfig. Also, remove CONFIG_Z3FOLD=y from defconfigs. [1]https://lore.kernel.org/lkml/CAJD7tkbRF6od-2x_L8-A1QL3=2Ww13sCj4S3i4bNndqF+3+_Vg@mail.gmail.com/ [2]https://lore.kernel.org/lkml/EF0ABD3E-A239-4111-A8AB-5C442E759CF3@gmail.com/ [3]https://lore.kernel.org/lkml/CAJD7tkbnmeVugfunffSovJf9FAgy9rhBVt_tx=nxUveLUfqVsA@mail.gmail.com/ [arnd@arndb.de: deprecate ZSWAP_ZPOOL_DEFAULT_Z3FOLD as well] Link: https://lkml.kernel.org/r/20240909202625.1054880-1-arnd@kernel.org Link: https://lkml.kernel.org/r/20240904233343.933462-1-yosryahmed@google.com Signed-off-by: Yosry Ahmed <yosryahmed@google.com> Signed-off-by: Arnd Bergmann <arnd@arndb.de> Acked-by: Chris Down <chris@chrisdown.name> Acked-by: Nhat Pham <nphamcs@gmail.com> Acked-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Vitaly Wool <vitaly.wool@konsulko.com> Acked-by: Christoph Hellwig <hch@lst.de> Cc: Aneesh Kumar K.V <aneesh.kumar@kernel.org> Cc: Christophe Leroy <christophe.leroy@csgroup.eu> Cc: Huacai Chen <chenhuacai@kernel.org> Cc: Miaohe Lin <linmiaohe@huawei.com> Cc: Michael Ellerman <mpe@ellerman.id.au> Cc: Naveen N. Rao <naveen.n.rao@linux.ibm.com> Cc: Nicholas Piggin <npiggin@gmail.com> Cc: Sergey Senozhatsky <senozhatsky@chromium.org> Cc: WANG Xuerui <kernel@xen0n.name> Cc: <stable@vger.kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2024-09-04 16:33:43 -07:00
config ZSWAP_ZPOOL_DEFAULT_Z3FOLD_DEPRECATED
bool "z3foldi (DEPRECATED)"
select Z3FOLD_DEPRECATED
help
Use the z3fold allocator as the default allocator.
mm: z3fold: deprecate CONFIG_Z3FOLD The z3fold compressed pages allocator is rarely used, most users use zsmalloc. The only disadvantage of zsmalloc in comparison is the dependency on MMU, and zbud is a more common option for !MMU as it was the default zswap allocator for a long time. Historically, zsmalloc had worse latency than zbud and z3fold but offered better memory savings. This is no longer the case as shown by a simple recent analysis [1]. That analysis showed that z3fold does not have any advantage over zsmalloc or zbud considering both performance and memory usage. In a kernel build test on tmpfs in a limited cgroup, z3fold took 3% more time and used 1.8% more memory. The latency of zswap_load() was 7% higher, and that of zswap_store() was 10% higher. Zsmalloc is better in all metrics. Moreover, z3fold apparently has latent bugs, which was made noticeable by a recent soft lockup bug report with z3fold [2]. Switching to zsmalloc not only fixed the problem, but also reduced the swap usage from 6~8G to 1~2G. Other users have also reported being bitten by mistakenly enabling z3fold. Other than hurting users, z3fold is repeatedly causing wasted engineering effort. Apart from investigating the above bug, it came up in multiple development discussions (e.g. [3]) as something we need to handle, when there aren't any legit users (at least not intentionally). The natural course of action is to deprecate z3fold, and remove in a few cycles if no objections are raised from active users. Next on the list should be zbud, as it offers marginal latency gains at the cost of huge memory waste when compared to zsmalloc. That one will need to wait until zsmalloc does not depend on MMU. Rename the user-visible config option from CONFIG_Z3FOLD to CONFIG_Z3FOLD_DEPRECATED so that users with CONFIG_Z3FOLD=y get a new prompt with explanation during make oldconfig. Also, remove CONFIG_Z3FOLD=y from defconfigs. [1]https://lore.kernel.org/lkml/CAJD7tkbRF6od-2x_L8-A1QL3=2Ww13sCj4S3i4bNndqF+3+_Vg@mail.gmail.com/ [2]https://lore.kernel.org/lkml/EF0ABD3E-A239-4111-A8AB-5C442E759CF3@gmail.com/ [3]https://lore.kernel.org/lkml/CAJD7tkbnmeVugfunffSovJf9FAgy9rhBVt_tx=nxUveLUfqVsA@mail.gmail.com/ [arnd@arndb.de: deprecate ZSWAP_ZPOOL_DEFAULT_Z3FOLD as well] Link: https://lkml.kernel.org/r/20240909202625.1054880-1-arnd@kernel.org Link: https://lkml.kernel.org/r/20240904233343.933462-1-yosryahmed@google.com Signed-off-by: Yosry Ahmed <yosryahmed@google.com> Signed-off-by: Arnd Bergmann <arnd@arndb.de> Acked-by: Chris Down <chris@chrisdown.name> Acked-by: Nhat Pham <nphamcs@gmail.com> Acked-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Vitaly Wool <vitaly.wool@konsulko.com> Acked-by: Christoph Hellwig <hch@lst.de> Cc: Aneesh Kumar K.V <aneesh.kumar@kernel.org> Cc: Christophe Leroy <christophe.leroy@csgroup.eu> Cc: Huacai Chen <chenhuacai@kernel.org> Cc: Miaohe Lin <linmiaohe@huawei.com> Cc: Michael Ellerman <mpe@ellerman.id.au> Cc: Naveen N. Rao <naveen.n.rao@linux.ibm.com> Cc: Nicholas Piggin <npiggin@gmail.com> Cc: Sergey Senozhatsky <senozhatsky@chromium.org> Cc: WANG Xuerui <kernel@xen0n.name> Cc: <stable@vger.kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2024-09-04 16:33:43 -07:00
Deprecated and scheduled for removal in a few cycles,
see CONFIG_Z3FOLD_DEPRECATED.
config ZSWAP_ZPOOL_DEFAULT_ZSMALLOC
bool "zsmalloc"
select ZSMALLOC
help
Use the zsmalloc allocator as the default allocator.
endchoice
config ZSWAP_ZPOOL_DEFAULT
string
depends on ZSWAP
default "zbud" if ZSWAP_ZPOOL_DEFAULT_ZBUD
mm: z3fold: deprecate CONFIG_Z3FOLD The z3fold compressed pages allocator is rarely used, most users use zsmalloc. The only disadvantage of zsmalloc in comparison is the dependency on MMU, and zbud is a more common option for !MMU as it was the default zswap allocator for a long time. Historically, zsmalloc had worse latency than zbud and z3fold but offered better memory savings. This is no longer the case as shown by a simple recent analysis [1]. That analysis showed that z3fold does not have any advantage over zsmalloc or zbud considering both performance and memory usage. In a kernel build test on tmpfs in a limited cgroup, z3fold took 3% more time and used 1.8% more memory. The latency of zswap_load() was 7% higher, and that of zswap_store() was 10% higher. Zsmalloc is better in all metrics. Moreover, z3fold apparently has latent bugs, which was made noticeable by a recent soft lockup bug report with z3fold [2]. Switching to zsmalloc not only fixed the problem, but also reduced the swap usage from 6~8G to 1~2G. Other users have also reported being bitten by mistakenly enabling z3fold. Other than hurting users, z3fold is repeatedly causing wasted engineering effort. Apart from investigating the above bug, it came up in multiple development discussions (e.g. [3]) as something we need to handle, when there aren't any legit users (at least not intentionally). The natural course of action is to deprecate z3fold, and remove in a few cycles if no objections are raised from active users. Next on the list should be zbud, as it offers marginal latency gains at the cost of huge memory waste when compared to zsmalloc. That one will need to wait until zsmalloc does not depend on MMU. Rename the user-visible config option from CONFIG_Z3FOLD to CONFIG_Z3FOLD_DEPRECATED so that users with CONFIG_Z3FOLD=y get a new prompt with explanation during make oldconfig. Also, remove CONFIG_Z3FOLD=y from defconfigs. [1]https://lore.kernel.org/lkml/CAJD7tkbRF6od-2x_L8-A1QL3=2Ww13sCj4S3i4bNndqF+3+_Vg@mail.gmail.com/ [2]https://lore.kernel.org/lkml/EF0ABD3E-A239-4111-A8AB-5C442E759CF3@gmail.com/ [3]https://lore.kernel.org/lkml/CAJD7tkbnmeVugfunffSovJf9FAgy9rhBVt_tx=nxUveLUfqVsA@mail.gmail.com/ [arnd@arndb.de: deprecate ZSWAP_ZPOOL_DEFAULT_Z3FOLD as well] Link: https://lkml.kernel.org/r/20240909202625.1054880-1-arnd@kernel.org Link: https://lkml.kernel.org/r/20240904233343.933462-1-yosryahmed@google.com Signed-off-by: Yosry Ahmed <yosryahmed@google.com> Signed-off-by: Arnd Bergmann <arnd@arndb.de> Acked-by: Chris Down <chris@chrisdown.name> Acked-by: Nhat Pham <nphamcs@gmail.com> Acked-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Vitaly Wool <vitaly.wool@konsulko.com> Acked-by: Christoph Hellwig <hch@lst.de> Cc: Aneesh Kumar K.V <aneesh.kumar@kernel.org> Cc: Christophe Leroy <christophe.leroy@csgroup.eu> Cc: Huacai Chen <chenhuacai@kernel.org> Cc: Miaohe Lin <linmiaohe@huawei.com> Cc: Michael Ellerman <mpe@ellerman.id.au> Cc: Naveen N. Rao <naveen.n.rao@linux.ibm.com> Cc: Nicholas Piggin <npiggin@gmail.com> Cc: Sergey Senozhatsky <senozhatsky@chromium.org> Cc: WANG Xuerui <kernel@xen0n.name> Cc: <stable@vger.kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2024-09-04 16:33:43 -07:00
default "z3fold" if ZSWAP_ZPOOL_DEFAULT_Z3FOLD_DEPRECATED
default "zsmalloc" if ZSWAP_ZPOOL_DEFAULT_ZSMALLOC
default ""
config ZBUD
tristate "2:1 compression allocator (zbud)"
depends on ZSWAP
help
A special purpose allocator for storing compressed pages.
It is designed to store up to two compressed pages per physical
page. While this design limits storage density, it has simple and
deterministic reclaim properties that make it preferable to a higher
density approach when reclaim will be used.
mm: z3fold: deprecate CONFIG_Z3FOLD The z3fold compressed pages allocator is rarely used, most users use zsmalloc. The only disadvantage of zsmalloc in comparison is the dependency on MMU, and zbud is a more common option for !MMU as it was the default zswap allocator for a long time. Historically, zsmalloc had worse latency than zbud and z3fold but offered better memory savings. This is no longer the case as shown by a simple recent analysis [1]. That analysis showed that z3fold does not have any advantage over zsmalloc or zbud considering both performance and memory usage. In a kernel build test on tmpfs in a limited cgroup, z3fold took 3% more time and used 1.8% more memory. The latency of zswap_load() was 7% higher, and that of zswap_store() was 10% higher. Zsmalloc is better in all metrics. Moreover, z3fold apparently has latent bugs, which was made noticeable by a recent soft lockup bug report with z3fold [2]. Switching to zsmalloc not only fixed the problem, but also reduced the swap usage from 6~8G to 1~2G. Other users have also reported being bitten by mistakenly enabling z3fold. Other than hurting users, z3fold is repeatedly causing wasted engineering effort. Apart from investigating the above bug, it came up in multiple development discussions (e.g. [3]) as something we need to handle, when there aren't any legit users (at least not intentionally). The natural course of action is to deprecate z3fold, and remove in a few cycles if no objections are raised from active users. Next on the list should be zbud, as it offers marginal latency gains at the cost of huge memory waste when compared to zsmalloc. That one will need to wait until zsmalloc does not depend on MMU. Rename the user-visible config option from CONFIG_Z3FOLD to CONFIG_Z3FOLD_DEPRECATED so that users with CONFIG_Z3FOLD=y get a new prompt with explanation during make oldconfig. Also, remove CONFIG_Z3FOLD=y from defconfigs. [1]https://lore.kernel.org/lkml/CAJD7tkbRF6od-2x_L8-A1QL3=2Ww13sCj4S3i4bNndqF+3+_Vg@mail.gmail.com/ [2]https://lore.kernel.org/lkml/EF0ABD3E-A239-4111-A8AB-5C442E759CF3@gmail.com/ [3]https://lore.kernel.org/lkml/CAJD7tkbnmeVugfunffSovJf9FAgy9rhBVt_tx=nxUveLUfqVsA@mail.gmail.com/ [arnd@arndb.de: deprecate ZSWAP_ZPOOL_DEFAULT_Z3FOLD as well] Link: https://lkml.kernel.org/r/20240909202625.1054880-1-arnd@kernel.org Link: https://lkml.kernel.org/r/20240904233343.933462-1-yosryahmed@google.com Signed-off-by: Yosry Ahmed <yosryahmed@google.com> Signed-off-by: Arnd Bergmann <arnd@arndb.de> Acked-by: Chris Down <chris@chrisdown.name> Acked-by: Nhat Pham <nphamcs@gmail.com> Acked-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Vitaly Wool <vitaly.wool@konsulko.com> Acked-by: Christoph Hellwig <hch@lst.de> Cc: Aneesh Kumar K.V <aneesh.kumar@kernel.org> Cc: Christophe Leroy <christophe.leroy@csgroup.eu> Cc: Huacai Chen <chenhuacai@kernel.org> Cc: Miaohe Lin <linmiaohe@huawei.com> Cc: Michael Ellerman <mpe@ellerman.id.au> Cc: Naveen N. Rao <naveen.n.rao@linux.ibm.com> Cc: Nicholas Piggin <npiggin@gmail.com> Cc: Sergey Senozhatsky <senozhatsky@chromium.org> Cc: WANG Xuerui <kernel@xen0n.name> Cc: <stable@vger.kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2024-09-04 16:33:43 -07:00
config Z3FOLD_DEPRECATED
tristate "3:1 compression allocator (z3fold) (DEPRECATED)"
depends on ZSWAP
help
mm: z3fold: deprecate CONFIG_Z3FOLD The z3fold compressed pages allocator is rarely used, most users use zsmalloc. The only disadvantage of zsmalloc in comparison is the dependency on MMU, and zbud is a more common option for !MMU as it was the default zswap allocator for a long time. Historically, zsmalloc had worse latency than zbud and z3fold but offered better memory savings. This is no longer the case as shown by a simple recent analysis [1]. That analysis showed that z3fold does not have any advantage over zsmalloc or zbud considering both performance and memory usage. In a kernel build test on tmpfs in a limited cgroup, z3fold took 3% more time and used 1.8% more memory. The latency of zswap_load() was 7% higher, and that of zswap_store() was 10% higher. Zsmalloc is better in all metrics. Moreover, z3fold apparently has latent bugs, which was made noticeable by a recent soft lockup bug report with z3fold [2]. Switching to zsmalloc not only fixed the problem, but also reduced the swap usage from 6~8G to 1~2G. Other users have also reported being bitten by mistakenly enabling z3fold. Other than hurting users, z3fold is repeatedly causing wasted engineering effort. Apart from investigating the above bug, it came up in multiple development discussions (e.g. [3]) as something we need to handle, when there aren't any legit users (at least not intentionally). The natural course of action is to deprecate z3fold, and remove in a few cycles if no objections are raised from active users. Next on the list should be zbud, as it offers marginal latency gains at the cost of huge memory waste when compared to zsmalloc. That one will need to wait until zsmalloc does not depend on MMU. Rename the user-visible config option from CONFIG_Z3FOLD to CONFIG_Z3FOLD_DEPRECATED so that users with CONFIG_Z3FOLD=y get a new prompt with explanation during make oldconfig. Also, remove CONFIG_Z3FOLD=y from defconfigs. [1]https://lore.kernel.org/lkml/CAJD7tkbRF6od-2x_L8-A1QL3=2Ww13sCj4S3i4bNndqF+3+_Vg@mail.gmail.com/ [2]https://lore.kernel.org/lkml/EF0ABD3E-A239-4111-A8AB-5C442E759CF3@gmail.com/ [3]https://lore.kernel.org/lkml/CAJD7tkbnmeVugfunffSovJf9FAgy9rhBVt_tx=nxUveLUfqVsA@mail.gmail.com/ [arnd@arndb.de: deprecate ZSWAP_ZPOOL_DEFAULT_Z3FOLD as well] Link: https://lkml.kernel.org/r/20240909202625.1054880-1-arnd@kernel.org Link: https://lkml.kernel.org/r/20240904233343.933462-1-yosryahmed@google.com Signed-off-by: Yosry Ahmed <yosryahmed@google.com> Signed-off-by: Arnd Bergmann <arnd@arndb.de> Acked-by: Chris Down <chris@chrisdown.name> Acked-by: Nhat Pham <nphamcs@gmail.com> Acked-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Vitaly Wool <vitaly.wool@konsulko.com> Acked-by: Christoph Hellwig <hch@lst.de> Cc: Aneesh Kumar K.V <aneesh.kumar@kernel.org> Cc: Christophe Leroy <christophe.leroy@csgroup.eu> Cc: Huacai Chen <chenhuacai@kernel.org> Cc: Miaohe Lin <linmiaohe@huawei.com> Cc: Michael Ellerman <mpe@ellerman.id.au> Cc: Naveen N. Rao <naveen.n.rao@linux.ibm.com> Cc: Nicholas Piggin <npiggin@gmail.com> Cc: Sergey Senozhatsky <senozhatsky@chromium.org> Cc: WANG Xuerui <kernel@xen0n.name> Cc: <stable@vger.kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2024-09-04 16:33:43 -07:00
Deprecated and scheduled for removal in a few cycles. If you have
a good reason for using Z3FOLD over ZSMALLOC, please contact
linux-mm@kvack.org and the zswap maintainers.
A special purpose allocator for storing compressed pages.
It is designed to store up to three compressed pages per physical
page. It is a ZBUD derivative so the simplicity and determinism are
still there.
mm: z3fold: deprecate CONFIG_Z3FOLD The z3fold compressed pages allocator is rarely used, most users use zsmalloc. The only disadvantage of zsmalloc in comparison is the dependency on MMU, and zbud is a more common option for !MMU as it was the default zswap allocator for a long time. Historically, zsmalloc had worse latency than zbud and z3fold but offered better memory savings. This is no longer the case as shown by a simple recent analysis [1]. That analysis showed that z3fold does not have any advantage over zsmalloc or zbud considering both performance and memory usage. In a kernel build test on tmpfs in a limited cgroup, z3fold took 3% more time and used 1.8% more memory. The latency of zswap_load() was 7% higher, and that of zswap_store() was 10% higher. Zsmalloc is better in all metrics. Moreover, z3fold apparently has latent bugs, which was made noticeable by a recent soft lockup bug report with z3fold [2]. Switching to zsmalloc not only fixed the problem, but also reduced the swap usage from 6~8G to 1~2G. Other users have also reported being bitten by mistakenly enabling z3fold. Other than hurting users, z3fold is repeatedly causing wasted engineering effort. Apart from investigating the above bug, it came up in multiple development discussions (e.g. [3]) as something we need to handle, when there aren't any legit users (at least not intentionally). The natural course of action is to deprecate z3fold, and remove in a few cycles if no objections are raised from active users. Next on the list should be zbud, as it offers marginal latency gains at the cost of huge memory waste when compared to zsmalloc. That one will need to wait until zsmalloc does not depend on MMU. Rename the user-visible config option from CONFIG_Z3FOLD to CONFIG_Z3FOLD_DEPRECATED so that users with CONFIG_Z3FOLD=y get a new prompt with explanation during make oldconfig. Also, remove CONFIG_Z3FOLD=y from defconfigs. [1]https://lore.kernel.org/lkml/CAJD7tkbRF6od-2x_L8-A1QL3=2Ww13sCj4S3i4bNndqF+3+_Vg@mail.gmail.com/ [2]https://lore.kernel.org/lkml/EF0ABD3E-A239-4111-A8AB-5C442E759CF3@gmail.com/ [3]https://lore.kernel.org/lkml/CAJD7tkbnmeVugfunffSovJf9FAgy9rhBVt_tx=nxUveLUfqVsA@mail.gmail.com/ [arnd@arndb.de: deprecate ZSWAP_ZPOOL_DEFAULT_Z3FOLD as well] Link: https://lkml.kernel.org/r/20240909202625.1054880-1-arnd@kernel.org Link: https://lkml.kernel.org/r/20240904233343.933462-1-yosryahmed@google.com Signed-off-by: Yosry Ahmed <yosryahmed@google.com> Signed-off-by: Arnd Bergmann <arnd@arndb.de> Acked-by: Chris Down <chris@chrisdown.name> Acked-by: Nhat Pham <nphamcs@gmail.com> Acked-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Vitaly Wool <vitaly.wool@konsulko.com> Acked-by: Christoph Hellwig <hch@lst.de> Cc: Aneesh Kumar K.V <aneesh.kumar@kernel.org> Cc: Christophe Leroy <christophe.leroy@csgroup.eu> Cc: Huacai Chen <chenhuacai@kernel.org> Cc: Miaohe Lin <linmiaohe@huawei.com> Cc: Michael Ellerman <mpe@ellerman.id.au> Cc: Naveen N. Rao <naveen.n.rao@linux.ibm.com> Cc: Nicholas Piggin <npiggin@gmail.com> Cc: Sergey Senozhatsky <senozhatsky@chromium.org> Cc: WANG Xuerui <kernel@xen0n.name> Cc: <stable@vger.kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2024-09-04 16:33:43 -07:00
config Z3FOLD
tristate
default y if Z3FOLD_DEPRECATED=y
default m if Z3FOLD_DEPRECATED=m
depends on Z3FOLD_DEPRECATED
config ZSMALLOC
tristate
prompt "N:1 compression allocator (zsmalloc)" if (ZSWAP || ZRAM)
depends on MMU
help
zsmalloc is a slab-based memory allocator designed to store
pages of various compression levels efficiently. It achieves
the highest storage density with the least amount of fragmentation.
config ZSMALLOC_STAT
bool "Export zsmalloc statistics"
depends on ZSMALLOC
select DEBUG_FS
help
This option enables code in the zsmalloc to collect various
statistics about what's happening in zsmalloc and exports that
information to userspace via debugfs.
If unsure, say N.
config ZSMALLOC_CHAIN_SIZE
int "Maximum number of physical pages per-zspage"
default 8
range 4 16
depends on ZSMALLOC
help
This option sets the upper limit on the number of physical pages
that a zmalloc page (zspage) can consist of. The optimal zspage
chain size is calculated for each size class during the
initialization of the pool.
Changing this option can alter the characteristics of size classes,
such as the number of pages per zspage and the number of objects
per zspage. This can also result in different configurations of
the pool, as zsmalloc merges size classes with similar
characteristics.
For more information, see zsmalloc documentation.
menu "Slab allocator options"
config SLUB
def_bool y
config SLUB_TINY
bool "Configure for minimal memory footprint"
depends on EXPERT
select SLAB_MERGE_DEFAULT
help
Configures the slab allocator in a way to achieve minimal memory
footprint, sacrificing scalability, debugging and other features.
This is intended only for the smallest system that had used the
SLOB allocator and is not recommended for systems with more than
16MB RAM.
If unsure, say N.
config SLAB_MERGE_DEFAULT
bool "Allow slab caches to be merged"
default y
help
For reduced kernel memory fragmentation, slab caches can be
merged when they share the same size and other characteristics.
This carries a risk of kernel heap overflows being able to
overwrite objects from merged caches (and more easily control
cache layout), which makes such heap attacks easier to exploit
by attackers. By keeping caches unmerged, these kinds of exploits
can usually only damage objects in the same cache. To disable
merging at runtime, "slab_nomerge" can be passed on the kernel
command line.
config SLAB_FREELIST_RANDOM
bool "Randomize slab freelist"
depends on !SLUB_TINY
help
Randomizes the freelist order used on creating new pages. This
security feature reduces the predictability of the kernel slab
allocator against heap overflows.
config SLAB_FREELIST_HARDENED
bool "Harden slab freelist metadata"
depends on !SLUB_TINY
help
Many kernel heap attacks try to target slab cache metadata and
other infrastructure. This options makes minor performance
sacrifices to harden the kernel slab allocator against common
freelist exploit methods.
config SLAB_BUCKETS
bool "Support allocation from separate kmalloc buckets"
depends on !SLUB_TINY
default SLAB_FREELIST_HARDENED
help
Kernel heap attacks frequently depend on being able to create
specifically-sized allocations with user-controlled contents
that will be allocated into the same kmalloc bucket as a
target object. To avoid sharing these allocation buckets,
provide an explicitly separated set of buckets to be used for
user-controlled allocations. This may very slightly increase
memory fragmentation, though in practice it's only a handful
of extra pages since the bulk of user-controlled allocations
are relatively long-lived.
If unsure, say Y.
config SLUB_STATS
default n
bool "Enable performance statistics"
depends on SYSFS && !SLUB_TINY
help
The statistics are useful to debug slab allocation behavior in
order find ways to optimize the allocator. This should never be
enabled for production use since keeping statistics slows down
the allocator by a few percentage points. The slabinfo command
supports the determination of the most active slabs to figure
out which slabs are relevant to a particular load.
Try running: slabinfo -DA
config SLUB_CPU_PARTIAL
default y
depends on SMP && !SLUB_TINY
bool "Enable per cpu partial caches"
help
Per cpu partial caches accelerate objects allocation and freeing
that is local to a processor at the price of more indeterminism
in the latency of the free. On overflow these caches will be cleared
which requires the taking of locks that may cause latency spikes.
Typically one would choose no for a realtime system.
Randomized slab caches for kmalloc() When exploiting memory vulnerabilities, "heap spraying" is a common technique targeting those related to dynamic memory allocation (i.e. the "heap"), and it plays an important role in a successful exploitation. Basically, it is to overwrite the memory area of vulnerable object by triggering allocation in other subsystems or modules and therefore getting a reference to the targeted memory location. It's usable on various types of vulnerablity including use after free (UAF), heap out- of-bound write and etc. There are (at least) two reasons why the heap can be sprayed: 1) generic slab caches are shared among different subsystems and modules, and 2) dedicated slab caches could be merged with the generic ones. Currently these two factors cannot be prevented at a low cost: the first one is a widely used memory allocation mechanism, and shutting down slab merging completely via `slub_nomerge` would be overkill. To efficiently prevent heap spraying, we propose the following approach: to create multiple copies of generic slab caches that will never be merged, and random one of them will be used at allocation. The random selection is based on the address of code that calls `kmalloc()`, which means it is static at runtime (rather than dynamically determined at each time of allocation, which could be bypassed by repeatedly spraying in brute force). In other words, the randomness of cache selection will be with respect to the code address rather than time, i.e. allocations in different code paths would most likely pick different caches, although kmalloc() at each place would use the same cache copy whenever it is executed. In this way, the vulnerable object and memory allocated in other subsystems and modules will (most probably) be on different slab caches, which prevents the object from being sprayed. Meanwhile, the static random selection is further enhanced with a per-boot random seed, which prevents the attacker from finding a usable kmalloc that happens to pick the same cache with the vulnerable subsystem/module by analyzing the open source code. In other words, with the per-boot seed, the random selection is static during each time the system starts and runs, but not across different system startups. The overhead of performance has been tested on a 40-core x86 server by comparing the results of `perf bench all` between the kernels with and without this patch based on the latest linux-next kernel, which shows minor difference. A subset of benchmarks are listed below: sched/ sched/ syscall/ mem/ mem/ messaging pipe basic memcpy memset (sec) (sec) (sec) (GB/sec) (GB/sec) control1 0.019 5.459 0.733 15.258789 51.398026 control2 0.019 5.439 0.730 16.009221 48.828125 control3 0.019 5.282 0.735 16.009221 48.828125 control_avg 0.019 5.393 0.733 15.759077 49.684759 experiment1 0.019 5.374 0.741 15.500992 46.502976 experiment2 0.019 5.440 0.746 16.276042 51.398026 experiment3 0.019 5.242 0.752 15.258789 51.398026 experiment_avg 0.019 5.352 0.746 15.678608 49.766343 The overhead of memory usage was measured by executing `free` after boot on a QEMU VM with 1GB total memory, and as expected, it's positively correlated with # of cache copies: control 4 copies 8 copies 16 copies total 969.8M 968.2M 968.2M 968.2M used 20.0M 21.9M 24.1M 26.7M free 936.9M 933.6M 931.4M 928.6M available 932.2M 928.8M 926.6M 923.9M Co-developed-by: Xiu Jianfeng <xiujianfeng@huawei.com> Signed-off-by: Xiu Jianfeng <xiujianfeng@huawei.com> Signed-off-by: GONG, Ruiqi <gongruiqi@huaweicloud.com> Reviewed-by: Kees Cook <keescook@chromium.org> Reviewed-by: Hyeonggon Yoo <42.hyeyoo@gmail.com> Acked-by: Dennis Zhou <dennis@kernel.org> # percpu Signed-off-by: Vlastimil Babka <vbabka@suse.cz>
2023-07-13 23:44:22 -07:00
config RANDOM_KMALLOC_CACHES
default n
depends on !SLUB_TINY
Randomized slab caches for kmalloc() When exploiting memory vulnerabilities, "heap spraying" is a common technique targeting those related to dynamic memory allocation (i.e. the "heap"), and it plays an important role in a successful exploitation. Basically, it is to overwrite the memory area of vulnerable object by triggering allocation in other subsystems or modules and therefore getting a reference to the targeted memory location. It's usable on various types of vulnerablity including use after free (UAF), heap out- of-bound write and etc. There are (at least) two reasons why the heap can be sprayed: 1) generic slab caches are shared among different subsystems and modules, and 2) dedicated slab caches could be merged with the generic ones. Currently these two factors cannot be prevented at a low cost: the first one is a widely used memory allocation mechanism, and shutting down slab merging completely via `slub_nomerge` would be overkill. To efficiently prevent heap spraying, we propose the following approach: to create multiple copies of generic slab caches that will never be merged, and random one of them will be used at allocation. The random selection is based on the address of code that calls `kmalloc()`, which means it is static at runtime (rather than dynamically determined at each time of allocation, which could be bypassed by repeatedly spraying in brute force). In other words, the randomness of cache selection will be with respect to the code address rather than time, i.e. allocations in different code paths would most likely pick different caches, although kmalloc() at each place would use the same cache copy whenever it is executed. In this way, the vulnerable object and memory allocated in other subsystems and modules will (most probably) be on different slab caches, which prevents the object from being sprayed. Meanwhile, the static random selection is further enhanced with a per-boot random seed, which prevents the attacker from finding a usable kmalloc that happens to pick the same cache with the vulnerable subsystem/module by analyzing the open source code. In other words, with the per-boot seed, the random selection is static during each time the system starts and runs, but not across different system startups. The overhead of performance has been tested on a 40-core x86 server by comparing the results of `perf bench all` between the kernels with and without this patch based on the latest linux-next kernel, which shows minor difference. A subset of benchmarks are listed below: sched/ sched/ syscall/ mem/ mem/ messaging pipe basic memcpy memset (sec) (sec) (sec) (GB/sec) (GB/sec) control1 0.019 5.459 0.733 15.258789 51.398026 control2 0.019 5.439 0.730 16.009221 48.828125 control3 0.019 5.282 0.735 16.009221 48.828125 control_avg 0.019 5.393 0.733 15.759077 49.684759 experiment1 0.019 5.374 0.741 15.500992 46.502976 experiment2 0.019 5.440 0.746 16.276042 51.398026 experiment3 0.019 5.242 0.752 15.258789 51.398026 experiment_avg 0.019 5.352 0.746 15.678608 49.766343 The overhead of memory usage was measured by executing `free` after boot on a QEMU VM with 1GB total memory, and as expected, it's positively correlated with # of cache copies: control 4 copies 8 copies 16 copies total 969.8M 968.2M 968.2M 968.2M used 20.0M 21.9M 24.1M 26.7M free 936.9M 933.6M 931.4M 928.6M available 932.2M 928.8M 926.6M 923.9M Co-developed-by: Xiu Jianfeng <xiujianfeng@huawei.com> Signed-off-by: Xiu Jianfeng <xiujianfeng@huawei.com> Signed-off-by: GONG, Ruiqi <gongruiqi@huaweicloud.com> Reviewed-by: Kees Cook <keescook@chromium.org> Reviewed-by: Hyeonggon Yoo <42.hyeyoo@gmail.com> Acked-by: Dennis Zhou <dennis@kernel.org> # percpu Signed-off-by: Vlastimil Babka <vbabka@suse.cz>
2023-07-13 23:44:22 -07:00
bool "Randomize slab caches for normal kmalloc"
help
A hardening feature that creates multiple copies of slab caches for
normal kmalloc allocation and makes kmalloc randomly pick one based
on code address, which makes the attackers more difficult to spray
vulnerable memory objects on the heap for the purpose of exploiting
memory vulnerabilities.
Currently the number of copies is set to 16, a reasonably large value
that effectively diverges the memory objects allocated for different
subsystems or modules into different caches, at the expense of a
limited degree of memory and CPU overhead that relates to hardware and
system workload.
endmenu # Slab allocator options
config SHUFFLE_PAGE_ALLOCATOR
bool "Page allocator randomization"
default SLAB_FREELIST_RANDOM && ACPI_NUMA
help
Randomization of the page allocator improves the average
utilization of a direct-mapped memory-side-cache. See section
5.2.27 Heterogeneous Memory Attribute Table (HMAT) in the ACPI
6.2a specification for an example of how a platform advertises
the presence of a memory-side-cache. There are also incidental
security benefits as it reduces the predictability of page
allocations to compliment SLAB_FREELIST_RANDOM, but the
default granularity of shuffling on the MAX_PAGE_ORDER i.e, 10th
order of pages is selected based on cache utilization benefits
on x86.
While the randomization improves cache utilization it may
negatively impact workloads on platforms without a cache. For
this reason, by default, the randomization is not enabled even
if SHUFFLE_PAGE_ALLOCATOR=y. The randomization may be force enabled
with the 'page_alloc.shuffle' kernel command line parameter.
Say Y if unsure.
config COMPAT_BRK
bool "Disable heap randomization"
default y
help
Randomizing heap placement makes heap exploits harder, but it
also breaks ancient binaries (including anything libc5 based).
This option changes the bootup default to heap randomization
disabled, and can be overridden at runtime by setting
/proc/sys/kernel/randomize_va_space to 2.
On non-ancient distros (post-2000 ones) N is usually a safe choice.
config MMAP_ALLOW_UNINITIALIZED
bool "Allow mmapped anonymous memory to be uninitialized"
depends on EXPERT && !MMU
default n
help
Normally, and according to the Linux spec, anonymous memory obtained
from mmap() has its contents cleared before it is passed to
userspace. Enabling this config option allows you to request that
mmap() skip that if it is given an MAP_UNINITIALIZED flag, thus
providing a huge performance boost. If this option is not enabled,
then the flag will be ignored.
This is taken advantage of by uClibc's malloc(), and also by
ELF-FDPIC binfmt's brk and stack allocator.
Because of the obvious security issues, this option should only be
enabled on embedded devices where you control what is run in
userspace. Since that isn't generally a problem on no-MMU systems,
it is normally safe to say Y here.
See Documentation/admin-guide/mm/nommu-mmap.rst for more information.
config SELECT_MEMORY_MODEL
def_bool y
depends on ARCH_SELECT_MEMORY_MODEL
choice
prompt "Memory model"
depends on SELECT_MEMORY_MODEL
[PATCH] sparsemem memory model Sparsemem abstracts the use of discontiguous mem_maps[]. This kind of mem_map[] is needed by discontiguous memory machines (like in the old CONFIG_DISCONTIGMEM case) as well as memory hotplug systems. Sparsemem replaces DISCONTIGMEM when enabled, and it is hoped that it can eventually become a complete replacement. A significant advantage over DISCONTIGMEM is that it's completely separated from CONFIG_NUMA. When producing this patch, it became apparent in that NUMA and DISCONTIG are often confused. Another advantage is that sparse doesn't require each NUMA node's ranges to be contiguous. It can handle overlapping ranges between nodes with no problems, where DISCONTIGMEM currently throws away that memory. Sparsemem uses an array to provide different pfn_to_page() translations for each SECTION_SIZE area of physical memory. This is what allows the mem_map[] to be chopped up. In order to do quick pfn_to_page() operations, the section number of the page is encoded in page->flags. Part of the sparsemem infrastructure enables sharing of these bits more dynamically (at compile-time) between the page_zone() and sparsemem operations. However, on 32-bit architectures, the number of bits is quite limited, and may require growing the size of the page->flags type in certain conditions. Several things might force this to occur: a decrease in the SECTION_SIZE (if you want to hotplug smaller areas of memory), an increase in the physical address space, or an increase in the number of used page->flags. One thing to note is that, once sparsemem is present, the NUMA node information no longer needs to be stored in the page->flags. It might provide speed increases on certain platforms and will be stored there if there is room. But, if out of room, an alternate (theoretically slower) mechanism is used. This patch introduces CONFIG_FLATMEM. It is used in almost all cases where there used to be an #ifndef DISCONTIG, because SPARSEMEM and DISCONTIGMEM often have to compile out the same areas of code. Signed-off-by: Andy Whitcroft <apw@shadowen.org> Signed-off-by: Dave Hansen <haveblue@us.ibm.com> Signed-off-by: Martin Bligh <mbligh@aracnet.com> Signed-off-by: Adrian Bunk <bunk@stusta.de> Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com> Signed-off-by: Bob Picco <bob.picco@hp.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 00:07:54 -07:00
default SPARSEMEM_MANUAL if ARCH_SPARSEMEM_DEFAULT
default FLATMEM_MANUAL
help
This option allows you to change some of the ways that
Linux manages its memory internally. Most users will
only have one option here selected by the architecture
configuration. This is normal.
config FLATMEM_MANUAL
bool "Flat Memory"
depends on !ARCH_SPARSEMEM_ENABLE || ARCH_FLATMEM_ENABLE
help
This option is best suited for non-NUMA systems with
flat address space. The FLATMEM is the most efficient
system in terms of performance and resource consumption
and it is the best option for smaller systems.
For systems that have holes in their physical address
spaces and for features like NUMA and memory hotplug,
choose "Sparse Memory".
[PATCH] sparsemem memory model Sparsemem abstracts the use of discontiguous mem_maps[]. This kind of mem_map[] is needed by discontiguous memory machines (like in the old CONFIG_DISCONTIGMEM case) as well as memory hotplug systems. Sparsemem replaces DISCONTIGMEM when enabled, and it is hoped that it can eventually become a complete replacement. A significant advantage over DISCONTIGMEM is that it's completely separated from CONFIG_NUMA. When producing this patch, it became apparent in that NUMA and DISCONTIG are often confused. Another advantage is that sparse doesn't require each NUMA node's ranges to be contiguous. It can handle overlapping ranges between nodes with no problems, where DISCONTIGMEM currently throws away that memory. Sparsemem uses an array to provide different pfn_to_page() translations for each SECTION_SIZE area of physical memory. This is what allows the mem_map[] to be chopped up. In order to do quick pfn_to_page() operations, the section number of the page is encoded in page->flags. Part of the sparsemem infrastructure enables sharing of these bits more dynamically (at compile-time) between the page_zone() and sparsemem operations. However, on 32-bit architectures, the number of bits is quite limited, and may require growing the size of the page->flags type in certain conditions. Several things might force this to occur: a decrease in the SECTION_SIZE (if you want to hotplug smaller areas of memory), an increase in the physical address space, or an increase in the number of used page->flags. One thing to note is that, once sparsemem is present, the NUMA node information no longer needs to be stored in the page->flags. It might provide speed increases on certain platforms and will be stored there if there is room. But, if out of room, an alternate (theoretically slower) mechanism is used. This patch introduces CONFIG_FLATMEM. It is used in almost all cases where there used to be an #ifndef DISCONTIG, because SPARSEMEM and DISCONTIGMEM often have to compile out the same areas of code. Signed-off-by: Andy Whitcroft <apw@shadowen.org> Signed-off-by: Dave Hansen <haveblue@us.ibm.com> Signed-off-by: Martin Bligh <mbligh@aracnet.com> Signed-off-by: Adrian Bunk <bunk@stusta.de> Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com> Signed-off-by: Bob Picco <bob.picco@hp.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 00:07:54 -07:00
If unsure, choose this option (Flat Memory) over any other.
[PATCH] sparsemem memory model Sparsemem abstracts the use of discontiguous mem_maps[]. This kind of mem_map[] is needed by discontiguous memory machines (like in the old CONFIG_DISCONTIGMEM case) as well as memory hotplug systems. Sparsemem replaces DISCONTIGMEM when enabled, and it is hoped that it can eventually become a complete replacement. A significant advantage over DISCONTIGMEM is that it's completely separated from CONFIG_NUMA. When producing this patch, it became apparent in that NUMA and DISCONTIG are often confused. Another advantage is that sparse doesn't require each NUMA node's ranges to be contiguous. It can handle overlapping ranges between nodes with no problems, where DISCONTIGMEM currently throws away that memory. Sparsemem uses an array to provide different pfn_to_page() translations for each SECTION_SIZE area of physical memory. This is what allows the mem_map[] to be chopped up. In order to do quick pfn_to_page() operations, the section number of the page is encoded in page->flags. Part of the sparsemem infrastructure enables sharing of these bits more dynamically (at compile-time) between the page_zone() and sparsemem operations. However, on 32-bit architectures, the number of bits is quite limited, and may require growing the size of the page->flags type in certain conditions. Several things might force this to occur: a decrease in the SECTION_SIZE (if you want to hotplug smaller areas of memory), an increase in the physical address space, or an increase in the number of used page->flags. One thing to note is that, once sparsemem is present, the NUMA node information no longer needs to be stored in the page->flags. It might provide speed increases on certain platforms and will be stored there if there is room. But, if out of room, an alternate (theoretically slower) mechanism is used. This patch introduces CONFIG_FLATMEM. It is used in almost all cases where there used to be an #ifndef DISCONTIG, because SPARSEMEM and DISCONTIGMEM often have to compile out the same areas of code. Signed-off-by: Andy Whitcroft <apw@shadowen.org> Signed-off-by: Dave Hansen <haveblue@us.ibm.com> Signed-off-by: Martin Bligh <mbligh@aracnet.com> Signed-off-by: Adrian Bunk <bunk@stusta.de> Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com> Signed-off-by: Bob Picco <bob.picco@hp.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 00:07:54 -07:00
config SPARSEMEM_MANUAL
bool "Sparse Memory"
depends on ARCH_SPARSEMEM_ENABLE
help
This will be the only option for some systems, including
memory hot-plug systems. This is normal.
[PATCH] sparsemem memory model Sparsemem abstracts the use of discontiguous mem_maps[]. This kind of mem_map[] is needed by discontiguous memory machines (like in the old CONFIG_DISCONTIGMEM case) as well as memory hotplug systems. Sparsemem replaces DISCONTIGMEM when enabled, and it is hoped that it can eventually become a complete replacement. A significant advantage over DISCONTIGMEM is that it's completely separated from CONFIG_NUMA. When producing this patch, it became apparent in that NUMA and DISCONTIG are often confused. Another advantage is that sparse doesn't require each NUMA node's ranges to be contiguous. It can handle overlapping ranges between nodes with no problems, where DISCONTIGMEM currently throws away that memory. Sparsemem uses an array to provide different pfn_to_page() translations for each SECTION_SIZE area of physical memory. This is what allows the mem_map[] to be chopped up. In order to do quick pfn_to_page() operations, the section number of the page is encoded in page->flags. Part of the sparsemem infrastructure enables sharing of these bits more dynamically (at compile-time) between the page_zone() and sparsemem operations. However, on 32-bit architectures, the number of bits is quite limited, and may require growing the size of the page->flags type in certain conditions. Several things might force this to occur: a decrease in the SECTION_SIZE (if you want to hotplug smaller areas of memory), an increase in the physical address space, or an increase in the number of used page->flags. One thing to note is that, once sparsemem is present, the NUMA node information no longer needs to be stored in the page->flags. It might provide speed increases on certain platforms and will be stored there if there is room. But, if out of room, an alternate (theoretically slower) mechanism is used. This patch introduces CONFIG_FLATMEM. It is used in almost all cases where there used to be an #ifndef DISCONTIG, because SPARSEMEM and DISCONTIGMEM often have to compile out the same areas of code. Signed-off-by: Andy Whitcroft <apw@shadowen.org> Signed-off-by: Dave Hansen <haveblue@us.ibm.com> Signed-off-by: Martin Bligh <mbligh@aracnet.com> Signed-off-by: Adrian Bunk <bunk@stusta.de> Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com> Signed-off-by: Bob Picco <bob.picco@hp.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 00:07:54 -07:00
This option provides efficient support for systems with
holes is their physical address space and allows memory
hot-plug and hot-remove.
[PATCH] sparsemem memory model Sparsemem abstracts the use of discontiguous mem_maps[]. This kind of mem_map[] is needed by discontiguous memory machines (like in the old CONFIG_DISCONTIGMEM case) as well as memory hotplug systems. Sparsemem replaces DISCONTIGMEM when enabled, and it is hoped that it can eventually become a complete replacement. A significant advantage over DISCONTIGMEM is that it's completely separated from CONFIG_NUMA. When producing this patch, it became apparent in that NUMA and DISCONTIG are often confused. Another advantage is that sparse doesn't require each NUMA node's ranges to be contiguous. It can handle overlapping ranges between nodes with no problems, where DISCONTIGMEM currently throws away that memory. Sparsemem uses an array to provide different pfn_to_page() translations for each SECTION_SIZE area of physical memory. This is what allows the mem_map[] to be chopped up. In order to do quick pfn_to_page() operations, the section number of the page is encoded in page->flags. Part of the sparsemem infrastructure enables sharing of these bits more dynamically (at compile-time) between the page_zone() and sparsemem operations. However, on 32-bit architectures, the number of bits is quite limited, and may require growing the size of the page->flags type in certain conditions. Several things might force this to occur: a decrease in the SECTION_SIZE (if you want to hotplug smaller areas of memory), an increase in the physical address space, or an increase in the number of used page->flags. One thing to note is that, once sparsemem is present, the NUMA node information no longer needs to be stored in the page->flags. It might provide speed increases on certain platforms and will be stored there if there is room. But, if out of room, an alternate (theoretically slower) mechanism is used. This patch introduces CONFIG_FLATMEM. It is used in almost all cases where there used to be an #ifndef DISCONTIG, because SPARSEMEM and DISCONTIGMEM often have to compile out the same areas of code. Signed-off-by: Andy Whitcroft <apw@shadowen.org> Signed-off-by: Dave Hansen <haveblue@us.ibm.com> Signed-off-by: Martin Bligh <mbligh@aracnet.com> Signed-off-by: Adrian Bunk <bunk@stusta.de> Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com> Signed-off-by: Bob Picco <bob.picco@hp.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 00:07:54 -07:00
If unsure, choose "Flat Memory" over this option.
[PATCH] sparsemem memory model Sparsemem abstracts the use of discontiguous mem_maps[]. This kind of mem_map[] is needed by discontiguous memory machines (like in the old CONFIG_DISCONTIGMEM case) as well as memory hotplug systems. Sparsemem replaces DISCONTIGMEM when enabled, and it is hoped that it can eventually become a complete replacement. A significant advantage over DISCONTIGMEM is that it's completely separated from CONFIG_NUMA. When producing this patch, it became apparent in that NUMA and DISCONTIG are often confused. Another advantage is that sparse doesn't require each NUMA node's ranges to be contiguous. It can handle overlapping ranges between nodes with no problems, where DISCONTIGMEM currently throws away that memory. Sparsemem uses an array to provide different pfn_to_page() translations for each SECTION_SIZE area of physical memory. This is what allows the mem_map[] to be chopped up. In order to do quick pfn_to_page() operations, the section number of the page is encoded in page->flags. Part of the sparsemem infrastructure enables sharing of these bits more dynamically (at compile-time) between the page_zone() and sparsemem operations. However, on 32-bit architectures, the number of bits is quite limited, and may require growing the size of the page->flags type in certain conditions. Several things might force this to occur: a decrease in the SECTION_SIZE (if you want to hotplug smaller areas of memory), an increase in the physical address space, or an increase in the number of used page->flags. One thing to note is that, once sparsemem is present, the NUMA node information no longer needs to be stored in the page->flags. It might provide speed increases on certain platforms and will be stored there if there is room. But, if out of room, an alternate (theoretically slower) mechanism is used. This patch introduces CONFIG_FLATMEM. It is used in almost all cases where there used to be an #ifndef DISCONTIG, because SPARSEMEM and DISCONTIGMEM often have to compile out the same areas of code. Signed-off-by: Andy Whitcroft <apw@shadowen.org> Signed-off-by: Dave Hansen <haveblue@us.ibm.com> Signed-off-by: Martin Bligh <mbligh@aracnet.com> Signed-off-by: Adrian Bunk <bunk@stusta.de> Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com> Signed-off-by: Bob Picco <bob.picco@hp.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 00:07:54 -07:00
endchoice
[PATCH] sparsemem memory model Sparsemem abstracts the use of discontiguous mem_maps[]. This kind of mem_map[] is needed by discontiguous memory machines (like in the old CONFIG_DISCONTIGMEM case) as well as memory hotplug systems. Sparsemem replaces DISCONTIGMEM when enabled, and it is hoped that it can eventually become a complete replacement. A significant advantage over DISCONTIGMEM is that it's completely separated from CONFIG_NUMA. When producing this patch, it became apparent in that NUMA and DISCONTIG are often confused. Another advantage is that sparse doesn't require each NUMA node's ranges to be contiguous. It can handle overlapping ranges between nodes with no problems, where DISCONTIGMEM currently throws away that memory. Sparsemem uses an array to provide different pfn_to_page() translations for each SECTION_SIZE area of physical memory. This is what allows the mem_map[] to be chopped up. In order to do quick pfn_to_page() operations, the section number of the page is encoded in page->flags. Part of the sparsemem infrastructure enables sharing of these bits more dynamically (at compile-time) between the page_zone() and sparsemem operations. However, on 32-bit architectures, the number of bits is quite limited, and may require growing the size of the page->flags type in certain conditions. Several things might force this to occur: a decrease in the SECTION_SIZE (if you want to hotplug smaller areas of memory), an increase in the physical address space, or an increase in the number of used page->flags. One thing to note is that, once sparsemem is present, the NUMA node information no longer needs to be stored in the page->flags. It might provide speed increases on certain platforms and will be stored there if there is room. But, if out of room, an alternate (theoretically slower) mechanism is used. This patch introduces CONFIG_FLATMEM. It is used in almost all cases where there used to be an #ifndef DISCONTIG, because SPARSEMEM and DISCONTIGMEM often have to compile out the same areas of code. Signed-off-by: Andy Whitcroft <apw@shadowen.org> Signed-off-by: Dave Hansen <haveblue@us.ibm.com> Signed-off-by: Martin Bligh <mbligh@aracnet.com> Signed-off-by: Adrian Bunk <bunk@stusta.de> Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com> Signed-off-by: Bob Picco <bob.picco@hp.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 00:07:54 -07:00
config SPARSEMEM
def_bool y
depends on (!SELECT_MEMORY_MODEL && ARCH_SPARSEMEM_ENABLE) || SPARSEMEM_MANUAL
[PATCH] sparsemem memory model Sparsemem abstracts the use of discontiguous mem_maps[]. This kind of mem_map[] is needed by discontiguous memory machines (like in the old CONFIG_DISCONTIGMEM case) as well as memory hotplug systems. Sparsemem replaces DISCONTIGMEM when enabled, and it is hoped that it can eventually become a complete replacement. A significant advantage over DISCONTIGMEM is that it's completely separated from CONFIG_NUMA. When producing this patch, it became apparent in that NUMA and DISCONTIG are often confused. Another advantage is that sparse doesn't require each NUMA node's ranges to be contiguous. It can handle overlapping ranges between nodes with no problems, where DISCONTIGMEM currently throws away that memory. Sparsemem uses an array to provide different pfn_to_page() translations for each SECTION_SIZE area of physical memory. This is what allows the mem_map[] to be chopped up. In order to do quick pfn_to_page() operations, the section number of the page is encoded in page->flags. Part of the sparsemem infrastructure enables sharing of these bits more dynamically (at compile-time) between the page_zone() and sparsemem operations. However, on 32-bit architectures, the number of bits is quite limited, and may require growing the size of the page->flags type in certain conditions. Several things might force this to occur: a decrease in the SECTION_SIZE (if you want to hotplug smaller areas of memory), an increase in the physical address space, or an increase in the number of used page->flags. One thing to note is that, once sparsemem is present, the NUMA node information no longer needs to be stored in the page->flags. It might provide speed increases on certain platforms and will be stored there if there is room. But, if out of room, an alternate (theoretically slower) mechanism is used. This patch introduces CONFIG_FLATMEM. It is used in almost all cases where there used to be an #ifndef DISCONTIG, because SPARSEMEM and DISCONTIGMEM often have to compile out the same areas of code. Signed-off-by: Andy Whitcroft <apw@shadowen.org> Signed-off-by: Dave Hansen <haveblue@us.ibm.com> Signed-off-by: Martin Bligh <mbligh@aracnet.com> Signed-off-by: Adrian Bunk <bunk@stusta.de> Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com> Signed-off-by: Bob Picco <bob.picco@hp.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 00:07:54 -07:00
config FLATMEM
def_bool y
depends on !SPARSEMEM || FLATMEM_MANUAL
[PATCH] sparsemem memory model Sparsemem abstracts the use of discontiguous mem_maps[]. This kind of mem_map[] is needed by discontiguous memory machines (like in the old CONFIG_DISCONTIGMEM case) as well as memory hotplug systems. Sparsemem replaces DISCONTIGMEM when enabled, and it is hoped that it can eventually become a complete replacement. A significant advantage over DISCONTIGMEM is that it's completely separated from CONFIG_NUMA. When producing this patch, it became apparent in that NUMA and DISCONTIG are often confused. Another advantage is that sparse doesn't require each NUMA node's ranges to be contiguous. It can handle overlapping ranges between nodes with no problems, where DISCONTIGMEM currently throws away that memory. Sparsemem uses an array to provide different pfn_to_page() translations for each SECTION_SIZE area of physical memory. This is what allows the mem_map[] to be chopped up. In order to do quick pfn_to_page() operations, the section number of the page is encoded in page->flags. Part of the sparsemem infrastructure enables sharing of these bits more dynamically (at compile-time) between the page_zone() and sparsemem operations. However, on 32-bit architectures, the number of bits is quite limited, and may require growing the size of the page->flags type in certain conditions. Several things might force this to occur: a decrease in the SECTION_SIZE (if you want to hotplug smaller areas of memory), an increase in the physical address space, or an increase in the number of used page->flags. One thing to note is that, once sparsemem is present, the NUMA node information no longer needs to be stored in the page->flags. It might provide speed increases on certain platforms and will be stored there if there is room. But, if out of room, an alternate (theoretically slower) mechanism is used. This patch introduces CONFIG_FLATMEM. It is used in almost all cases where there used to be an #ifndef DISCONTIG, because SPARSEMEM and DISCONTIGMEM often have to compile out the same areas of code. Signed-off-by: Andy Whitcroft <apw@shadowen.org> Signed-off-by: Dave Hansen <haveblue@us.ibm.com> Signed-off-by: Martin Bligh <mbligh@aracnet.com> Signed-off-by: Adrian Bunk <bunk@stusta.de> Signed-off-by: Yasunori Goto <y-goto@jp.fujitsu.com> Signed-off-by: Bob Picco <bob.picco@hp.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-06-23 00:07:54 -07:00
#
# SPARSEMEM_EXTREME (which is the default) does some bootmem
# allocations when sparse_init() is called. If this cannot
# be done on your architecture, select this option. However,
# statically allocating the mem_section[] array can potentially
# consume vast quantities of .bss, so be careful.
#
# This option will also potentially produce smaller runtime code
# with gcc 3.4 and later.
#
config SPARSEMEM_STATIC
bool
#
# Architecture platforms which require a two level mem_section in SPARSEMEM
# must select this option. This is usually for architecture platforms with
# an extremely sparse physical address space.
#
config SPARSEMEM_EXTREME
def_bool y
depends on SPARSEMEM && !SPARSEMEM_STATIC
[PATCH] mm: split page table lock Christoph Lameter demonstrated very poor scalability on the SGI 512-way, with a many-threaded application which concurrently initializes different parts of a large anonymous area. This patch corrects that, by using a separate spinlock per page table page, to guard the page table entries in that page, instead of using the mm's single page_table_lock. (But even then, page_table_lock is still used to guard page table allocation, and anon_vma allocation.) In this implementation, the spinlock is tucked inside the struct page of the page table page: with a BUILD_BUG_ON in case it overflows - which it would in the case of 32-bit PA-RISC with spinlock debugging enabled. Splitting the lock is not quite for free: another cacheline access. Ideally, I suppose we would use split ptlock only for multi-threaded processes on multi-cpu machines; but deciding that dynamically would have its own costs. So for now enable it by config, at some number of cpus - since the Kconfig language doesn't support inequalities, let preprocessor compare that with NR_CPUS. But I don't think it's worth being user-configurable: for good testing of both split and unsplit configs, split now at 4 cpus, and perhaps change that to 8 later. There is a benefit even for singly threaded processes: kswapd can be attacking one part of the mm while another part is busy faulting. Signed-off-by: Hugh Dickins <hugh@veritas.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-10-29 18:16:40 -07:00
config SPARSEMEM_VMEMMAP_ENABLE
bool
config SPARSEMEM_VMEMMAP
bool "Sparse Memory virtual memmap"
depends on SPARSEMEM && SPARSEMEM_VMEMMAP_ENABLE
default y
help
SPARSEMEM_VMEMMAP uses a virtually mapped memmap to optimise
pfn_to_page and page_to_pfn operations. This is the most
efficient option when sufficient kernel resources are available.
#
# Select this config option from the architecture Kconfig, if it is preferred
# to enable the feature of HugeTLB/dev_dax vmemmap optimization.
#
config ARCH_WANT_OPTIMIZE_DAX_VMEMMAP
bool
config ARCH_WANT_OPTIMIZE_HUGETLB_VMEMMAP
bool
config HAVE_MEMBLOCK_PHYS_MAP
bool
config HAVE_GUP_FAST
depends on MMU
bool
mm: introduce a general RCU get_user_pages_fast() This series implements general forms of get_user_pages_fast and __get_user_pages_fast in core code and activates them for arm and arm64. These are required for Transparent HugePages to function correctly, as a futex on a THP tail will otherwise result in an infinite loop (due to the core implementation of __get_user_pages_fast always returning 0). Unfortunately, a futex on THP tail can be quite common for certain workloads; thus THP is unreliable without a __get_user_pages_fast implementation. This series may also be beneficial for direct-IO heavy workloads and certain KVM workloads. This patch (of 6): get_user_pages_fast() attempts to pin user pages by walking the page tables directly and avoids taking locks. Thus the walker needs to be protected from page table pages being freed from under it, and needs to block any THP splits. One way to achieve this is to have the walker disable interrupts, and rely on IPIs from the TLB flushing code blocking before the page table pages are freed. On some platforms we have hardware broadcast of TLB invalidations, thus the TLB flushing code doesn't necessarily need to broadcast IPIs; and spuriously broadcasting IPIs can hurt system performance if done too often. This problem has been solved on PowerPC and Sparc by batching up page table pages belonging to more than one mm_user, then scheduling an rcu_sched callback to free the pages. This RCU page table free logic has been promoted to core code and is activated when one enables HAVE_RCU_TABLE_FREE. Unfortunately, these architectures implement their own get_user_pages_fast routines. The RCU page table free logic coupled with an IPI broadcast on THP split (which is a rare event), allows one to protect a page table walker by merely disabling the interrupts during the walk. This patch provides a general RCU implementation of get_user_pages_fast that can be used by architectures that perform hardware broadcast of TLB invalidations. It is based heavily on the PowerPC implementation by Nick Piggin. [akpm@linux-foundation.org: various comment fixes] Signed-off-by: Steve Capper <steve.capper@linaro.org> Tested-by: Dann Frazier <dann.frazier@canonical.com> Reviewed-by: Catalin Marinas <catalin.marinas@arm.com> Acked-by: Hugh Dickins <hughd@google.com> Cc: Russell King <rmk@arm.linux.org.uk> Cc: Mark Rutland <mark.rutland@arm.com> Cc: Mel Gorman <mel@csn.ul.ie> Cc: Will Deacon <will.deacon@arm.com> Cc: Christoffer Dall <christoffer.dall@linaro.org> Cc: Andrea Arcangeli <aarcange@redhat.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2014-10-09 15:29:14 -07:00
# Don't discard allocated memory used to track "memory" and "reserved" memblocks
# after early boot, so it can still be used to test for validity of memory.
# Also, memblocks are updated with memory hot(un)plug.
mm: memblock: make keeping memblock memory opt-in rather than opt-out Most architectures do not need the memblock memory after the page allocator is initialized, but only few enable ARCH_DISCARD_MEMBLOCK in the arch Kconfig. Replacing ARCH_DISCARD_MEMBLOCK with ARCH_KEEP_MEMBLOCK and inverting the logic makes it clear which architectures actually use memblock after system initialization and skips the necessity to add ARCH_DISCARD_MEMBLOCK to the architectures that are still missing that option. Link: http://lkml.kernel.org/r/1556102150-32517-1-git-send-email-rppt@linux.ibm.com Signed-off-by: Mike Rapoport <rppt@linux.ibm.com> Acked-by: Michael Ellerman <mpe@ellerman.id.au> (powerpc) Cc: Russell King <linux@armlinux.org.uk> Cc: Catalin Marinas <catalin.marinas@arm.com> Cc: Will Deacon <will.deacon@arm.com> Cc: Richard Kuo <rkuo@codeaurora.org> Cc: Tony Luck <tony.luck@intel.com> Cc: Fenghua Yu <fenghua.yu@intel.com> Cc: Geert Uytterhoeven <geert@linux-m68k.org> Cc: Ralf Baechle <ralf@linux-mips.org> Cc: Paul Burton <paul.burton@mips.com> Cc: James Hogan <jhogan@kernel.org> Cc: Ley Foon Tan <lftan@altera.com> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: Paul Mackerras <paulus@samba.org> Cc: Martin Schwidefsky <schwidefsky@de.ibm.com> Cc: Heiko Carstens <heiko.carstens@de.ibm.com> Cc: Yoshinori Sato <ysato@users.sourceforge.jp> Cc: Rich Felker <dalias@libc.org> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Ingo Molnar <mingo@redhat.com> Cc: Borislav Petkov <bp@alien8.de> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Eric Biederman <ebiederm@xmission.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2019-05-13 17:22:59 -07:00
config ARCH_KEEP_MEMBLOCK
bool
# Keep arch NUMA mapping infrastructure post-init.
config NUMA_KEEP_MEMINFO
bool
config MEMORY_ISOLATION
bool
kernel/resource: disallow access to exclusive system RAM regions virtio-mem dynamically exposes memory inside a device memory region as system RAM to Linux, coordinating with the hypervisor which parts are actually "plugged" and consequently usable/accessible. On the one hand, the virtio-mem driver adds/removes whole memory blocks, creating/removing busy IORESOURCE_SYSTEM_RAM resources, on the other hand, it logically (un)plugs memory inside added memory blocks, dynamically either exposing them to the buddy or hiding them from the buddy and marking them PG_offline. In contrast to physical devices, like a DIMM, the virtio-mem driver is required to actually make use of any of the device-provided memory, because it performs the handshake with the hypervisor. virtio-mem memory cannot simply be access via /dev/mem without a driver. There is no safe way to: a) Access plugged memory blocks via /dev/mem, as they might contain unplugged holes or might get silently unplugged by the virtio-mem driver and consequently turned inaccessible. b) Access unplugged memory blocks via /dev/mem because the virtio-mem driver is required to make them actually accessible first. The virtio-spec states that unplugged memory blocks MUST NOT be written, and only selected unplugged memory blocks MAY be read. We want to make sure, this is the case in sane environments -- where the virtio-mem driver was loaded. We want to make sure that in a sane environment, nobody "accidentially" accesses unplugged memory inside the device managed region. For example, a user might spot a memory region in /proc/iomem and try accessing it via /dev/mem via gdb or dumping it via something else. By the time the mmap() happens, the memory might already have been removed by the virtio-mem driver silently: the mmap() would succeeed and user space might accidentially access unplugged memory. So once the driver was loaded and detected the device along the device-managed region, we just want to disallow any access via /dev/mem to it. In an ideal world, we would mark the whole region as busy ("owned by a driver") and exclude it; however, that would be wrong, as we don't really have actual system RAM at these ranges added to Linux ("busy system RAM"). Instead, we want to mark such ranges as "not actual busy system RAM but still soft-reserved and prepared by a driver for future use." Let's teach iomem_is_exclusive() to reject access to any range with "IORESOURCE_SYSTEM_RAM | IORESOURCE_EXCLUSIVE", even if not busy and even if "iomem=relaxed" is set. Introduce EXCLUSIVE_SYSTEM_RAM to make it easier for applicable drivers to depend on this setting in their Kconfig. For now, there are no applicable ranges and we'll modify virtio-mem next to properly set IORESOURCE_EXCLUSIVE on the parent resource container it creates to contain all actual busy system RAM added via add_memory_driver_managed(). Link: https://lkml.kernel.org/r/20210920142856.17758-3-david@redhat.com Signed-off-by: David Hildenbrand <david@redhat.com> Reviewed-by: Dan Williams <dan.j.williams@intel.com> Cc: Andy Shevchenko <andy.shevchenko@gmail.com> Cc: Arnd Bergmann <arnd@arndb.de> Cc: Greg Kroah-Hartman <gregkh@linuxfoundation.org> Cc: Hanjun Guo <guohanjun@huawei.com> Cc: Jason Wang <jasowang@redhat.com> Cc: "Michael S. Tsirkin" <mst@redhat.com> Cc: "Rafael J. Wysocki" <rafael.j.wysocki@intel.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2021-11-08 19:35:50 -07:00
# IORESOURCE_SYSTEM_RAM regions in the kernel resource tree that are marked
# IORESOURCE_EXCLUSIVE cannot be mapped to user space, for example, via
# /dev/mem.
config EXCLUSIVE_SYSTEM_RAM
def_bool y
depends on !DEVMEM || STRICT_DEVMEM
memory-hotplug: implement register_page_bootmem_info_section of sparse-vmemmap For removing memmap region of sparse-vmemmap which is allocated bootmem, memmap region of sparse-vmemmap needs to be registered by get_page_bootmem(). So the patch searches pages of virtual mapping and registers the pages by get_page_bootmem(). NOTE: register_page_bootmem_memmap() is not implemented for ia64, ppc, s390, and sparc. So introduce CONFIG_HAVE_BOOTMEM_INFO_NODE and revert register_page_bootmem_info_node() when platform doesn't support it. It's implemented by adding a new Kconfig option named CONFIG_HAVE_BOOTMEM_INFO_NODE, which will be automatically selected by memory-hotplug feature fully supported archs(currently only on x86_64). Since we have 2 config options called MEMORY_HOTPLUG and MEMORY_HOTREMOVE used for memory hot-add and hot-remove separately, and codes in function register_page_bootmem_info_node() are only used for collecting infomation for hot-remove, so reside it under MEMORY_HOTREMOVE. Besides page_isolation.c selected by MEMORY_ISOLATION under MEMORY_HOTPLUG is also such case, move it too. [mhocko@suse.cz: put register_page_bootmem_memmap inside CONFIG_MEMORY_HOTPLUG_SPARSE] [linfeng@cn.fujitsu.com: introduce CONFIG_HAVE_BOOTMEM_INFO_NODE and revert register_page_bootmem_info_node()] [mhocko@suse.cz: remove the arch specific functions without any implementation] [linfeng@cn.fujitsu.com: mm/Kconfig: move auto selects from MEMORY_HOTPLUG to MEMORY_HOTREMOVE as needed] [rientjes@google.com: fix defined but not used warning] Signed-off-by: Wen Congyang <wency@cn.fujitsu.com> Signed-off-by: Yasuaki Ishimatsu <isimatu.yasuaki@jp.fujitsu.com> Signed-off-by: Tang Chen <tangchen@cn.fujitsu.com> Reviewed-by: Wu Jianguo <wujianguo@huawei.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Jiang Liu <jiang.liu@huawei.com> Cc: Jianguo Wu <wujianguo@huawei.com> Cc: Kamezawa Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: Lai Jiangshan <laijs@cn.fujitsu.com> Cc: Ingo Molnar <mingo@elte.hu> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: "H. Peter Anvin" <hpa@zytor.com> Signed-off-by: Michal Hocko <mhocko@suse.cz> Signed-off-by: Lin Feng <linfeng@cn.fujitsu.com> Signed-off-by: David Rientjes <rientjes@google.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2013-02-22 17:33:00 -07:00
#
# Only be set on architectures that have completely implemented memory hotplug
# feature. If you are not sure, don't touch it.
#
config HAVE_BOOTMEM_INFO_NODE
def_bool n
mm: generalize ARCH_ENABLE_MEMORY_[HOTPLUG|HOTREMOVE] ARCH_ENABLE_MEMORY_[HOTPLUG|HOTREMOVE] configs have duplicate definitions on platforms that subscribe them. Instead, just make them generic options which can be selected on applicable platforms. Link: https://lkml.kernel.org/r/1617259448-22529-4-git-send-email-anshuman.khandual@arm.com Signed-off-by: Anshuman Khandual <anshuman.khandual@arm.com> Acked-by: Catalin Marinas <catalin.marinas@arm.com> [arm64] Acked-by: Heiko Carstens <hca@linux.ibm.com> [s390] Cc: Will Deacon <will@kernel.org> Cc: Michael Ellerman <mpe@ellerman.id.au> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: Paul Mackerras <paulus@samba.org> Cc: Vasily Gorbik <gor@linux.ibm.com> Cc: Christian Borntraeger <borntraeger@de.ibm.com> Cc: Yoshinori Sato <ysato@users.sourceforge.jp> Cc: Rich Felker <dalias@libc.org> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Ingo Molnar <mingo@redhat.com> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Albert Ou <aou@eecs.berkeley.edu> Cc: Alexander Viro <viro@zeniv.linux.org.uk> Cc: Arnd Bergmann <arnd@arndb.de> Cc: Borislav Petkov <bp@alien8.de> Cc: Helge Deller <deller@gmx.de> Cc: "James E.J. Bottomley" <James.Bottomley@HansenPartnership.com> Cc: Palmer Dabbelt <palmerdabbelt@google.com> Cc: Paul Walmsley <paul.walmsley@sifive.com> Cc: Russell King <linux@armlinux.org.uk> Cc: Thomas Bogendoerfer <tsbogend@alpha.franken.de> Cc: Vineet Gupta <vgupta@synopsys.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2021-05-04 18:38:17 -07:00
config ARCH_ENABLE_MEMORY_HOTPLUG
bool
config ARCH_ENABLE_MEMORY_HOTREMOVE
bool
# eventually, we can have this option just 'select SPARSEMEM'
menuconfig MEMORY_HOTPLUG
bool "Memory hotplug"
mm/memory_hotplug: mark pageblocks MIGRATE_ISOLATE while onlining memory Currently, it can happen that pages are allocated (and freed) via the buddy before we finished basic memory onlining. For example, pages are exposed to the buddy and can be allocated before we actually mark the sections online. Allocated pages could suddenly fail pfn_to_online_page() checks. We had similar issues with pcp handling, when pages are allocated+freed before we reach zone_pcp_update() in online_pages() [1]. Instead, mark all pageblocks MIGRATE_ISOLATE, such that allocations are impossible. Once done with the heavy lifting, use undo_isolate_page_range() to move the pages to the MIGRATE_MOVABLE freelist, marking them ready for allocation. Similar to offline_pages(), we have to manually adjust zone->nr_isolate_pageblock. [1] https://lkml.kernel.org/r/1597150703-19003-1-git-send-email-charante@codeaurora.org Signed-off-by: David Hildenbrand <david@redhat.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Reviewed-by: Oscar Salvador <osalvador@suse.de> Acked-by: Michal Hocko <mhocko@suse.com> Cc: Wei Yang <richard.weiyang@linux.alibaba.com> Cc: Baoquan He <bhe@redhat.com> Cc: Pankaj Gupta <pankaj.gupta.linux@gmail.com> Cc: Charan Teja Reddy <charante@codeaurora.org> Cc: Dan Williams <dan.j.williams@intel.com> Cc: Fenghua Yu <fenghua.yu@intel.com> Cc: Logan Gunthorpe <logang@deltatee.com> Cc: "Matthew Wilcox (Oracle)" <willy@infradead.org> Cc: Mel Gorman <mgorman@suse.de> Cc: Mel Gorman <mgorman@techsingularity.net> Cc: Michel Lespinasse <walken@google.com> Cc: Mike Rapoport <rppt@kernel.org> Cc: Tony Luck <tony.luck@intel.com> Link: https://lkml.kernel.org/r/20200819175957.28465-11-david@redhat.com Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2020-10-15 20:08:23 -07:00
select MEMORY_ISOLATION
mm/memory_hotplug: remove CONFIG_X86_64_ACPI_NUMA dependency from CONFIG_MEMORY_HOTPLUG Patch series "mm/memory_hotplug: Kconfig and 32 bit cleanups". Some cleanups around CONFIG_MEMORY_HOTPLUG, including removing 32 bit leftovers of memory hotplug support. This patch (of 6): SPARSEMEM is the only possible memory model for x86-64, FLATMEM is not possible: config ARCH_FLATMEM_ENABLE def_bool y depends on X86_32 && !NUMA And X86_64_ACPI_NUMA (obviously) only supports x86-64: config X86_64_ACPI_NUMA def_bool y depends on X86_64 && NUMA && ACPI && PCI Let's just remove the CONFIG_X86_64_ACPI_NUMA dependency, as it does no longer make sense. Link: https://lkml.kernel.org/r/20210929143600.49379-2-david@redhat.com Signed-off-by: David Hildenbrand <david@redhat.com> Reviewed-by: Oscar Salvador <osalvador@suse.de> Cc: Jonathan Corbet <corbet@lwn.net> Cc: Alex Shi <alexs@kernel.org> Cc: Michael Ellerman <mpe@ellerman.id.au> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: Paul Mackerras <paulus@samba.org> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Ingo Molnar <mingo@redhat.com> Cc: Borislav Petkov <bp@alien8.de> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Dave Hansen <dave.hansen@linux.intel.com> Cc: Andy Lutomirski <luto@kernel.org> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Greg Kroah-Hartman <gregkh@linuxfoundation.org> Cc: "Rafael J. Wysocki" <rafael@kernel.org> Cc: "Michael S. Tsirkin" <mst@redhat.com> Cc: Jason Wang <jasowang@redhat.com> Cc: Shuah Khan <shuah@kernel.org> Cc: Michal Hocko <mhocko@suse.com> Cc: Mike Rapoport <rppt@kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2021-11-05 13:44:20 -07:00
depends on SPARSEMEM
depends on ARCH_ENABLE_MEMORY_HOTPLUG
mm/memory_hotplug: restrict CONFIG_MEMORY_HOTPLUG to 64 bit 32 bit support is broken in various ways: for example, we can online memory that should actually go to ZONE_HIGHMEM to ZONE_MOVABLE or in some cases even to one of the other kernel zones. We marked it BROKEN in commit b59d02ed0869 ("mm/memory_hotplug: disable the functionality for 32b") almost one year ago. According to that commit it might be broken at least since 2017. Further, there is hardly a sane use case nowadays. Let's just depend completely on 64bit, dropping the "BROKEN" dependency to make clear that we are not going to support it again. Next, we'll remove some HIGHMEM leftovers from memory hotplug code to clean up. Link: https://lkml.kernel.org/r/20210929143600.49379-4-david@redhat.com Signed-off-by: David Hildenbrand <david@redhat.com> Reviewed-by: Oscar Salvador <osalvador@suse.de> Cc: Alex Shi <alexs@kernel.org> Cc: Andy Lutomirski <luto@kernel.org> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: Borislav Petkov <bp@alien8.de> Cc: Dave Hansen <dave.hansen@linux.intel.com> Cc: Greg Kroah-Hartman <gregkh@linuxfoundation.org> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Ingo Molnar <mingo@redhat.com> Cc: Jason Wang <jasowang@redhat.com> Cc: Jonathan Corbet <corbet@lwn.net> Cc: Michael Ellerman <mpe@ellerman.id.au> Cc: "Michael S. Tsirkin" <mst@redhat.com> Cc: Michal Hocko <mhocko@suse.com> Cc: Mike Rapoport <rppt@kernel.org> Cc: Paul Mackerras <paulus@samba.org> Cc: Peter Zijlstra <peterz@infradead.org> Cc: "Rafael J. Wysocki" <rafael@kernel.org> Cc: Shuah Khan <shuah@kernel.org> Cc: Thomas Gleixner <tglx@linutronix.de> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2021-11-05 13:44:28 -07:00
depends on 64BIT
select NUMA_KEEP_MEMINFO if NUMA
if MEMORY_HOTPLUG
memory_hotplug: introduce CONFIG_MEMORY_HOTPLUG_DEFAULT_ONLINE This patchset continues the work I started with commit 31bc3858ea3e ("memory-hotplug: add automatic onlining policy for the newly added memory"). Initially I was going to stop there and bring the policy setting logic to userspace. I met two issues on this way: 1) It is possible to have memory hotplugged at boot (e.g. with QEMU). These blocks stay offlined if we turn the onlining policy on by userspace. 2) My attempt to bring this policy setting to systemd failed, systemd maintainers suggest to change the default in kernel or ... to use tmpfiles.d to alter the policy (which looks like a hack to me): https://github.com/systemd/systemd/pull/2938 Here I suggest to add a config option to set the default value for the policy and a kernel command line parameter to make the override. This patch (of 2): Introduce config option to set the default value for memory hotplug onlining policy (/sys/devices/system/memory/auto_online_blocks). The reason one would want to turn this option on are to have early onlining for hotpluggable memory available at boot and to not require any userspace actions to make memory hotplug work. [akpm@linux-foundation.org: tweak Kconfig text] Signed-off-by: Vitaly Kuznetsov <vkuznets@redhat.com> Cc: Jonathan Corbet <corbet@lwn.net> Cc: Dan Williams <dan.j.williams@intel.com> Cc: "Kirill A. Shutemov" <kirill.shutemov@linux.intel.com> Cc: Mel Gorman <mgorman@techsingularity.net> Cc: David Vrabel <david.vrabel@citrix.com> Cc: David Rientjes <rientjes@google.com> Cc: Igor Mammedov <imammedo@redhat.com> Cc: Lennart Poettering <lennart@poettering.net> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-05-19 17:13:03 -07:00
config MEMORY_HOTPLUG_DEFAULT_ONLINE
bool "Online the newly added memory blocks by default"
depends on MEMORY_HOTPLUG
help
memory_hotplug: introduce CONFIG_MEMORY_HOTPLUG_DEFAULT_ONLINE This patchset continues the work I started with commit 31bc3858ea3e ("memory-hotplug: add automatic onlining policy for the newly added memory"). Initially I was going to stop there and bring the policy setting logic to userspace. I met two issues on this way: 1) It is possible to have memory hotplugged at boot (e.g. with QEMU). These blocks stay offlined if we turn the onlining policy on by userspace. 2) My attempt to bring this policy setting to systemd failed, systemd maintainers suggest to change the default in kernel or ... to use tmpfiles.d to alter the policy (which looks like a hack to me): https://github.com/systemd/systemd/pull/2938 Here I suggest to add a config option to set the default value for the policy and a kernel command line parameter to make the override. This patch (of 2): Introduce config option to set the default value for memory hotplug onlining policy (/sys/devices/system/memory/auto_online_blocks). The reason one would want to turn this option on are to have early onlining for hotpluggable memory available at boot and to not require any userspace actions to make memory hotplug work. [akpm@linux-foundation.org: tweak Kconfig text] Signed-off-by: Vitaly Kuznetsov <vkuznets@redhat.com> Cc: Jonathan Corbet <corbet@lwn.net> Cc: Dan Williams <dan.j.williams@intel.com> Cc: "Kirill A. Shutemov" <kirill.shutemov@linux.intel.com> Cc: Mel Gorman <mgorman@techsingularity.net> Cc: David Vrabel <david.vrabel@citrix.com> Cc: David Rientjes <rientjes@google.com> Cc: Igor Mammedov <imammedo@redhat.com> Cc: Lennart Poettering <lennart@poettering.net> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-05-19 17:13:03 -07:00
This option sets the default policy setting for memory hotplug
onlining policy (/sys/devices/system/memory/auto_online_blocks) which
determines what happens to newly added memory regions. Policy setting
can always be changed at runtime.
See Documentation/admin-guide/mm/memory-hotplug.rst for more information.
memory_hotplug: introduce CONFIG_MEMORY_HOTPLUG_DEFAULT_ONLINE This patchset continues the work I started with commit 31bc3858ea3e ("memory-hotplug: add automatic onlining policy for the newly added memory"). Initially I was going to stop there and bring the policy setting logic to userspace. I met two issues on this way: 1) It is possible to have memory hotplugged at boot (e.g. with QEMU). These blocks stay offlined if we turn the onlining policy on by userspace. 2) My attempt to bring this policy setting to systemd failed, systemd maintainers suggest to change the default in kernel or ... to use tmpfiles.d to alter the policy (which looks like a hack to me): https://github.com/systemd/systemd/pull/2938 Here I suggest to add a config option to set the default value for the policy and a kernel command line parameter to make the override. This patch (of 2): Introduce config option to set the default value for memory hotplug onlining policy (/sys/devices/system/memory/auto_online_blocks). The reason one would want to turn this option on are to have early onlining for hotpluggable memory available at boot and to not require any userspace actions to make memory hotplug work. [akpm@linux-foundation.org: tweak Kconfig text] Signed-off-by: Vitaly Kuznetsov <vkuznets@redhat.com> Cc: Jonathan Corbet <corbet@lwn.net> Cc: Dan Williams <dan.j.williams@intel.com> Cc: "Kirill A. Shutemov" <kirill.shutemov@linux.intel.com> Cc: Mel Gorman <mgorman@techsingularity.net> Cc: David Vrabel <david.vrabel@citrix.com> Cc: David Rientjes <rientjes@google.com> Cc: Igor Mammedov <imammedo@redhat.com> Cc: Lennart Poettering <lennart@poettering.net> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-05-19 17:13:03 -07:00
Say Y here if you want all hot-plugged memory blocks to appear in
'online' state by default.
Say N here if you want the default policy to keep all hot-plugged
memory blocks in 'offline' state.
config MEMORY_HOTREMOVE
bool "Allow for memory hot remove"
select HAVE_BOOTMEM_INFO_NODE if (X86_64 || PPC64)
depends on MEMORY_HOTPLUG && ARCH_ENABLE_MEMORY_HOTREMOVE
depends on MIGRATION
mm,memory_hotplug: allocate memmap from the added memory range Physical memory hotadd has to allocate a memmap (struct page array) for the newly added memory section. Currently, alloc_pages_node() is used for those allocations. This has some disadvantages: a) an existing memory is consumed for that purpose (eg: ~2MB per 128MB memory section on x86_64) This can even lead to extreme cases where system goes OOM because the physically hotplugged memory depletes the available memory before it is onlined. b) if the whole node is movable then we have off-node struct pages which has performance drawbacks. c) It might be there are no PMD_ALIGNED chunks so memmap array gets populated with base pages. This can be improved when CONFIG_SPARSEMEM_VMEMMAP is enabled. Vmemap page tables can map arbitrary memory. That means that we can reserve a part of the physically hotadded memory to back vmemmap page tables. This implementation uses the beginning of the hotplugged memory for that purpose. There are some non-obviously things to consider though. Vmemmap pages are allocated/freed during the memory hotplug events (add_memory_resource(), try_remove_memory()) when the memory is added/removed. This means that the reserved physical range is not online although it is used. The most obvious side effect is that pfn_to_online_page() returns NULL for those pfns. The current design expects that this should be OK as the hotplugged memory is considered a garbage until it is onlined. For example hibernation wouldn't save the content of those vmmemmaps into the image so it wouldn't be restored on resume but this should be OK as there no real content to recover anyway while metadata is reachable from other data structures (e.g. vmemmap page tables). The reserved space is therefore (de)initialized during the {on,off}line events (mhp_{de}init_memmap_on_memory). That is done by extracting page allocator independent initialization from the regular onlining path. The primary reason to handle the reserved space outside of {on,off}line_pages is to make each initialization specific to the purpose rather than special case them in a single function. As per above, the functions that are introduced are: - mhp_init_memmap_on_memory: Initializes vmemmap pages by calling move_pfn_range_to_zone(), calls kasan_add_zero_shadow(), and onlines as many sections as vmemmap pages fully span. - mhp_deinit_memmap_on_memory: Offlines as many sections as vmemmap pages fully span, removes the range from zhe zone by remove_pfn_range_from_zone(), and calls kasan_remove_zero_shadow() for the range. The new function memory_block_online() calls mhp_init_memmap_on_memory() before doing the actual online_pages(). Should online_pages() fail, we clean up by calling mhp_deinit_memmap_on_memory(). Adjusting of present_pages is done at the end once we know that online_pages() succedeed. On offline, memory_block_offline() needs to unaccount vmemmap pages from present_pages() before calling offline_pages(). This is necessary because offline_pages() tears down some structures based on the fact whether the node or the zone become empty. If offline_pages() fails, we account back vmemmap pages. If it succeeds, we call mhp_deinit_memmap_on_memory(). Hot-remove: We need to be careful when removing memory, as adding and removing memory needs to be done with the same granularity. To check that this assumption is not violated, we check the memory range we want to remove and if a) any memory block has vmemmap pages and b) the range spans more than a single memory block, we scream out loud and refuse to proceed. If all is good and the range was using memmap on memory (aka vmemmap pages), we construct an altmap structure so free_hugepage_table does the right thing and calls vmem_altmap_free instead of free_pagetable. Link: https://lkml.kernel.org/r/20210421102701.25051-5-osalvador@suse.de Signed-off-by: Oscar Salvador <osalvador@suse.de> Reviewed-by: David Hildenbrand <david@redhat.com> Acked-by: Michal Hocko <mhocko@suse.com> Cc: Anshuman Khandual <anshuman.khandual@arm.com> Cc: Pavel Tatashin <pasha.tatashin@soleen.com> Cc: Vlastimil Babka <vbabka@suse.cz> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2021-05-04 18:39:42 -07:00
config MHP_MEMMAP_ON_MEMORY
def_bool y
depends on MEMORY_HOTPLUG && SPARSEMEM_VMEMMAP
depends on ARCH_MHP_MEMMAP_ON_MEMORY_ENABLE
endif # MEMORY_HOTPLUG
config ARCH_MHP_MEMMAP_ON_MEMORY_ENABLE
bool
[PATCH] mm: split page table lock Christoph Lameter demonstrated very poor scalability on the SGI 512-way, with a many-threaded application which concurrently initializes different parts of a large anonymous area. This patch corrects that, by using a separate spinlock per page table page, to guard the page table entries in that page, instead of using the mm's single page_table_lock. (But even then, page_table_lock is still used to guard page table allocation, and anon_vma allocation.) In this implementation, the spinlock is tucked inside the struct page of the page table page: with a BUILD_BUG_ON in case it overflows - which it would in the case of 32-bit PA-RISC with spinlock debugging enabled. Splitting the lock is not quite for free: another cacheline access. Ideally, I suppose we would use split ptlock only for multi-threaded processes on multi-cpu machines; but deciding that dynamically would have its own costs. So for now enable it by config, at some number of cpus - since the Kconfig language doesn't support inequalities, let preprocessor compare that with NR_CPUS. But I don't think it's worth being user-configurable: for good testing of both split and unsplit configs, split now at 4 cpus, and perhaps change that to 8 later. There is a benefit even for singly threaded processes: kswapd can be attacking one part of the mm while another part is busy faulting. Signed-off-by: Hugh Dickins <hugh@veritas.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-10-29 18:16:40 -07:00
# Heavily threaded applications may benefit from splitting the mm-wide
# page_table_lock, so that faults on different parts of the user address
# space can be handled with less contention: split it at this NR_CPUS.
# Default to 4 for wider testing, though 8 might be more appropriate.
# ARM's adjust_pte (unused if VIPT) depends on mm-wide page_table_lock.
# PA-RISC 7xxx's spinlock_t would enlarge struct page from 32 to 44 bytes.
# SPARC32 allocates multiple pte tables within a single page, and therefore
# a per-page lock leads to problems when multiple tables need to be locked
# at the same time (e.g. copy_page_range()).
mm: stop ptlock enlarging struct page CONFIG_DEBUG_SPINLOCK adds 12 or 16 bytes to a 32- or 64-bit spinlock_t, and CONFIG_DEBUG_LOCK_ALLOC adds another 12 or 24 bytes to it: lockdep enables both of those, and CONFIG_LOCK_STAT adds 8 or 16 bytes to that. When 2.6.15 placed the split page table lock inside struct page (usually sized 32 or 56 bytes), only CONFIG_DEBUG_SPINLOCK was a possibility, and we ignored the enlargement (but fitted in CONFIG_GENERIC_LOCKBREAK's 4 by letting the spinlock_t occupy both page->private and page->mapping). Should these debugging options be allowed to double the size of a struct page, when only one minority use of the page (as a page table) needs to fit a spinlock in there? Perhaps not. Take the easy way out: switch off SPLIT_PTLOCK_CPUS when DEBUG_SPINLOCK or DEBUG_LOCK_ALLOC is in force. I've sometimes tried to be cleverer, kmallocing a cacheline for the spinlock when it doesn't fit, but given up each time. Falling back to mm->page_table_lock (as we do when ptlock is not split) lets lockdep check out the strictest path anyway. And now that some arches allow 8192 cpus, use 999999 for infinity. (What has this got to do with KSM swapping? It doesn't care about the size of struct page, but may care about random junk in page->mapping - to be explained separately later.) Signed-off-by: Hugh Dickins <hugh.dickins@tiscali.co.uk> Cc: Izik Eidus <ieidus@redhat.com> Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Nick Piggin <npiggin@suse.de> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Rik van Riel <riel@redhat.com> Cc: Lee Schermerhorn <Lee.Schermerhorn@hp.com> Cc: Andi Kleen <andi@firstfloor.org> Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: Wu Fengguang <fengguang.wu@intel.com> Cc: Minchan Kim <minchan.kim@gmail.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-12-14 18:59:02 -07:00
# DEBUG_SPINLOCK and DEBUG_LOCK_ALLOC spinlock_t also enlarge struct page.
[PATCH] mm: split page table lock Christoph Lameter demonstrated very poor scalability on the SGI 512-way, with a many-threaded application which concurrently initializes different parts of a large anonymous area. This patch corrects that, by using a separate spinlock per page table page, to guard the page table entries in that page, instead of using the mm's single page_table_lock. (But even then, page_table_lock is still used to guard page table allocation, and anon_vma allocation.) In this implementation, the spinlock is tucked inside the struct page of the page table page: with a BUILD_BUG_ON in case it overflows - which it would in the case of 32-bit PA-RISC with spinlock debugging enabled. Splitting the lock is not quite for free: another cacheline access. Ideally, I suppose we would use split ptlock only for multi-threaded processes on multi-cpu machines; but deciding that dynamically would have its own costs. So for now enable it by config, at some number of cpus - since the Kconfig language doesn't support inequalities, let preprocessor compare that with NR_CPUS. But I don't think it's worth being user-configurable: for good testing of both split and unsplit configs, split now at 4 cpus, and perhaps change that to 8 later. There is a benefit even for singly threaded processes: kswapd can be attacking one part of the mm while another part is busy faulting. Signed-off-by: Hugh Dickins <hugh@veritas.com> Signed-off-by: Andrew Morton <akpm@osdl.org> Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-10-29 18:16:40 -07:00
#
mm: turn USE_SPLIT_PTE_PTLOCKS / USE_SPLIT_PTE_PTLOCKS into Kconfig options Patch series "mm: split PTE/PMD PT table Kconfig cleanups+clarifications". This series is a follow up to the fixes: "[PATCH v1 0/2] mm/hugetlb: fix hugetlb vs. core-mm PT locking" When working on the fixes, I wondered why 8xx is fine (-> never uses split PT locks) and how PT locking even works properly with PMD page table sharing (-> always requires split PMD PT locks). Let's improve the split PT lock detection, make hugetlb properly depend on it and make 8xx bail out if it would ever get enabled by accident. As an alternative to patch #3 we could extend the Kconfig SPLIT_PTE_PTLOCKS option from patch #2 -- but enforcing it closer to the code that actually implements it feels a bit nicer for documentation purposes, and there is no need to actually disable it because it should always be disabled (!SMP). Did a bunch of cross-compilations to make sure that split PTE/PMD PT locks are still getting used where we would expect them. [1] https://lkml.kernel.org/r/20240725183955.2268884-1-david@redhat.com This patch (of 3): Let's clean that up a bit and prepare for depending on CONFIG_SPLIT_PMD_PTLOCKS in other Kconfig options. More cleanups would be reasonable (like the arch-specific "depends on" for CONFIG_SPLIT_PTE_PTLOCKS), but we'll leave that for another day. Link: https://lkml.kernel.org/r/20240726150728.3159964-1-david@redhat.com Link: https://lkml.kernel.org/r/20240726150728.3159964-2-david@redhat.com Signed-off-by: David Hildenbrand <david@redhat.com> Acked-by: Mike Rapoport (Microsoft) <rppt@kernel.org> Reviewed-by: Russell King (Oracle) <rmk+kernel@armlinux.org.uk> Reviewed-by: Qi Zheng <zhengqi.arch@bytedance.com> Cc: Alexander Viro <viro@zeniv.linux.org.uk> Cc: Borislav Petkov <bp@alien8.de> Cc: Boris Ostrovsky <boris.ostrovsky@oracle.com> Cc: Christian Brauner <brauner@kernel.org> Cc: Christophe Leroy <christophe.leroy@csgroup.eu> Cc: Dave Hansen <dave.hansen@linux.intel.com> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Ingo Molnar <mingo@redhat.com> Cc: Juergen Gross <jgross@suse.com> Cc: Michael Ellerman <mpe@ellerman.id.au> Cc: Muchun Song <muchun.song@linux.dev> Cc: "Naveen N. Rao" <naveen.n.rao@linux.ibm.com> Cc: Nicholas Piggin <npiggin@gmail.com> Cc: Oscar Salvador <osalvador@suse.de> Cc: Peter Xu <peterx@redhat.com> Cc: Thomas Gleixner <tglx@linutronix.de> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2024-07-26 08:07:26 -07:00
config SPLIT_PTE_PTLOCKS
def_bool y
depends on MMU
depends on SMP
mm: turn USE_SPLIT_PTE_PTLOCKS / USE_SPLIT_PTE_PTLOCKS into Kconfig options Patch series "mm: split PTE/PMD PT table Kconfig cleanups+clarifications". This series is a follow up to the fixes: "[PATCH v1 0/2] mm/hugetlb: fix hugetlb vs. core-mm PT locking" When working on the fixes, I wondered why 8xx is fine (-> never uses split PT locks) and how PT locking even works properly with PMD page table sharing (-> always requires split PMD PT locks). Let's improve the split PT lock detection, make hugetlb properly depend on it and make 8xx bail out if it would ever get enabled by accident. As an alternative to patch #3 we could extend the Kconfig SPLIT_PTE_PTLOCKS option from patch #2 -- but enforcing it closer to the code that actually implements it feels a bit nicer for documentation purposes, and there is no need to actually disable it because it should always be disabled (!SMP). Did a bunch of cross-compilations to make sure that split PTE/PMD PT locks are still getting used where we would expect them. [1] https://lkml.kernel.org/r/20240725183955.2268884-1-david@redhat.com This patch (of 3): Let's clean that up a bit and prepare for depending on CONFIG_SPLIT_PMD_PTLOCKS in other Kconfig options. More cleanups would be reasonable (like the arch-specific "depends on" for CONFIG_SPLIT_PTE_PTLOCKS), but we'll leave that for another day. Link: https://lkml.kernel.org/r/20240726150728.3159964-1-david@redhat.com Link: https://lkml.kernel.org/r/20240726150728.3159964-2-david@redhat.com Signed-off-by: David Hildenbrand <david@redhat.com> Acked-by: Mike Rapoport (Microsoft) <rppt@kernel.org> Reviewed-by: Russell King (Oracle) <rmk+kernel@armlinux.org.uk> Reviewed-by: Qi Zheng <zhengqi.arch@bytedance.com> Cc: Alexander Viro <viro@zeniv.linux.org.uk> Cc: Borislav Petkov <bp@alien8.de> Cc: Boris Ostrovsky <boris.ostrovsky@oracle.com> Cc: Christian Brauner <brauner@kernel.org> Cc: Christophe Leroy <christophe.leroy@csgroup.eu> Cc: Dave Hansen <dave.hansen@linux.intel.com> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Ingo Molnar <mingo@redhat.com> Cc: Juergen Gross <jgross@suse.com> Cc: Michael Ellerman <mpe@ellerman.id.au> Cc: Muchun Song <muchun.song@linux.dev> Cc: "Naveen N. Rao" <naveen.n.rao@linux.ibm.com> Cc: Nicholas Piggin <npiggin@gmail.com> Cc: Oscar Salvador <osalvador@suse.de> Cc: Peter Xu <peterx@redhat.com> Cc: Thomas Gleixner <tglx@linutronix.de> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2024-07-26 08:07:26 -07:00
depends on NR_CPUS >= 4
depends on !ARM || CPU_CACHE_VIPT
depends on !PARISC || PA20
depends on !SPARC32
mm: implement split page table lock for PMD level The basic idea is the same as with PTE level: the lock is embedded into struct page of table's page. We can't use mm->pmd_huge_pte to store pgtables for THP, since we don't take mm->page_table_lock anymore. Let's reuse page->lru of table's page for that. pgtable_pmd_page_ctor() returns true, if initialization is successful and false otherwise. Current implementation never fails, but assumption that constructor can fail will help to port it to -rt where spinlock_t is rather huge and cannot be embedded into struct page -- dynamic allocation is required. Signed-off-by: Naoya Horiguchi <n-horiguchi@ah.jp.nec.com> Signed-off-by: Kirill A. Shutemov <kirill.shutemov@linux.intel.com> Tested-by: Alex Thorlton <athorlton@sgi.com> Cc: Ingo Molnar <mingo@redhat.com> Cc: "Eric W . Biederman" <ebiederm@xmission.com> Cc: "Paul E . McKenney" <paulmck@linux.vnet.ibm.com> Cc: Al Viro <viro@zeniv.linux.org.uk> Cc: Andi Kleen <ak@linux.intel.com> Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Dave Hansen <dave.hansen@intel.com> Cc: Dave Jones <davej@redhat.com> Cc: David Howells <dhowells@redhat.com> Cc: Frederic Weisbecker <fweisbec@gmail.com> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Kees Cook <keescook@chromium.org> Cc: Mel Gorman <mgorman@suse.de> Cc: Michael Kerrisk <mtk.manpages@gmail.com> Cc: Oleg Nesterov <oleg@redhat.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Rik van Riel <riel@redhat.com> Cc: Robin Holt <robinmholt@gmail.com> Cc: Sedat Dilek <sedat.dilek@gmail.com> Cc: Srikar Dronamraju <srikar@linux.vnet.ibm.com> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Hugh Dickins <hughd@google.com> Reviewed-by: Steven Rostedt <rostedt@goodmis.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2013-11-14 15:31:07 -07:00
config ARCH_ENABLE_SPLIT_PMD_PTLOCK
bool
mm: implement split page table lock for PMD level The basic idea is the same as with PTE level: the lock is embedded into struct page of table's page. We can't use mm->pmd_huge_pte to store pgtables for THP, since we don't take mm->page_table_lock anymore. Let's reuse page->lru of table's page for that. pgtable_pmd_page_ctor() returns true, if initialization is successful and false otherwise. Current implementation never fails, but assumption that constructor can fail will help to port it to -rt where spinlock_t is rather huge and cannot be embedded into struct page -- dynamic allocation is required. Signed-off-by: Naoya Horiguchi <n-horiguchi@ah.jp.nec.com> Signed-off-by: Kirill A. Shutemov <kirill.shutemov@linux.intel.com> Tested-by: Alex Thorlton <athorlton@sgi.com> Cc: Ingo Molnar <mingo@redhat.com> Cc: "Eric W . Biederman" <ebiederm@xmission.com> Cc: "Paul E . McKenney" <paulmck@linux.vnet.ibm.com> Cc: Al Viro <viro@zeniv.linux.org.uk> Cc: Andi Kleen <ak@linux.intel.com> Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Dave Hansen <dave.hansen@intel.com> Cc: Dave Jones <davej@redhat.com> Cc: David Howells <dhowells@redhat.com> Cc: Frederic Weisbecker <fweisbec@gmail.com> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Kees Cook <keescook@chromium.org> Cc: Mel Gorman <mgorman@suse.de> Cc: Michael Kerrisk <mtk.manpages@gmail.com> Cc: Oleg Nesterov <oleg@redhat.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Rik van Riel <riel@redhat.com> Cc: Robin Holt <robinmholt@gmail.com> Cc: Sedat Dilek <sedat.dilek@gmail.com> Cc: Srikar Dronamraju <srikar@linux.vnet.ibm.com> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Hugh Dickins <hughd@google.com> Reviewed-by: Steven Rostedt <rostedt@goodmis.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2013-11-14 15:31:07 -07:00
mm: turn USE_SPLIT_PTE_PTLOCKS / USE_SPLIT_PTE_PTLOCKS into Kconfig options Patch series "mm: split PTE/PMD PT table Kconfig cleanups+clarifications". This series is a follow up to the fixes: "[PATCH v1 0/2] mm/hugetlb: fix hugetlb vs. core-mm PT locking" When working on the fixes, I wondered why 8xx is fine (-> never uses split PT locks) and how PT locking even works properly with PMD page table sharing (-> always requires split PMD PT locks). Let's improve the split PT lock detection, make hugetlb properly depend on it and make 8xx bail out if it would ever get enabled by accident. As an alternative to patch #3 we could extend the Kconfig SPLIT_PTE_PTLOCKS option from patch #2 -- but enforcing it closer to the code that actually implements it feels a bit nicer for documentation purposes, and there is no need to actually disable it because it should always be disabled (!SMP). Did a bunch of cross-compilations to make sure that split PTE/PMD PT locks are still getting used where we would expect them. [1] https://lkml.kernel.org/r/20240725183955.2268884-1-david@redhat.com This patch (of 3): Let's clean that up a bit and prepare for depending on CONFIG_SPLIT_PMD_PTLOCKS in other Kconfig options. More cleanups would be reasonable (like the arch-specific "depends on" for CONFIG_SPLIT_PTE_PTLOCKS), but we'll leave that for another day. Link: https://lkml.kernel.org/r/20240726150728.3159964-1-david@redhat.com Link: https://lkml.kernel.org/r/20240726150728.3159964-2-david@redhat.com Signed-off-by: David Hildenbrand <david@redhat.com> Acked-by: Mike Rapoport (Microsoft) <rppt@kernel.org> Reviewed-by: Russell King (Oracle) <rmk+kernel@armlinux.org.uk> Reviewed-by: Qi Zheng <zhengqi.arch@bytedance.com> Cc: Alexander Viro <viro@zeniv.linux.org.uk> Cc: Borislav Petkov <bp@alien8.de> Cc: Boris Ostrovsky <boris.ostrovsky@oracle.com> Cc: Christian Brauner <brauner@kernel.org> Cc: Christophe Leroy <christophe.leroy@csgroup.eu> Cc: Dave Hansen <dave.hansen@linux.intel.com> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Ingo Molnar <mingo@redhat.com> Cc: Juergen Gross <jgross@suse.com> Cc: Michael Ellerman <mpe@ellerman.id.au> Cc: Muchun Song <muchun.song@linux.dev> Cc: "Naveen N. Rao" <naveen.n.rao@linux.ibm.com> Cc: Nicholas Piggin <npiggin@gmail.com> Cc: Oscar Salvador <osalvador@suse.de> Cc: Peter Xu <peterx@redhat.com> Cc: Thomas Gleixner <tglx@linutronix.de> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2024-07-26 08:07:26 -07:00
config SPLIT_PMD_PTLOCKS
def_bool y
depends on SPLIT_PTE_PTLOCKS && ARCH_ENABLE_SPLIT_PMD_PTLOCK
#
# support for memory balloon
config MEMORY_BALLOON
bool
#
# support for memory balloon compaction
config BALLOON_COMPACTION
bool "Allow for balloon memory compaction/migration"
default y
depends on COMPACTION && MEMORY_BALLOON
help
Memory fragmentation introduced by ballooning might reduce
significantly the number of 2MB contiguous memory blocks that can be
used within a guest, thus imposing performance penalties associated
with the reduced number of transparent huge pages that could be used
by the guest workload. Allowing the compaction & migration for memory
pages enlisted as being part of memory balloon devices avoids the
scenario aforementioned and helps improving memory defragmentation.
#
# support for memory compaction
config COMPACTION
bool "Allow for memory compaction"
default y
select MIGRATION
depends on MMU
help
Compaction is the only memory management component to form
high order (larger physically contiguous) memory blocks
reliably. The page allocator relies on compaction heavily and
the lack of the feature can lead to unexpected OOM killer
invocations for high order memory requests. You shouldn't
disable this option unless there really is a strong reason for
it and then we would be really interested to hear about that at
linux-mm@kvack.org.
config COMPACT_UNEVICTABLE_DEFAULT
int
depends on COMPACTION
default 0 if PREEMPT_RT
default 1
mm: introduce Reported pages In order to pave the way for free page reporting in virtualized environments we will need a way to get pages out of the free lists and identify those pages after they have been returned. To accomplish this, this patch adds the concept of a Reported Buddy, which is essentially meant to just be the Uptodate flag used in conjunction with the Buddy page type. To prevent the reported pages from leaking outside of the buddy lists I added a check to clear the PageReported bit in the del_page_from_free_list function. As a result any reported page that is split, merged, or allocated will have the flag cleared prior to the PageBuddy value being cleared. The process for reporting pages is fairly simple. Once we free a page that meets the minimum order for page reporting we will schedule a worker thread to start 2s or more in the future. That worker thread will begin working from the lowest supported page reporting order up to MAX_ORDER - 1 pulling unreported pages from the free list and storing them in the scatterlist. When processing each individual free list it is necessary for the worker thread to release the zone lock when it needs to stop and report the full scatterlist of pages. To reduce the work of the next iteration the worker thread will rotate the free list so that the first unreported page in the free list becomes the first entry in the list. It will then call a reporting function providing information on how many entries are in the scatterlist. Once the function completes it will return the pages to the free area from which they were allocated and start over pulling more pages from the free areas until there are no longer enough pages to report on to keep the worker busy, or we have processed as many pages as were contained in the free area when we started processing the list. The worker thread will work in a round-robin fashion making its way though each zone requesting reporting, and through each reportable free list within that zone. Once all free areas within the zone have been processed it will check to see if there have been any requests for reporting while it was processing. If so it will reschedule the worker thread to start up again in roughly 2s and exit. Signed-off-by: Alexander Duyck <alexander.h.duyck@linux.intel.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Acked-by: Mel Gorman <mgorman@techsingularity.net> Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Dan Williams <dan.j.williams@intel.com> Cc: Dave Hansen <dave.hansen@intel.com> Cc: David Hildenbrand <david@redhat.com> Cc: Konrad Rzeszutek Wilk <konrad.wilk@oracle.com> Cc: Luiz Capitulino <lcapitulino@redhat.com> Cc: Matthew Wilcox <willy@infradead.org> Cc: Michael S. Tsirkin <mst@redhat.com> Cc: Michal Hocko <mhocko@kernel.org> Cc: Nitesh Narayan Lal <nitesh@redhat.com> Cc: Oscar Salvador <osalvador@suse.de> Cc: Pankaj Gupta <pagupta@redhat.com> Cc: Paolo Bonzini <pbonzini@redhat.com> Cc: Rik van Riel <riel@surriel.com> Cc: Vlastimil Babka <vbabka@suse.cz> Cc: Wei Wang <wei.w.wang@intel.com> Cc: Yang Zhang <yang.zhang.wz@gmail.com> Cc: wei qi <weiqi4@huawei.com> Link: http://lkml.kernel.org/r/20200211224635.29318.19750.stgit@localhost.localdomain Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2020-04-06 20:04:56 -07:00
#
# support for free page reporting
config PAGE_REPORTING
bool "Free page reporting"
help
Free page reporting allows for the incremental acquisition of
free pages from the buddy allocator for the purpose of reporting
those pages to another entity, such as a hypervisor, so that the
memory can be freed within the host for other uses.
#
# support for page migration
#
config MIGRATION
bool "Page migration"
default y
depends on (NUMA || ARCH_ENABLE_MEMORY_HOTREMOVE || COMPACTION || CMA) && MMU
help
Allows the migration of the physical location of pages of processes
while the virtual addresses are not changed. This is useful in
two situations. The first is on NUMA systems to put pages nearer
to the processors accessing. The second is when allocating huge
pages as migration can relocate pages to satisfy a huge page
allocation instead of reclaiming.
config DEVICE_MIGRATION
def_bool MIGRATION && ZONE_DEVICE
config ARCH_ENABLE_HUGEPAGE_MIGRATION
bool
config ARCH_ENABLE_THP_MIGRATION
bool
config HUGETLB_PAGE_SIZE_VARIABLE
def_bool n
help
Allows the pageblock_order value to be dynamic instead of just standard
HUGETLB_PAGE_ORDER when there are multiple HugeTLB page sizes available
on a platform.
Note that the pageblock_order cannot exceed MAX_PAGE_ORDER and will be
clamped down to MAX_PAGE_ORDER.
mm: enforce pageblock_order < MAX_ORDER Some places in the kernel don't really expect pageblock_order >= MAX_ORDER, and it looks like this is only possible in corner cases: 1) CONFIG_DEFERRED_STRUCT_PAGE_INIT we'll end up freeing pageblock_order pages via __free_pages_core(), which cannot possibly work. 2) find_zone_movable_pfns_for_nodes() will roundup the ZONE_MOVABLE start PFN to MAX_ORDER_NR_PAGES. Consequently with a bigger pageblock_order, we could have a single pageblock partially managed by two zones. 3) compaction code runs into __fragmentation_index() with order >= MAX_ORDER, when checking WARN_ON_ONCE(order >= MAX_ORDER). [1] 4) mm/page_reporting.c won't be reporting any pages with default page_reporting_order == pageblock_order, as we'll be skipping the reporting loop inside page_reporting_process_zone(). 5) __rmqueue_fallback() will never be able to steal with ALLOC_NOFRAGMENT. pageblock_order >= MAX_ORDER is weird either way: it's a pure optimization for making alloc_contig_range(), as used for allcoation of gigantic pages, a little more reliable to succeed. However, if there is demand for somewhat reliable allocation of gigantic pages, affected setups should be using CMA or boottime allocations instead. So let's make sure that pageblock_order < MAX_ORDER and simplify. [1] https://lkml.kernel.org/r/87r189a2ks.fsf@linux.ibm.com Link: https://lkml.kernel.org/r/20220214174132.219303-3-david@redhat.com Signed-off-by: David Hildenbrand <david@redhat.com> Reviewed-by: Zi Yan <ziy@nvidia.com> Cc: Aneesh Kumar K.V <aneesh.kumar@linux.ibm.com> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: Christoph Hellwig <hch@lst.de> Cc: Frank Rowand <frowand.list@gmail.com> Cc: John Garry via iommu <iommu@lists.linux-foundation.org> Cc: Marek Szyprowski <m.szyprowski@samsung.com> Cc: Michael Ellerman <mpe@ellerman.id.au> Cc: Michael S. Tsirkin <mst@redhat.com> Cc: Minchan Kim <minchan@kernel.org> Cc: Paul Mackerras <paulus@samba.org> Cc: Rob Herring <robh+dt@kernel.org> Cc: Robin Murphy <robin.murphy@arm.com> Cc: Vlastimil Babka <vbabka@suse.cz> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2022-03-22 14:43:20 -07:00
config CONTIG_ALLOC
def_bool (MEMORY_ISOLATION && COMPACTION) || CMA
mm: restrict the pcp batch scale factor to avoid too long latency In page allocator, PCP (Per-CPU Pageset) is refilled and drained in batches to increase page allocation throughput, reduce page allocation/freeing latency per page, and reduce zone lock contention. But too large batch size will cause too long maximal allocation/freeing latency, which may punish arbitrary users. So the default batch size is chosen carefully (in zone_batchsize(), the value is 63 for zone > 1GB) to avoid that. In commit 3b12e7e97938 ("mm/page_alloc: scale the number of pages that are batch freed"), the batch size will be scaled for large number of page freeing to improve page freeing performance and reduce zone lock contention. Similar optimization can be used for large number of pages allocation too. To find out a suitable max batch scale factor (that is, max effective batch size), some tests and measurement on some machines were done as follows. A set of debug patches are implemented as follows, - Set PCP high to be 2 * batch to reduce the effect of PCP high - Disable free batch size scaling to get the raw performance. - The code with zone lock held is extracted from rmqueue_bulk() and free_pcppages_bulk() to 2 separate functions to make it easy to measure the function run time with ftrace function_graph tracer. - The batch size is hard coded to be 63 (default), 127, 255, 511, 1023, 2047, 4095. Then will-it-scale/page_fault1 is used to generate the page allocation/freeing workload. The page allocation/freeing throughput (page/s) is measured via will-it-scale. The page allocation/freeing average latency (alloc/free latency avg, in us) and allocation/freeing latency at 99 percentile (alloc/free latency 99%, in us) are measured with ftrace function_graph tracer. The test results are as follows, Sapphire Rapids Server ====================== Batch throughput free latency free latency alloc latency alloc latency page/s avg / us 99% / us avg / us 99% / us ----- ---------- ------------ ------------ ------------- ------------- 63 513633.4 2.33 3.57 2.67 6.83 127 517616.7 4.35 6.65 4.22 13.03 255 520822.8 8.29 13.32 7.52 25.24 511 524122.0 15.79 23.42 14.02 49.35 1023 525980.5 30.25 44.19 25.36 94.88 2047 526793.6 59.39 84.50 45.22 140.81 Ice Lake Server =============== Batch throughput free latency free latency alloc latency alloc latency page/s avg / us 99% / us avg / us 99% / us ----- ---------- ------------ ------------ ------------- ------------- 63 620210.3 2.21 3.68 2.02 4.35 127 627003.0 4.09 6.86 3.51 8.28 255 630777.5 7.70 13.50 6.17 15.97 511 633651.5 14.85 22.62 11.66 31.08 1023 637071.1 28.55 42.02 20.81 54.36 2047 638089.7 56.54 84.06 39.28 91.68 Cascade Lake Server =================== Batch throughput free latency free latency alloc latency alloc latency page/s avg / us 99% / us avg / us 99% / us ----- ---------- ------------ ------------ ------------- ------------- 63 404706.7 3.29 5.03 3.53 4.75 127 422475.2 6.12 9.09 6.36 8.76 255 411522.2 11.68 16.97 10.90 16.39 511 428124.1 22.54 31.28 19.86 32.25 1023 414718.4 43.39 62.52 40.00 66.33 2047 429848.7 86.64 120.34 71.14 106.08 Commet Lake Desktop =================== Batch throughput free latency free latency alloc latency alloc latency page/s avg / us 99% / us avg / us 99% / us ----- ---------- ------------ ------------ ------------- ------------- 63 795183.13 2.18 3.55 2.03 3.05 127 803067.85 3.91 6.56 3.85 5.52 255 812771.10 7.35 10.80 7.14 10.20 511 817723.48 14.17 27.54 13.43 30.31 1023 818870.19 27.72 40.10 27.89 46.28 Coffee Lake Desktop =================== Batch throughput free latency free latency alloc latency alloc latency page/s avg / us 99% / us avg / us 99% / us ----- ---------- ------------ ------------ ------------- ------------- 63 510542.8 3.13 4.40 2.48 3.43 127 514288.6 5.97 7.89 4.65 6.04 255 516889.7 11.86 15.58 8.96 12.55 511 519802.4 23.10 28.81 16.95 26.19 1023 520802.7 45.30 52.51 33.19 45.95 2047 519997.1 90.63 104.00 65.26 81.74 From the above data, to restrict the allocation/freeing latency to be less than 100 us in most times, the max batch scale factor needs to be less than or equal to 5. Although it is reasonable to use 5 as max batch scale factor for the systems tested, there are also slower systems. Where smaller value should be used to constrain the page allocation/freeing latency. So, in this patch, a new kconfig option (PCP_BATCH_SCALE_MAX) is added to set the max batch scale factor. Whose default value is 5, and users can reduce it when necessary. Link: https://lkml.kernel.org/r/20231016053002.756205-5-ying.huang@intel.com Signed-off-by: "Huang, Ying" <ying.huang@intel.com> Acked-by: Andrew Morton <akpm@linux-foundation.org> Acked-by: Mel Gorman <mgorman@techsingularity.net> Cc: Vlastimil Babka <vbabka@suse.cz> Cc: David Hildenbrand <david@redhat.com> Cc: Johannes Weiner <jweiner@redhat.com> Cc: Dave Hansen <dave.hansen@linux.intel.com> Cc: Michal Hocko <mhocko@suse.com> Cc: Pavel Tatashin <pasha.tatashin@soleen.com> Cc: Matthew Wilcox <willy@infradead.org> Cc: Christoph Lameter <cl@linux.com> Cc: Arjan van de Ven <arjan@linux.intel.com> Cc: Sudeep Holla <sudeep.holla@arm.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2023-10-15 22:29:57 -07:00
config PCP_BATCH_SCALE_MAX
int "Maximum scale factor of PCP (Per-CPU pageset) batch allocate/free"
default 5
range 0 6
help
In page allocator, PCP (Per-CPU pageset) is refilled and drained in
batches. The batch number is scaled automatically to improve page
allocation/free throughput. But too large scale factor may hurt
latency. This option sets the upper limit of scale factor to limit
the maximum latency.
config PHYS_ADDR_T_64BIT
def_bool 64BIT
config BOUNCE
bool "Enable bounce buffers"
default y
depends on BLOCK && MMU && HIGHMEM
help
Enable bounce buffers for devices that cannot access the full range of
memory available to the CPU. Enabled by default when HIGHMEM is
selected, but you may say n to override this.
mmu-notifiers: core With KVM/GFP/XPMEM there isn't just the primary CPU MMU pointing to pages. There are secondary MMUs (with secondary sptes and secondary tlbs) too. sptes in the kvm case are shadow pagetables, but when I say spte in mmu-notifier context, I mean "secondary pte". In GRU case there's no actual secondary pte and there's only a secondary tlb because the GRU secondary MMU has no knowledge about sptes and every secondary tlb miss event in the MMU always generates a page fault that has to be resolved by the CPU (this is not the case of KVM where the a secondary tlb miss will walk sptes in hardware and it will refill the secondary tlb transparently to software if the corresponding spte is present). The same way zap_page_range has to invalidate the pte before freeing the page, the spte (and secondary tlb) must also be invalidated before any page is freed and reused. Currently we take a page_count pin on every page mapped by sptes, but that means the pages can't be swapped whenever they're mapped by any spte because they're part of the guest working set. Furthermore a spte unmap event can immediately lead to a page to be freed when the pin is released (so requiring the same complex and relatively slow tlb_gather smp safe logic we have in zap_page_range and that can be avoided completely if the spte unmap event doesn't require an unpin of the page previously mapped in the secondary MMU). The mmu notifiers allow kvm/GRU/XPMEM to attach to the tsk->mm and know when the VM is swapping or freeing or doing anything on the primary MMU so that the secondary MMU code can drop sptes before the pages are freed, avoiding all page pinning and allowing 100% reliable swapping of guest physical address space. Furthermore it avoids the code that teardown the mappings of the secondary MMU, to implement a logic like tlb_gather in zap_page_range that would require many IPI to flush other cpu tlbs, for each fixed number of spte unmapped. To make an example: if what happens on the primary MMU is a protection downgrade (from writeable to wrprotect) the secondary MMU mappings will be invalidated, and the next secondary-mmu-page-fault will call get_user_pages and trigger a do_wp_page through get_user_pages if it called get_user_pages with write=1, and it'll re-establishing an updated spte or secondary-tlb-mapping on the copied page. Or it will setup a readonly spte or readonly tlb mapping if it's a guest-read, if it calls get_user_pages with write=0. This is just an example. This allows to map any page pointed by any pte (and in turn visible in the primary CPU MMU), into a secondary MMU (be it a pure tlb like GRU, or an full MMU with both sptes and secondary-tlb like the shadow-pagetable layer with kvm), or a remote DMA in software like XPMEM (hence needing of schedule in XPMEM code to send the invalidate to the remote node, while no need to schedule in kvm/gru as it's an immediate event like invalidating primary-mmu pte). At least for KVM without this patch it's impossible to swap guests reliably. And having this feature and removing the page pin allows several other optimizations that simplify life considerably. Dependencies: 1) mm_take_all_locks() to register the mmu notifier when the whole VM isn't doing anything with "mm". This allows mmu notifier users to keep track if the VM is in the middle of the invalidate_range_begin/end critical section with an atomic counter incraese in range_begin and decreased in range_end. No secondary MMU page fault is allowed to map any spte or secondary tlb reference, while the VM is in the middle of range_begin/end as any page returned by get_user_pages in that critical section could later immediately be freed without any further ->invalidate_page notification (invalidate_range_begin/end works on ranges and ->invalidate_page isn't called immediately before freeing the page). To stop all page freeing and pagetable overwrites the mmap_sem must be taken in write mode and all other anon_vma/i_mmap locks must be taken too. 2) It'd be a waste to add branches in the VM if nobody could possibly run KVM/GRU/XPMEM on the kernel, so mmu notifiers will only enabled if CONFIG_KVM=m/y. In the current kernel kvm won't yet take advantage of mmu notifiers, but this already allows to compile a KVM external module against a kernel with mmu notifiers enabled and from the next pull from kvm.git we'll start using them. And GRU/XPMEM will also be able to continue the development by enabling KVM=m in their config, until they submit all GRU/XPMEM GPLv2 code to the mainline kernel. Then they can also enable MMU_NOTIFIERS in the same way KVM does it (even if KVM=n). This guarantees nobody selects MMU_NOTIFIER=y if KVM and GRU and XPMEM are all =n. The mmu_notifier_register call can fail because mm_take_all_locks may be interrupted by a signal and return -EINTR. Because mmu_notifier_reigster is used when a driver startup, a failure can be gracefully handled. Here an example of the change applied to kvm to register the mmu notifiers. Usually when a driver startups other allocations are required anyway and -ENOMEM failure paths exists already. struct kvm *kvm_arch_create_vm(void) { struct kvm *kvm = kzalloc(sizeof(struct kvm), GFP_KERNEL); + int err; if (!kvm) return ERR_PTR(-ENOMEM); INIT_LIST_HEAD(&kvm->arch.active_mmu_pages); + kvm->arch.mmu_notifier.ops = &kvm_mmu_notifier_ops; + err = mmu_notifier_register(&kvm->arch.mmu_notifier, current->mm); + if (err) { + kfree(kvm); + return ERR_PTR(err); + } + return kvm; } mmu_notifier_unregister returns void and it's reliable. The patch also adds a few needed but missing includes that would prevent kernel to compile after these changes on non-x86 archs (x86 didn't need them by luck). [akpm@linux-foundation.org: coding-style fixes] [akpm@linux-foundation.org: fix mm/filemap_xip.c build] [akpm@linux-foundation.org: fix mm/mmu_notifier.c build] Signed-off-by: Andrea Arcangeli <andrea@qumranet.com> Signed-off-by: Nick Piggin <npiggin@suse.de> Signed-off-by: Christoph Lameter <cl@linux-foundation.org> Cc: Jack Steiner <steiner@sgi.com> Cc: Robin Holt <holt@sgi.com> Cc: Nick Piggin <npiggin@suse.de> Cc: Peter Zijlstra <a.p.zijlstra@chello.nl> Cc: Kanoj Sarcar <kanojsarcar@yahoo.com> Cc: Roland Dreier <rdreier@cisco.com> Cc: Steve Wise <swise@opengridcomputing.com> Cc: Avi Kivity <avi@qumranet.com> Cc: Hugh Dickins <hugh@veritas.com> Cc: Rusty Russell <rusty@rustcorp.com.au> Cc: Anthony Liguori <aliguori@us.ibm.com> Cc: Chris Wright <chrisw@redhat.com> Cc: Marcelo Tosatti <marcelo@kvack.org> Cc: Eric Dumazet <dada1@cosmosbay.com> Cc: "Paul E. McKenney" <paulmck@us.ibm.com> Cc: Izik Eidus <izike@qumranet.com> Cc: Anthony Liguori <aliguori@us.ibm.com> Cc: Rik van Riel <riel@redhat.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2008-07-28 15:46:29 -07:00
config MMU_NOTIFIER
bool
mm/mmu_notifier: add an interval tree notifier Of the 13 users of mmu_notifiers, 8 of them use only invalidate_range_start/end() and immediately intersect the mmu_notifier_range with some kind of internal list of VAs. 4 use an interval tree (i915_gem, radeon_mn, umem_odp, hfi1). 4 use a linked list of some kind (scif_dma, vhost, gntdev, hmm) And the remaining 5 either don't use invalidate_range_start() or do some special thing with it. It turns out that building a correct scheme with an interval tree is pretty complicated, particularly if the use case is synchronizing against another thread doing get_user_pages(). Many of these implementations have various subtle and difficult to fix races. This approach puts the interval tree as common code at the top of the mmu notifier call tree and implements a shareable locking scheme. It includes: - An interval tree tracking VA ranges, with per-range callbacks - A read/write locking scheme for the interval tree that avoids sleeping in the notifier path (for OOM killer) - A sequence counter based collision-retry locking scheme to tell device page fault that a VA range is being concurrently invalidated. This is based on various ideas: - hmm accumulates invalidated VA ranges and releases them when all invalidates are done, via active_invalidate_ranges count. This approach avoids having to intersect the interval tree twice (as umem_odp does) at the potential cost of a longer device page fault. - kvm/umem_odp use a sequence counter to drive the collision retry, via invalidate_seq - a deferred work todo list on unlock scheme like RTNL, via deferred_list. This makes adding/removing interval tree members more deterministic - seqlock, except this version makes the seqlock idea multi-holder on the write side by protecting it with active_invalidate_ranges and a spinlock To minimize MM overhead when only the interval tree is being used, the entire SRCU and hlist overheads are dropped using some simple branches. Similarly the interval tree overhead is dropped when in hlist mode. The overhead from the mandatory spinlock is broadly the same as most of existing users which already had a lock (or two) of some sort on the invalidation path. Link: https://lore.kernel.org/r/20191112202231.3856-3-jgg@ziepe.ca Acked-by: Christian König <christian.koenig@amd.com> Tested-by: Philip Yang <Philip.Yang@amd.com> Tested-by: Ralph Campbell <rcampbell@nvidia.com> Reviewed-by: John Hubbard <jhubbard@nvidia.com> Reviewed-by: Christoph Hellwig <hch@lst.de> Signed-off-by: Jason Gunthorpe <jgg@mellanox.com>
2019-11-12 13:22:19 -07:00
select INTERVAL_TREE
nommu: make the initial mmap allocation excess behaviour Kconfig configurable NOMMU mmap() has an option controlled by a sysctl variable that determines whether the allocations made by do_mmap_private() should have the excess space trimmed off and returned to the allocator. Make the initial setting of this variable a Kconfig configuration option. The reason there can be excess space is that the allocator only allocates in power-of-2 size chunks, but mmap()'s can be made in sizes that aren't a power of 2. There are two alternatives: (1) Keep the excess as dead space. The dead space then remains unused for the lifetime of the mapping. Mappings of shared objects such as libc, ld.so or busybox's text segment may retain their dead space forever. (2) Return the excess to the allocator. This means that the dead space is limited to less than a page per mapping, but it means that for a transient process, there's more chance of fragmentation as the excess space may be reused fairly quickly. During the boot process, a lot of transient processes are created, and this can cause a lot of fragmentation as the pagecache and various slabs grow greatly during this time. By turning off the trimming of excess space during boot and disabling batching of frees, Coldfire can manage to boot. A better way of doing things might be to have /sbin/init turn this option off. By that point libc, ld.so and init - which are all long-duration processes - have all been loaded and trimmed. Reported-by: Lanttor Guo <lanttor.guo@freescale.com> Signed-off-by: David Howells <dhowells@redhat.com> Tested-by: Lanttor Guo <lanttor.guo@freescale.com> Cc: Greg Ungerer <gerg@snapgear.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-05-06 16:03:05 -07:00
ksm: the mm interface to ksm This patch presents the mm interface to a dummy version of ksm.c, for better scrutiny of that interface: the real ksm.c follows later. When CONFIG_KSM is not set, madvise(2) reject MADV_MERGEABLE and MADV_UNMERGEABLE with EINVAL, since that seems more helpful than pretending that they can be serviced. But when CONFIG_KSM=y, accept them even if KSM is not currently running, and even on areas which KSM will not touch (e.g. hugetlb or shared file or special driver mappings). Like other madvices, report ENOMEM despite success if any area in the range is unmapped, and use EAGAIN to report out of memory. Define vma flag VM_MERGEABLE to identify an area on which KSM may try merging pages: leave it to ksm_madvise() to decide whether to set it. Define mm flag MMF_VM_MERGEABLE to identify an mm which might contain VM_MERGEABLE areas, to minimize callouts when forking or exiting. Based upon earlier patches by Chris Wright and Izik Eidus. Signed-off-by: Hugh Dickins <hugh.dickins@tiscali.co.uk> Signed-off-by: Chris Wright <chrisw@redhat.com> Signed-off-by: Izik Eidus <ieidus@redhat.com> Cc: Michael Kerrisk <mtk.manpages@gmail.com> Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Rik van Riel <riel@redhat.com> Cc: Wu Fengguang <fengguang.wu@intel.com> Cc: Balbir Singh <balbir@in.ibm.com> Cc: Hugh Dickins <hugh.dickins@tiscali.co.uk> Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: Lee Schermerhorn <lee.schermerhorn@hp.com> Cc: Avi Kivity <avi@redhat.com> Cc: Nick Piggin <nickpiggin@yahoo.com.au> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-09-21 17:01:57 -07:00
config KSM
bool "Enable KSM for page merging"
depends on MMU
select XXHASH
ksm: the mm interface to ksm This patch presents the mm interface to a dummy version of ksm.c, for better scrutiny of that interface: the real ksm.c follows later. When CONFIG_KSM is not set, madvise(2) reject MADV_MERGEABLE and MADV_UNMERGEABLE with EINVAL, since that seems more helpful than pretending that they can be serviced. But when CONFIG_KSM=y, accept them even if KSM is not currently running, and even on areas which KSM will not touch (e.g. hugetlb or shared file or special driver mappings). Like other madvices, report ENOMEM despite success if any area in the range is unmapped, and use EAGAIN to report out of memory. Define vma flag VM_MERGEABLE to identify an area on which KSM may try merging pages: leave it to ksm_madvise() to decide whether to set it. Define mm flag MMF_VM_MERGEABLE to identify an mm which might contain VM_MERGEABLE areas, to minimize callouts when forking or exiting. Based upon earlier patches by Chris Wright and Izik Eidus. Signed-off-by: Hugh Dickins <hugh.dickins@tiscali.co.uk> Signed-off-by: Chris Wright <chrisw@redhat.com> Signed-off-by: Izik Eidus <ieidus@redhat.com> Cc: Michael Kerrisk <mtk.manpages@gmail.com> Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Rik van Riel <riel@redhat.com> Cc: Wu Fengguang <fengguang.wu@intel.com> Cc: Balbir Singh <balbir@in.ibm.com> Cc: Hugh Dickins <hugh.dickins@tiscali.co.uk> Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: Lee Schermerhorn <lee.schermerhorn@hp.com> Cc: Avi Kivity <avi@redhat.com> Cc: Nick Piggin <nickpiggin@yahoo.com.au> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-09-21 17:01:57 -07:00
help
Enable Kernel Samepage Merging: KSM periodically scans those areas
of an application's address space that an app has advised may be
mergeable. When it finds pages of identical content, it replaces
the many instances by a single page with that content, so
ksm: the mm interface to ksm This patch presents the mm interface to a dummy version of ksm.c, for better scrutiny of that interface: the real ksm.c follows later. When CONFIG_KSM is not set, madvise(2) reject MADV_MERGEABLE and MADV_UNMERGEABLE with EINVAL, since that seems more helpful than pretending that they can be serviced. But when CONFIG_KSM=y, accept them even if KSM is not currently running, and even on areas which KSM will not touch (e.g. hugetlb or shared file or special driver mappings). Like other madvices, report ENOMEM despite success if any area in the range is unmapped, and use EAGAIN to report out of memory. Define vma flag VM_MERGEABLE to identify an area on which KSM may try merging pages: leave it to ksm_madvise() to decide whether to set it. Define mm flag MMF_VM_MERGEABLE to identify an mm which might contain VM_MERGEABLE areas, to minimize callouts when forking or exiting. Based upon earlier patches by Chris Wright and Izik Eidus. Signed-off-by: Hugh Dickins <hugh.dickins@tiscali.co.uk> Signed-off-by: Chris Wright <chrisw@redhat.com> Signed-off-by: Izik Eidus <ieidus@redhat.com> Cc: Michael Kerrisk <mtk.manpages@gmail.com> Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Rik van Riel <riel@redhat.com> Cc: Wu Fengguang <fengguang.wu@intel.com> Cc: Balbir Singh <balbir@in.ibm.com> Cc: Hugh Dickins <hugh.dickins@tiscali.co.uk> Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: Lee Schermerhorn <lee.schermerhorn@hp.com> Cc: Avi Kivity <avi@redhat.com> Cc: Nick Piggin <nickpiggin@yahoo.com.au> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-09-21 17:01:57 -07:00
saving memory until one or another app needs to modify the content.
Recommended for use with KVM, or with other duplicative applications.
See Documentation/mm/ksm.rst for more information: KSM is inactive
until a program has madvised that an area is MADV_MERGEABLE, and
root has set /sys/kernel/mm/ksm/run to 1 (if CONFIG_SYSFS is set).
ksm: the mm interface to ksm This patch presents the mm interface to a dummy version of ksm.c, for better scrutiny of that interface: the real ksm.c follows later. When CONFIG_KSM is not set, madvise(2) reject MADV_MERGEABLE and MADV_UNMERGEABLE with EINVAL, since that seems more helpful than pretending that they can be serviced. But when CONFIG_KSM=y, accept them even if KSM is not currently running, and even on areas which KSM will not touch (e.g. hugetlb or shared file or special driver mappings). Like other madvices, report ENOMEM despite success if any area in the range is unmapped, and use EAGAIN to report out of memory. Define vma flag VM_MERGEABLE to identify an area on which KSM may try merging pages: leave it to ksm_madvise() to decide whether to set it. Define mm flag MMF_VM_MERGEABLE to identify an mm which might contain VM_MERGEABLE areas, to minimize callouts when forking or exiting. Based upon earlier patches by Chris Wright and Izik Eidus. Signed-off-by: Hugh Dickins <hugh.dickins@tiscali.co.uk> Signed-off-by: Chris Wright <chrisw@redhat.com> Signed-off-by: Izik Eidus <ieidus@redhat.com> Cc: Michael Kerrisk <mtk.manpages@gmail.com> Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Rik van Riel <riel@redhat.com> Cc: Wu Fengguang <fengguang.wu@intel.com> Cc: Balbir Singh <balbir@in.ibm.com> Cc: Hugh Dickins <hugh.dickins@tiscali.co.uk> Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: Lee Schermerhorn <lee.schermerhorn@hp.com> Cc: Avi Kivity <avi@redhat.com> Cc: Nick Piggin <nickpiggin@yahoo.com.au> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-09-21 17:01:57 -07:00
config DEFAULT_MMAP_MIN_ADDR
int "Low address space to protect from user allocation"
depends on MMU
default 4096
help
This is the portion of low virtual memory which should be protected
from userspace allocation. Keeping a user from writing to low pages
can help reduce the impact of kernel NULL pointer bugs.
For most arm64, ppc64 and x86 users with lots of address space
a value of 65536 is reasonable and should cause no problems.
On arm and other archs it should not be higher than 32768.
Programs which use vm86 functionality or have some need to map
this low address space will need CAP_SYS_RAWIO or disable this
protection by setting the value to 0.
This value can be changed after boot using the
/proc/sys/vm/mmap_min_addr tunable.
config ARCH_SUPPORTS_MEMORY_FAILURE
bool
HWPOISON: The high level memory error handler in the VM v7 Add the high level memory handler that poisons pages that got corrupted by hardware (typically by a two bit flip in a DIMM or a cache) on the Linux level. The goal is to prevent everyone from accessing these pages in the future. This done at the VM level by marking a page hwpoisoned and doing the appropriate action based on the type of page it is. The code that does this is portable and lives in mm/memory-failure.c To quote the overview comment: High level machine check handler. Handles pages reported by the hardware as being corrupted usually due to a 2bit ECC memory or cache failure. This focuses on pages detected as corrupted in the background. When the current CPU tries to consume corruption the currently running process can just be killed directly instead. This implies that if the error cannot be handled for some reason it's safe to just ignore it because no corruption has been consumed yet. Instead when that happens another machine check will happen. Handles page cache pages in various states. The tricky part here is that we can access any page asynchronous to other VM users, because memory failures could happen anytime and anywhere, possibly violating some of their assumptions. This is why this code has to be extremely careful. Generally it tries to use normal locking rules, as in get the standard locks, even if that means the error handling takes potentially a long time. Some of the operations here are somewhat inefficient and have non linear algorithmic complexity, because the data structures have not been optimized for this case. This is in particular the case for the mapping from a vma to a process. Since this case is expected to be rare we hope we can get away with this. There are in principle two strategies to kill processes on poison: - just unmap the data and wait for an actual reference before killing - kill as soon as corruption is detected. Both have advantages and disadvantages and should be used in different situations. Right now both are implemented and can be switched with a new sysctl vm.memory_failure_early_kill The default is early kill. The patch does some rmap data structure walking on its own to collect processes to kill. This is unusual because normally all rmap data structure knowledge is in rmap.c only. I put it here for now to keep everything together and rmap knowledge has been seeping out anyways Includes contributions from Johannes Weiner, Chris Mason, Fengguang Wu, Nick Piggin (who did a lot of great work) and others. Cc: npiggin@suse.de Cc: riel@redhat.com Signed-off-by: Andi Kleen <ak@linux.intel.com> Acked-by: Rik van Riel <riel@redhat.com> Reviewed-by: Hidehiro Kawai <hidehiro.kawai.ez@hitachi.com>
2009-09-16 02:50:15 -07:00
config MEMORY_FAILURE
depends on MMU
depends on ARCH_SUPPORTS_MEMORY_FAILURE
HWPOISON: The high level memory error handler in the VM v7 Add the high level memory handler that poisons pages that got corrupted by hardware (typically by a two bit flip in a DIMM or a cache) on the Linux level. The goal is to prevent everyone from accessing these pages in the future. This done at the VM level by marking a page hwpoisoned and doing the appropriate action based on the type of page it is. The code that does this is portable and lives in mm/memory-failure.c To quote the overview comment: High level machine check handler. Handles pages reported by the hardware as being corrupted usually due to a 2bit ECC memory or cache failure. This focuses on pages detected as corrupted in the background. When the current CPU tries to consume corruption the currently running process can just be killed directly instead. This implies that if the error cannot be handled for some reason it's safe to just ignore it because no corruption has been consumed yet. Instead when that happens another machine check will happen. Handles page cache pages in various states. The tricky part here is that we can access any page asynchronous to other VM users, because memory failures could happen anytime and anywhere, possibly violating some of their assumptions. This is why this code has to be extremely careful. Generally it tries to use normal locking rules, as in get the standard locks, even if that means the error handling takes potentially a long time. Some of the operations here are somewhat inefficient and have non linear algorithmic complexity, because the data structures have not been optimized for this case. This is in particular the case for the mapping from a vma to a process. Since this case is expected to be rare we hope we can get away with this. There are in principle two strategies to kill processes on poison: - just unmap the data and wait for an actual reference before killing - kill as soon as corruption is detected. Both have advantages and disadvantages and should be used in different situations. Right now both are implemented and can be switched with a new sysctl vm.memory_failure_early_kill The default is early kill. The patch does some rmap data structure walking on its own to collect processes to kill. This is unusual because normally all rmap data structure knowledge is in rmap.c only. I put it here for now to keep everything together and rmap knowledge has been seeping out anyways Includes contributions from Johannes Weiner, Chris Mason, Fengguang Wu, Nick Piggin (who did a lot of great work) and others. Cc: npiggin@suse.de Cc: riel@redhat.com Signed-off-by: Andi Kleen <ak@linux.intel.com> Acked-by: Rik van Riel <riel@redhat.com> Reviewed-by: Hidehiro Kawai <hidehiro.kawai.ez@hitachi.com>
2009-09-16 02:50:15 -07:00
bool "Enable recovery from hardware memory errors"
select MEMORY_ISOLATION
select RAS
HWPOISON: The high level memory error handler in the VM v7 Add the high level memory handler that poisons pages that got corrupted by hardware (typically by a two bit flip in a DIMM or a cache) on the Linux level. The goal is to prevent everyone from accessing these pages in the future. This done at the VM level by marking a page hwpoisoned and doing the appropriate action based on the type of page it is. The code that does this is portable and lives in mm/memory-failure.c To quote the overview comment: High level machine check handler. Handles pages reported by the hardware as being corrupted usually due to a 2bit ECC memory or cache failure. This focuses on pages detected as corrupted in the background. When the current CPU tries to consume corruption the currently running process can just be killed directly instead. This implies that if the error cannot be handled for some reason it's safe to just ignore it because no corruption has been consumed yet. Instead when that happens another machine check will happen. Handles page cache pages in various states. The tricky part here is that we can access any page asynchronous to other VM users, because memory failures could happen anytime and anywhere, possibly violating some of their assumptions. This is why this code has to be extremely careful. Generally it tries to use normal locking rules, as in get the standard locks, even if that means the error handling takes potentially a long time. Some of the operations here are somewhat inefficient and have non linear algorithmic complexity, because the data structures have not been optimized for this case. This is in particular the case for the mapping from a vma to a process. Since this case is expected to be rare we hope we can get away with this. There are in principle two strategies to kill processes on poison: - just unmap the data and wait for an actual reference before killing - kill as soon as corruption is detected. Both have advantages and disadvantages and should be used in different situations. Right now both are implemented and can be switched with a new sysctl vm.memory_failure_early_kill The default is early kill. The patch does some rmap data structure walking on its own to collect processes to kill. This is unusual because normally all rmap data structure knowledge is in rmap.c only. I put it here for now to keep everything together and rmap knowledge has been seeping out anyways Includes contributions from Johannes Weiner, Chris Mason, Fengguang Wu, Nick Piggin (who did a lot of great work) and others. Cc: npiggin@suse.de Cc: riel@redhat.com Signed-off-by: Andi Kleen <ak@linux.intel.com> Acked-by: Rik van Riel <riel@redhat.com> Reviewed-by: Hidehiro Kawai <hidehiro.kawai.ez@hitachi.com>
2009-09-16 02:50:15 -07:00
help
Enables code to recover from some memory failures on systems
with MCA recovery. This allows a system to continue running
even when some of its memory has uncorrected errors. This requires
special hardware support and typically ECC memory.
config HWPOISON_INJECT
tristate "HWPoison pages injector"
depends on MEMORY_FAILURE && DEBUG_KERNEL && PROC_FS
select PROC_PAGE_MONITOR
nommu: make the initial mmap allocation excess behaviour Kconfig configurable NOMMU mmap() has an option controlled by a sysctl variable that determines whether the allocations made by do_mmap_private() should have the excess space trimmed off and returned to the allocator. Make the initial setting of this variable a Kconfig configuration option. The reason there can be excess space is that the allocator only allocates in power-of-2 size chunks, but mmap()'s can be made in sizes that aren't a power of 2. There are two alternatives: (1) Keep the excess as dead space. The dead space then remains unused for the lifetime of the mapping. Mappings of shared objects such as libc, ld.so or busybox's text segment may retain their dead space forever. (2) Return the excess to the allocator. This means that the dead space is limited to less than a page per mapping, but it means that for a transient process, there's more chance of fragmentation as the excess space may be reused fairly quickly. During the boot process, a lot of transient processes are created, and this can cause a lot of fragmentation as the pagecache and various slabs grow greatly during this time. By turning off the trimming of excess space during boot and disabling batching of frees, Coldfire can manage to boot. A better way of doing things might be to have /sbin/init turn this option off. By that point libc, ld.so and init - which are all long-duration processes - have all been loaded and trimmed. Reported-by: Lanttor Guo <lanttor.guo@freescale.com> Signed-off-by: David Howells <dhowells@redhat.com> Tested-by: Lanttor Guo <lanttor.guo@freescale.com> Cc: Greg Ungerer <gerg@snapgear.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-05-06 16:03:05 -07:00
config NOMMU_INITIAL_TRIM_EXCESS
int "Turn on mmap() excess space trimming before booting"
depends on !MMU
default 1
help
The NOMMU mmap() frequently needs to allocate large contiguous chunks
of memory on which to store mappings, but it can only ask the system
allocator for chunks in 2^N*PAGE_SIZE amounts - which is frequently
more than it requires. To deal with this, mmap() is able to trim off
the excess and return it to the allocator.
If trimming is enabled, the excess is trimmed off and returned to the
system allocator, which can cause extra fragmentation, particularly
if there are a lot of transient processes.
If trimming is disabled, the excess is kept, but not used, which for
long-term mappings means that the space is wasted.
Trimming can be dynamically controlled through a sysctl option
(/proc/sys/vm/nr_trim_pages) which specifies the minimum number of
excess pages there must be before trimming should occur, or zero if
no trimming is to occur.
This option specifies the initial value of this option. The default
of 1 says that all excess pages should be trimmed.
See Documentation/admin-guide/mm/nommu-mmap.rst for more information.
config ARCH_WANT_GENERAL_HUGETLB
bool
config ARCH_WANTS_THP_SWAP
def_bool n
menuconfig TRANSPARENT_HUGEPAGE
bool "Transparent Hugepage Support"
depends on HAVE_ARCH_TRANSPARENT_HUGEPAGE && !PREEMPT_RT
select COMPACTION
select XARRAY_MULTI
help
Transparent Hugepages allows the kernel to use huge pages and
huge tlb transparently to the applications whenever possible.
This feature can improve computing performance to certain
applications by speeding up page faults during memory
allocation, by reducing the number of tlb misses and by speeding
up the pagetable walking.
If memory constrained on embedded, you may want to say N.
if TRANSPARENT_HUGEPAGE
choice
prompt "Transparent Hugepage Support sysfs defaults"
depends on TRANSPARENT_HUGEPAGE
default TRANSPARENT_HUGEPAGE_ALWAYS
help
Selects the sysfs defaults for Transparent Hugepage Support.
config TRANSPARENT_HUGEPAGE_ALWAYS
bool "always"
help
Enabling Transparent Hugepage always, can increase the
memory footprint of applications without a guaranteed
benefit but it will work automatically for all applications.
config TRANSPARENT_HUGEPAGE_MADVISE
bool "madvise"
help
Enabling Transparent Hugepage madvise, will only provide a
performance improvement benefit to the applications using
madvise(MADV_HUGEPAGE) but it won't risk to increase the
memory footprint of applications without a guaranteed
benefit.
mm/thp: add CONFIG_TRANSPARENT_HUGEPAGE_NEVER option Currently enabling THP support (CONFIG_TRANSPARENT_HUGEPAGE) requires enabling either CONFIG_TRANSPARENT_HUGEPAGE_ALWAYS or CONFIG_TRANSPARENT_HUGEPAGE_MADVISE, which both cause khugepaged starting by default at kernel bootup. Add the third choice CONFIG_TRANSPARENT_HUGEPAGE_NEVER, in line with the existing kernel command line setting transparent_hugepage=never, to disable THP by default (in particular, to prevent starting khugepaged by default) but still allow enabling it at runtime via sysfs. Rationale: khugepaged has its own non-negligible memory cost even if it is not used by any applications, since it bumps up vm.min_free_kbytes to its own required minimum in set_recommended_min_free_kbytes(). For example, on a machine with 4GB RAM, with 3 mm zones and pageblock_order == MAX_ORDER, starting khugepaged causes vm.min_free_kbytes increase from 8MB to 132MB. So if we use THP on machines with e.g. >=8GB of memory for better performance, but avoid using it on lower-memory machines to avoid its memory overhead, then for the same reason we also want to avoid even starting khugepaged on those <8GB machines. So with CONFIG_TRANSPARENT_HUGEPAGE_NEVER we can use the same kernel image on both >=8GB and <8GB machines, with THP support enabled but khugepaged not started by default. The userspace can then decide to enable THP via sysfs if needed, based on the total amount of memory. This could also be achieved with the existing transparent_hugepage=never setting in the kernel command line instead. But it seems cleaner to avoid tweaking the command line for such a basic setting. P.S. I see that CONFIG_TRANSPARENT_HUGEPAGE_NEVER was already proposed in the past [1] but without an explanation of the purpose. [1] https://lore.kernel.org/all/202211301651462590168@zte.com.cn/ Link: https://lkml.kernel.org/r/20231205170244.2746210-1-dmaluka@chromium.org Link: https://lore.kernel.org/all/20231204163254.2636289-1-dmaluka@chromium.org/ Signed-off-by: Dmytro Maluka <dmaluka@chromium.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2023-12-05 10:02:44 -07:00
config TRANSPARENT_HUGEPAGE_NEVER
bool "never"
help
Disable Transparent Hugepage by default. It can still be
enabled at runtime via sysfs.
endchoice
mm, THP, swap: delay splitting THP during swap out Patch series "THP swap: Delay splitting THP during swapping out", v11. This patchset is to optimize the performance of Transparent Huge Page (THP) swap. Recently, the performance of the storage devices improved so fast that we cannot saturate the disk bandwidth with single logical CPU when do page swap out even on a high-end server machine. Because the performance of the storage device improved faster than that of single logical CPU. And it seems that the trend will not change in the near future. On the other hand, the THP becomes more and more popular because of increased memory size. So it becomes necessary to optimize THP swap performance. The advantages of the THP swap support include: - Batch the swap operations for the THP to reduce lock acquiring/releasing, including allocating/freeing the swap space, adding/deleting to/from the swap cache, and writing/reading the swap space, etc. This will help improve the performance of the THP swap. - The THP swap space read/write will be 2M sequential IO. It is particularly helpful for the swap read, which are usually 4k random IO. This will improve the performance of the THP swap too. - It will help the memory fragmentation, especially when the THP is heavily used by the applications. The 2M continuous pages will be free up after THP swapping out. - It will improve the THP utilization on the system with the swap turned on. Because the speed for khugepaged to collapse the normal pages into the THP is quite slow. After the THP is split during the swapping out, it will take quite long time for the normal pages to collapse back into the THP after being swapped in. The high THP utilization helps the efficiency of the page based memory management too. There are some concerns regarding THP swap in, mainly because possible enlarged read/write IO size (for swap in/out) may put more overhead on the storage device. To deal with that, the THP swap in should be turned on only when necessary. For example, it can be selected via "always/never/madvise" logic, to be turned on globally, turned off globally, or turned on only for VMA with MADV_HUGEPAGE, etc. This patchset is the first step for the THP swap support. The plan is to delay splitting THP step by step, finally avoid splitting THP during the THP swapping out and swap out/in the THP as a whole. As the first step, in this patchset, the splitting huge page is delayed from almost the first step of swapping out to after allocating the swap space for the THP and adding the THP into the swap cache. This will reduce lock acquiring/releasing for the locks used for the swap cache management. With the patchset, the swap out throughput improves 15.5% (from about 3.73GB/s to about 4.31GB/s) in the vm-scalability swap-w-seq test case with 8 processes. The test is done on a Xeon E5 v3 system. The swap device used is a RAM simulated PMEM (persistent memory) device. To test the sequential swapping out, the test case creates 8 processes, which sequentially allocate and write to the anonymous pages until the RAM and part of the swap device is used up. This patch (of 5): In this patch, splitting huge page is delayed from almost the first step of swapping out to after allocating the swap space for the THP (Transparent Huge Page) and adding the THP into the swap cache. This will batch the corresponding operation, thus improve THP swap out throughput. This is the first step for the THP swap optimization. The plan is to delay splitting the THP step by step and avoid splitting the THP finally. In this patch, one swap cluster is used to hold the contents of each THP swapped out. So, the size of the swap cluster is changed to that of the THP (Transparent Huge Page) on x86_64 architecture (512). For other architectures which want such THP swap optimization, ARCH_USES_THP_SWAP_CLUSTER needs to be selected in the Kconfig file for the architecture. In effect, this will enlarge swap cluster size by 2 times on x86_64. Which may make it harder to find a free cluster when the swap space becomes fragmented. So that, this may reduce the continuous swap space allocation and sequential write in theory. The performance test in 0day shows no regressions caused by this. In the future of THP swap optimization, some information of the swapped out THP (such as compound map count) will be recorded in the swap_cluster_info data structure. The mem cgroup swap accounting functions are enhanced to support charge or uncharge a swap cluster backing a THP as a whole. The swap cluster allocate/free functions are added to allocate/free a swap cluster for a THP. A fair simple algorithm is used for swap cluster allocation, that is, only the first swap device in priority list will be tried to allocate the swap cluster. The function will fail if the trying is not successful, and the caller will fallback to allocate a single swap slot instead. This works good enough for normal cases. If the difference of the number of the free swap clusters among multiple swap devices is significant, it is possible that some THPs are split earlier than necessary. For example, this could be caused by big size difference among multiple swap devices. The swap cache functions is enhanced to support add/delete THP to/from the swap cache as a set of (HPAGE_PMD_NR) sub-pages. This may be enhanced in the future with multi-order radix tree. But because we will split the THP soon during swapping out, that optimization doesn't make much sense for this first step. The THP splitting functions are enhanced to support to split THP in swap cache during swapping out. The page lock will be held during allocating the swap cluster, adding the THP into the swap cache and splitting the THP. So in the code path other than swapping out, if the THP need to be split, the PageSwapCache(THP) will be always false. The swap cluster is only available for SSD, so the THP swap optimization in this patchset has no effect for HDD. [ying.huang@intel.com: fix two issues in THP optimize patch] Link: http://lkml.kernel.org/r/87k25ed8zo.fsf@yhuang-dev.intel.com [hannes@cmpxchg.org: extensive cleanups and simplifications, reduce code size] Link: http://lkml.kernel.org/r/20170515112522.32457-2-ying.huang@intel.com Signed-off-by: "Huang, Ying" <ying.huang@intel.com> Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Suggested-by: Andrew Morton <akpm@linux-foundation.org> [for config option] Acked-by: Kirill A. Shutemov <kirill.shutemov@linux.intel.com> [for changes in huge_memory.c and huge_mm.h] Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Ebru Akagunduz <ebru.akagunduz@gmail.com> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Michal Hocko <mhocko@kernel.org> Cc: Tejun Heo <tj@kernel.org> Cc: Hugh Dickins <hughd@google.com> Cc: Shaohua Li <shli@kernel.org> Cc: Minchan Kim <minchan@kernel.org> Cc: Rik van Riel <riel@redhat.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2017-07-06 15:37:18 -07:00
config THP_SWAP
def_bool y
mm,hugetlb: use folio fields in second tail page Patch series "mm,huge,rmap: unify and speed up compound mapcounts". This patch (of 3): We want to declare one more int in the first tail of a compound page: that first tail page being valuable property, since every compound page has a first tail, but perhaps no more than that. No problem on 64-bit: there is already space for it. No problem with 32-bit THPs: 5.18 commit 5232c63f46fd ("mm: Make compound_pincount always available") kindly cleared the space for it, apparently not realizing that only 64-bit architectures enable CONFIG_THP_SWAP (whose use of tail page->private might conflict) - but make sure of that in its Kconfig. But hugetlb pages use tail page->private of the first tail page for a subpool pointer, which will conflict; and they also use page->private of the 2nd, 3rd and 4th tails. Undo "mm: add private field of first tail to struct page and struct folio"'s recent addition of private_1 to the folio tail: instead add hugetlb_subpool, hugetlb_cgroup, hugetlb_cgroup_rsvd, hugetlb_hwpoison to a second tail page of the folio: THP has long been using several fields of that tail, so make better use of it for hugetlb too. This is not how a generic folio should be declared in future, but it is an effective transitional way to make use of it. Delete the SUBPAGE_INDEX stuff, but keep __NR_USED_SUBPAGE: now 3. [hughd@google.com: prefix folio's page_1 and page_2 with double underscore, give folio's _flags_2 and _head_2 a line documentation each] Link: https://lkml.kernel.org/r/9e2cb6b-5b58-d3f2-b5ee-5f8a14e8f10@google.com Link: https://lkml.kernel.org/r/5f52de70-975-e94f-f141-543765736181@google.com Link: https://lkml.kernel.org/r/3818cc9a-9999-d064-d778-9c94c5911e6@google.com Signed-off-by: Hugh Dickins <hughd@google.com> Acked-by: Kirill A. Shutemov <kirill.shutemov@linux.intel.com> Cc: David Hildenbrand <david@redhat.com> Cc: James Houghton <jthoughton@google.com> Cc: John Hubbard <jhubbard@nvidia.com> Cc: Matthew Wilcox (Oracle) <willy@infradead.org> Cc: Miaohe Lin <linmiaohe@huawei.com> Cc: Mike Kravetz <mike.kravetz@oracle.com> Cc: Mina Almasry <almasrymina@google.com> Cc: Muchun Song <songmuchun@bytedance.com> Cc: Naoya Horiguchi <naoya.horiguchi@linux.dev> Cc: Peter Xu <peterx@redhat.com> Cc: Sidhartha Kumar <sidhartha.kumar@oracle.com> Cc: Vlastimil Babka <vbabka@suse.cz> Cc: Yang Shi <shy828301@gmail.com> Cc: Zach O'Keefe <zokeefe@google.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2022-11-02 18:48:45 -07:00
depends on TRANSPARENT_HUGEPAGE && ARCH_WANTS_THP_SWAP && SWAP && 64BIT
mm, THP, swap: delay splitting THP during swap out Patch series "THP swap: Delay splitting THP during swapping out", v11. This patchset is to optimize the performance of Transparent Huge Page (THP) swap. Recently, the performance of the storage devices improved so fast that we cannot saturate the disk bandwidth with single logical CPU when do page swap out even on a high-end server machine. Because the performance of the storage device improved faster than that of single logical CPU. And it seems that the trend will not change in the near future. On the other hand, the THP becomes more and more popular because of increased memory size. So it becomes necessary to optimize THP swap performance. The advantages of the THP swap support include: - Batch the swap operations for the THP to reduce lock acquiring/releasing, including allocating/freeing the swap space, adding/deleting to/from the swap cache, and writing/reading the swap space, etc. This will help improve the performance of the THP swap. - The THP swap space read/write will be 2M sequential IO. It is particularly helpful for the swap read, which are usually 4k random IO. This will improve the performance of the THP swap too. - It will help the memory fragmentation, especially when the THP is heavily used by the applications. The 2M continuous pages will be free up after THP swapping out. - It will improve the THP utilization on the system with the swap turned on. Because the speed for khugepaged to collapse the normal pages into the THP is quite slow. After the THP is split during the swapping out, it will take quite long time for the normal pages to collapse back into the THP after being swapped in. The high THP utilization helps the efficiency of the page based memory management too. There are some concerns regarding THP swap in, mainly because possible enlarged read/write IO size (for swap in/out) may put more overhead on the storage device. To deal with that, the THP swap in should be turned on only when necessary. For example, it can be selected via "always/never/madvise" logic, to be turned on globally, turned off globally, or turned on only for VMA with MADV_HUGEPAGE, etc. This patchset is the first step for the THP swap support. The plan is to delay splitting THP step by step, finally avoid splitting THP during the THP swapping out and swap out/in the THP as a whole. As the first step, in this patchset, the splitting huge page is delayed from almost the first step of swapping out to after allocating the swap space for the THP and adding the THP into the swap cache. This will reduce lock acquiring/releasing for the locks used for the swap cache management. With the patchset, the swap out throughput improves 15.5% (from about 3.73GB/s to about 4.31GB/s) in the vm-scalability swap-w-seq test case with 8 processes. The test is done on a Xeon E5 v3 system. The swap device used is a RAM simulated PMEM (persistent memory) device. To test the sequential swapping out, the test case creates 8 processes, which sequentially allocate and write to the anonymous pages until the RAM and part of the swap device is used up. This patch (of 5): In this patch, splitting huge page is delayed from almost the first step of swapping out to after allocating the swap space for the THP (Transparent Huge Page) and adding the THP into the swap cache. This will batch the corresponding operation, thus improve THP swap out throughput. This is the first step for the THP swap optimization. The plan is to delay splitting the THP step by step and avoid splitting the THP finally. In this patch, one swap cluster is used to hold the contents of each THP swapped out. So, the size of the swap cluster is changed to that of the THP (Transparent Huge Page) on x86_64 architecture (512). For other architectures which want such THP swap optimization, ARCH_USES_THP_SWAP_CLUSTER needs to be selected in the Kconfig file for the architecture. In effect, this will enlarge swap cluster size by 2 times on x86_64. Which may make it harder to find a free cluster when the swap space becomes fragmented. So that, this may reduce the continuous swap space allocation and sequential write in theory. The performance test in 0day shows no regressions caused by this. In the future of THP swap optimization, some information of the swapped out THP (such as compound map count) will be recorded in the swap_cluster_info data structure. The mem cgroup swap accounting functions are enhanced to support charge or uncharge a swap cluster backing a THP as a whole. The swap cluster allocate/free functions are added to allocate/free a swap cluster for a THP. A fair simple algorithm is used for swap cluster allocation, that is, only the first swap device in priority list will be tried to allocate the swap cluster. The function will fail if the trying is not successful, and the caller will fallback to allocate a single swap slot instead. This works good enough for normal cases. If the difference of the number of the free swap clusters among multiple swap devices is significant, it is possible that some THPs are split earlier than necessary. For example, this could be caused by big size difference among multiple swap devices. The swap cache functions is enhanced to support add/delete THP to/from the swap cache as a set of (HPAGE_PMD_NR) sub-pages. This may be enhanced in the future with multi-order radix tree. But because we will split the THP soon during swapping out, that optimization doesn't make much sense for this first step. The THP splitting functions are enhanced to support to split THP in swap cache during swapping out. The page lock will be held during allocating the swap cluster, adding the THP into the swap cache and splitting the THP. So in the code path other than swapping out, if the THP need to be split, the PageSwapCache(THP) will be always false. The swap cluster is only available for SSD, so the THP swap optimization in this patchset has no effect for HDD. [ying.huang@intel.com: fix two issues in THP optimize patch] Link: http://lkml.kernel.org/r/87k25ed8zo.fsf@yhuang-dev.intel.com [hannes@cmpxchg.org: extensive cleanups and simplifications, reduce code size] Link: http://lkml.kernel.org/r/20170515112522.32457-2-ying.huang@intel.com Signed-off-by: "Huang, Ying" <ying.huang@intel.com> Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Suggested-by: Andrew Morton <akpm@linux-foundation.org> [for config option] Acked-by: Kirill A. Shutemov <kirill.shutemov@linux.intel.com> [for changes in huge_memory.c and huge_mm.h] Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Ebru Akagunduz <ebru.akagunduz@gmail.com> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Michal Hocko <mhocko@kernel.org> Cc: Tejun Heo <tj@kernel.org> Cc: Hugh Dickins <hughd@google.com> Cc: Shaohua Li <shli@kernel.org> Cc: Minchan Kim <minchan@kernel.org> Cc: Rik van Riel <riel@redhat.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2017-07-06 15:37:18 -07:00
help
Swap transparent huge pages in one piece, without splitting.
XXX: For now, swap cluster backing transparent huge page
will be split after swapout.
mm, THP, swap: delay splitting THP during swap out Patch series "THP swap: Delay splitting THP during swapping out", v11. This patchset is to optimize the performance of Transparent Huge Page (THP) swap. Recently, the performance of the storage devices improved so fast that we cannot saturate the disk bandwidth with single logical CPU when do page swap out even on a high-end server machine. Because the performance of the storage device improved faster than that of single logical CPU. And it seems that the trend will not change in the near future. On the other hand, the THP becomes more and more popular because of increased memory size. So it becomes necessary to optimize THP swap performance. The advantages of the THP swap support include: - Batch the swap operations for the THP to reduce lock acquiring/releasing, including allocating/freeing the swap space, adding/deleting to/from the swap cache, and writing/reading the swap space, etc. This will help improve the performance of the THP swap. - The THP swap space read/write will be 2M sequential IO. It is particularly helpful for the swap read, which are usually 4k random IO. This will improve the performance of the THP swap too. - It will help the memory fragmentation, especially when the THP is heavily used by the applications. The 2M continuous pages will be free up after THP swapping out. - It will improve the THP utilization on the system with the swap turned on. Because the speed for khugepaged to collapse the normal pages into the THP is quite slow. After the THP is split during the swapping out, it will take quite long time for the normal pages to collapse back into the THP after being swapped in. The high THP utilization helps the efficiency of the page based memory management too. There are some concerns regarding THP swap in, mainly because possible enlarged read/write IO size (for swap in/out) may put more overhead on the storage device. To deal with that, the THP swap in should be turned on only when necessary. For example, it can be selected via "always/never/madvise" logic, to be turned on globally, turned off globally, or turned on only for VMA with MADV_HUGEPAGE, etc. This patchset is the first step for the THP swap support. The plan is to delay splitting THP step by step, finally avoid splitting THP during the THP swapping out and swap out/in the THP as a whole. As the first step, in this patchset, the splitting huge page is delayed from almost the first step of swapping out to after allocating the swap space for the THP and adding the THP into the swap cache. This will reduce lock acquiring/releasing for the locks used for the swap cache management. With the patchset, the swap out throughput improves 15.5% (from about 3.73GB/s to about 4.31GB/s) in the vm-scalability swap-w-seq test case with 8 processes. The test is done on a Xeon E5 v3 system. The swap device used is a RAM simulated PMEM (persistent memory) device. To test the sequential swapping out, the test case creates 8 processes, which sequentially allocate and write to the anonymous pages until the RAM and part of the swap device is used up. This patch (of 5): In this patch, splitting huge page is delayed from almost the first step of swapping out to after allocating the swap space for the THP (Transparent Huge Page) and adding the THP into the swap cache. This will batch the corresponding operation, thus improve THP swap out throughput. This is the first step for the THP swap optimization. The plan is to delay splitting the THP step by step and avoid splitting the THP finally. In this patch, one swap cluster is used to hold the contents of each THP swapped out. So, the size of the swap cluster is changed to that of the THP (Transparent Huge Page) on x86_64 architecture (512). For other architectures which want such THP swap optimization, ARCH_USES_THP_SWAP_CLUSTER needs to be selected in the Kconfig file for the architecture. In effect, this will enlarge swap cluster size by 2 times on x86_64. Which may make it harder to find a free cluster when the swap space becomes fragmented. So that, this may reduce the continuous swap space allocation and sequential write in theory. The performance test in 0day shows no regressions caused by this. In the future of THP swap optimization, some information of the swapped out THP (such as compound map count) will be recorded in the swap_cluster_info data structure. The mem cgroup swap accounting functions are enhanced to support charge or uncharge a swap cluster backing a THP as a whole. The swap cluster allocate/free functions are added to allocate/free a swap cluster for a THP. A fair simple algorithm is used for swap cluster allocation, that is, only the first swap device in priority list will be tried to allocate the swap cluster. The function will fail if the trying is not successful, and the caller will fallback to allocate a single swap slot instead. This works good enough for normal cases. If the difference of the number of the free swap clusters among multiple swap devices is significant, it is possible that some THPs are split earlier than necessary. For example, this could be caused by big size difference among multiple swap devices. The swap cache functions is enhanced to support add/delete THP to/from the swap cache as a set of (HPAGE_PMD_NR) sub-pages. This may be enhanced in the future with multi-order radix tree. But because we will split the THP soon during swapping out, that optimization doesn't make much sense for this first step. The THP splitting functions are enhanced to support to split THP in swap cache during swapping out. The page lock will be held during allocating the swap cluster, adding the THP into the swap cache and splitting the THP. So in the code path other than swapping out, if the THP need to be split, the PageSwapCache(THP) will be always false. The swap cluster is only available for SSD, so the THP swap optimization in this patchset has no effect for HDD. [ying.huang@intel.com: fix two issues in THP optimize patch] Link: http://lkml.kernel.org/r/87k25ed8zo.fsf@yhuang-dev.intel.com [hannes@cmpxchg.org: extensive cleanups and simplifications, reduce code size] Link: http://lkml.kernel.org/r/20170515112522.32457-2-ying.huang@intel.com Signed-off-by: "Huang, Ying" <ying.huang@intel.com> Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Suggested-by: Andrew Morton <akpm@linux-foundation.org> [for config option] Acked-by: Kirill A. Shutemov <kirill.shutemov@linux.intel.com> [for changes in huge_memory.c and huge_mm.h] Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Ebru Akagunduz <ebru.akagunduz@gmail.com> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Michal Hocko <mhocko@kernel.org> Cc: Tejun Heo <tj@kernel.org> Cc: Hugh Dickins <hughd@google.com> Cc: Shaohua Li <shli@kernel.org> Cc: Minchan Kim <minchan@kernel.org> Cc: Rik van Riel <riel@redhat.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2017-07-06 15:37:18 -07:00
For selection by architectures with reasonable THP sizes.
config READ_ONLY_THP_FOR_FS
bool "Read-only THP for filesystems (EXPERIMENTAL)"
depends on TRANSPARENT_HUGEPAGE && SHMEM
help
Allow khugepaged to put read-only file-backed pages in THP.
This is marked experimental because it is a new feature. Write
support of file THPs will be developed in the next few release
cycles.
endif # TRANSPARENT_HUGEPAGE
mm/Kconfig: CONFIG_PGTABLE_HAS_HUGE_LEAVES Patch series "mm/gup: Unify hugetlb, part 2", v4. The series removes the hugetlb slow gup path after a previous refactor work [1], so that slow gup now uses the exact same path to process all kinds of memory including hugetlb. For the long term, we may want to remove most, if not all, call sites of huge_pte_offset(). It'll be ideal if that API can be completely dropped from arch hugetlb API. This series is one small step towards merging hugetlb specific codes into generic mm paths. From that POV, this series removes one reference to huge_pte_offset() out of many others. One goal of such a route is that we can reconsider merging hugetlb features like High Granularity Mapping (HGM). It was not accepted in the past because it may add lots of hugetlb specific codes and make the mm code even harder to maintain. With a merged codeset, features like HGM can hopefully share some code with THP, legacy (PMD+) or modern (continuous PTEs). To make it work, the generic slow gup code will need to at least understand hugepd, which is already done like so in fast-gup. Due to the specialty of hugepd to be software-only solution (no hardware recognizes the hugepd format, so it's purely artificial structures), there's chance we can merge some or all hugepd formats with cont_pte in the future. That question is yet unsettled from Power side to have an acknowledgement. As of now for this series, I kept the hugepd handling because we may still need to do so before getting a clearer picture of the future of hugepd. The other reason is simply that we did it already for fast-gup and most codes are still around to be reused. It'll make more sense to keep slow/fast gup behave the same before a decision is made to remove hugepd. There's one major difference for slow-gup on cont_pte / cont_pmd handling, currently supported on three architectures (aarch64, riscv, ppc). Before the series, slow gup will be able to recognize e.g. cont_pte entries with the help of huge_pte_offset() when hstate is around. Now it's gone but still working, by looking up pgtable entries one by one. It's not ideal, but hopefully this change should not affect yet on major workloads. There's some more information in the commit message of the last patch. If this would be a concern, we can consider teaching slow gup to recognize cont pte/pmd entries, and that should recover the lost performance. But I doubt its necessity for now, so I kept it as simple as it can be. Patch layout ============= Patch 1-8: Preparation works, or cleanups in relevant code paths Patch 9-11: Teach slow gup with all kinds of huge entries (pXd, hugepd) Patch 12: Drop hugetlb_follow_page_mask() More information can be found in the commit messages of each patch. [1] https://lore.kernel.org/all/20230628215310.73782-1-peterx@redhat.com [2] https://lore.kernel.org/r/20240321215047.678172-1-peterx@redhat.com Introduce a config option that will be selected as long as huge leaves are involved in pgtable (thp or hugetlbfs). It would be useful to mark any code with this new config that can process either hugetlb or thp pages in any level that is higher than pte level. Link: https://lkml.kernel.org/r/20240327152332.950956-1-peterx@redhat.com Link: https://lkml.kernel.org/r/20240327152332.950956-2-peterx@redhat.com Signed-off-by: Peter Xu <peterx@redhat.com> Reviewed-by: Jason Gunthorpe <jgg@nvidia.com> Tested-by: Ryan Roberts <ryan.roberts@arm.com> Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Andrew Jones <andrew.jones@linux.dev> Cc: Aneesh Kumar K.V (IBM) <aneesh.kumar@kernel.org> Cc: Axel Rasmussen <axelrasmussen@google.com> Cc: Christophe Leroy <christophe.leroy@csgroup.eu> Cc: Christoph Hellwig <hch@infradead.org> Cc: David Hildenbrand <david@redhat.com> Cc: James Houghton <jthoughton@google.com> Cc: John Hubbard <jhubbard@nvidia.com> Cc: Kirill A. Shutemov <kirill@shutemov.name> Cc: Lorenzo Stoakes <lstoakes@gmail.com> Cc: Matthew Wilcox (Oracle) <willy@infradead.org> Cc: Michael Ellerman <mpe@ellerman.id.au> Cc: "Mike Rapoport (IBM)" <rppt@kernel.org> Cc: Muchun Song <muchun.song@linux.dev> Cc: Rik van Riel <riel@surriel.com> Cc: Vlastimil Babka <vbabka@suse.cz> Cc: Yang Shi <shy828301@gmail.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2024-03-27 08:23:20 -07:00
#
# The architecture supports pgtable leaves that is larger than PAGE_SIZE
#
config PGTABLE_HAS_HUGE_LEAVES
def_bool TRANSPARENT_HUGEPAGE || HUGETLB_PAGE
mm: introduce ARCH_SUPPORTS_HUGE_PFNMAP and special bits to pmd/pud Patch series "mm: Support huge pfnmaps", v2. Overview ======== This series implements huge pfnmaps support for mm in general. Huge pfnmap allows e.g. VM_PFNMAP vmas to map in either PMD or PUD levels, similar to what we do with dax / thp / hugetlb so far to benefit from TLB hits. Now we extend that idea to PFN mappings, e.g. PCI MMIO bars where it can grow as large as 8GB or even bigger. Currently, only x86_64 (1G+2M) and arm64 (2M) are supported. The last patch (from Alex Williamson) will be the first user of huge pfnmap, so as to enable vfio-pci driver to fault in huge pfn mappings. Implementation ============== In reality, it's relatively simple to add such support comparing to many other types of mappings, because of PFNMAP's specialties when there's no vmemmap backing it, so that most of the kernel routines on huge mappings should simply already fail for them, like GUPs or old-school follow_page() (which is recently rewritten to be folio_walk* APIs by David). One trick here is that we're still unmature on PUDs in generic paths here and there, as DAX is so far the only user. This patchset will add the 2nd user of it. Hugetlb can be a 3rd user if the hugetlb unification work can go on smoothly, but to be discussed later. The other trick is how to allow gup-fast working for such huge mappings even if there's no direct sign of knowing whether it's a normal page or MMIO mapping. This series chose to keep the pte_special solution, so that it reuses similar idea on setting a special bit to pfnmap PMDs/PUDs so that gup-fast will be able to identify them and fail properly. Along the way, we'll also notice that the major pgtable pfn walker, aka, follow_pte(), will need to retire soon due to the fact that it only works with ptes. A new set of simple API is introduced (follow_pfnmap* API) to be able to do whatever follow_pte() can already do, plus that it can also process huge pfnmaps now. Half of this series is about that and converting all existing pfnmap walkers to use the new API properly. Hopefully the new API also looks better to avoid exposing e.g. pgtable lock details into the callers, so that it can be used in an even more straightforward way. Here, three more options will be introduced and involved in huge pfnmap: - ARCH_SUPPORTS_HUGE_PFNMAP Arch developers will need to select this option when huge pfnmap is supported in arch's Kconfig. After this patchset applied, both x86_64 and arm64 will start to enable it by default. - ARCH_SUPPORTS_PMD_PFNMAP / ARCH_SUPPORTS_PUD_PFNMAP These options are for driver developers to identify whether current arch / config supports huge pfnmaps, making decision on whether it can use the huge pfnmap APIs to inject them. One can refer to the last vfio-pci patch from Alex on the use of them properly in a device driver. So after the whole set applied, and if one would enable some dynamic debug lines in vfio-pci core files, we should observe things like: vfio-pci 0000:00:06.0: vfio_pci_mmap_huge_fault(,order = 9) BAR 0 page offset 0x0: 0x100 vfio-pci 0000:00:06.0: vfio_pci_mmap_huge_fault(,order = 9) BAR 0 page offset 0x200: 0x100 vfio-pci 0000:00:06.0: vfio_pci_mmap_huge_fault(,order = 9) BAR 0 page offset 0x400: 0x100 In this specific case, it says that vfio-pci faults in PMDs properly for a few BAR0 offsets. Patch Layout ============ Patch 1: Introduce the new options mentioned above for huge PFNMAPs Patch 2: A tiny cleanup Patch 3-8: Preparation patches for huge pfnmap (include introduce special bit for pmd/pud) Patch 9-16: Introduce follow_pfnmap*() API, use it everywhere, and then drop follow_pte() API Patch 17: Add huge pfnmap support for x86_64 Patch 18: Add huge pfnmap support for arm64 Patch 19: Add vfio-pci support for all kinds of huge pfnmaps (Alex) TODO ==== More architectures / More page sizes ------------------------------------ Currently only x86_64 (2M+1G) and arm64 (2M) are supported. There seems to have plan to support arm64 1G later on top of this series [2]. Any arch will need to first support THP / THP_1G, then provide a special bit in pmds/puds to support huge pfnmaps. remap_pfn_range() support ------------------------- Currently, remap_pfn_range() still only maps PTEs. With the new option, remap_pfn_range() can logically start to inject either PMDs or PUDs when the alignment requirements match on the VAs. When the support is there, it should be able to silently benefit all drivers that is using remap_pfn_range() in its mmap() handler on better TLB hit rate and overall faster MMIO accesses similar to processor on hugepages. More driver support ------------------- VFIO is so far the only consumer for the huge pfnmaps after this series applied. Besides above remap_pfn_range() generic optimization, device driver can also try to optimize its mmap() on a better VA alignment for either PMD/PUD sizes. This may, iiuc, normally require userspace changes, as the driver doesn't normally decide the VA to map a bar. But I don't think I know all the drivers to know the full picture. Credits all go to Alex on help testing the GPU/NIC use cases above. [0] https://lore.kernel.org/r/73ad9540-3fb8-4154-9a4f-30a0a2b03d41@lucifer.local [1] https://lore.kernel.org/r/20240807194812.819412-1-peterx@redhat.com [2] https://lore.kernel.org/r/498e0731-81a4-4f75-95b4-a8ad0bcc7665@huawei.com This patch (of 19): This patch introduces the option to introduce special pte bit into pmd/puds. Archs can start to define pmd_special / pud_special when supported by selecting the new option. Per-arch support will be added later. Before that, create fallbacks for these helpers so that they are always available. Link: https://lkml.kernel.org/r/20240826204353.2228736-1-peterx@redhat.com Link: https://lkml.kernel.org/r/20240826204353.2228736-2-peterx@redhat.com Signed-off-by: Peter Xu <peterx@redhat.com> Cc: Alexander Gordeev <agordeev@linux.ibm.com> Cc: Alex Williamson <alex.williamson@redhat.com> Cc: Aneesh Kumar K.V <aneesh.kumar@linux.ibm.com> Cc: Borislav Petkov <bp@alien8.de> Cc: Catalin Marinas <catalin.marinas@arm.com> Cc: Christian Borntraeger <borntraeger@linux.ibm.com> Cc: Dave Hansen <dave.hansen@linux.intel.com> Cc: David Hildenbrand <david@redhat.com> Cc: Gavin Shan <gshan@redhat.com> Cc: Gerald Schaefer <gerald.schaefer@linux.ibm.com> Cc: Heiko Carstens <hca@linux.ibm.com> Cc: Ingo Molnar <mingo@redhat.com> Cc: Jason Gunthorpe <jgg@nvidia.com> Cc: Matthew Wilcox <willy@infradead.org> Cc: Niklas Schnelle <schnelle@linux.ibm.com> Cc: Paolo Bonzini <pbonzini@redhat.com> Cc: Ryan Roberts <ryan.roberts@arm.com> Cc: Sean Christopherson <seanjc@google.com> Cc: Sven Schnelle <svens@linux.ibm.com> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Vasily Gorbik <gor@linux.ibm.com> Cc: Will Deacon <will@kernel.org> Cc: Zi Yan <ziy@nvidia.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2024-08-26 13:43:35 -07:00
# TODO: Allow to be enabled without THP
config ARCH_SUPPORTS_HUGE_PFNMAP
def_bool n
depends on TRANSPARENT_HUGEPAGE
config ARCH_SUPPORTS_PMD_PFNMAP
def_bool y
depends on ARCH_SUPPORTS_HUGE_PFNMAP && HAVE_ARCH_TRANSPARENT_HUGEPAGE
config ARCH_SUPPORTS_PUD_PFNMAP
def_bool y
depends on ARCH_SUPPORTS_HUGE_PFNMAP && HAVE_ARCH_TRANSPARENT_HUGEPAGE_PUD
#
# UP and nommu archs use km based percpu allocator
#
config NEED_PER_CPU_KM
depends on !SMP || !MMU
bool
default y
mm: cleancache core ops functions and config This third patch of eight in this cleancache series provides the core code for cleancache that interfaces between the hooks in VFS and individual filesystems and a cleancache backend. It also includes build and config patches. Two new files are added: mm/cleancache.c and include/linux/cleancache.h. Note that CONFIG_CLEANCACHE can default to on; in systems that do not provide a cleancache backend, all hooks devolve to a simple check of a global enable flag, so performance impact should be negligible but can be reduced to zero impact if config'ed off. However for this first commit, it defaults to off. Details and a FAQ can be found in Documentation/vm/cleancache.txt Credits: Cleancache_ops design derived from Jeremy Fitzhardinge design for tmem [v8: dan.magenheimer@oracle.com: fix exportfs call affecting btrfs] [v8: akpm@linux-foundation.org: use static inline function, not macro] [v7: dan.magenheimer@oracle.com: cleanup sysfs and remove cleancache prefix] [v6: JBeulich@novell.com: robustly handle buggy fs encode_fh actor definition] [v5: jeremy@goop.org: clean up global usage and static var names] [v5: jeremy@goop.org: simplify init hook and any future fs init changes] [v5: hch@infradead.org: cleaner non-global interface for ops registration] [v4: adilger@sun.com: interface must support exportfs FS's] [v4: hch@infradead.org: interface must support 64-bit FS on 32-bit kernel] [v3: akpm@linux-foundation.org: use one ops struct to avoid pointer hops] [v3: akpm@linux-foundation.org: document and ensure PageLocked reqts are met] [v3: ngupta@vflare.org: fix success/fail codes, change funcs to void] [v2: viro@ZenIV.linux.org.uk: use sane types] Signed-off-by: Dan Magenheimer <dan.magenheimer@oracle.com> Reviewed-by: Jeremy Fitzhardinge <jeremy@goop.org> Reviewed-by: Konrad Rzeszutek Wilk <konrad.wilk@oracle.com> Acked-by: Al Viro <viro@ZenIV.linux.org.uk> Acked-by: Andrew Morton <akpm@linux-foundation.org> Acked-by: Nitin Gupta <ngupta@vflare.org> Acked-by: Minchan Kim <minchan.kim@gmail.com> Acked-by: Andreas Dilger <adilger@sun.com> Acked-by: Jan Beulich <JBeulich@novell.com> Cc: Matthew Wilcox <matthew@wil.cx> Cc: Nick Piggin <npiggin@kernel.dk> Cc: Mel Gorman <mel@csn.ul.ie> Cc: Rik Van Riel <riel@redhat.com> Cc: Chris Mason <chris.mason@oracle.com> Cc: Ted Ts'o <tytso@mit.edu> Cc: Mark Fasheh <mfasheh@suse.com> Cc: Joel Becker <joel.becker@oracle.com>
2011-05-26 09:01:36 -07:00
mm: percpu: generalize percpu related config Patch series "mm: percpu: Cleanup percpu first chunk function". When supporting page mapping percpu first chunk allocator on arm64, we found there are lots of duplicated codes in percpu embed/page first chunk allocator. This patchset is aimed to cleanup them and should no function change. The currently supported status about 'embed' and 'page' in Archs shows below, embed: NEED_PER_CPU_PAGE_FIRST_CHUNK page: NEED_PER_CPU_EMBED_FIRST_CHUNK embed page ------------------------ arm64 Y Y mips Y N powerpc Y Y riscv Y N sparc Y Y x86 Y Y ------------------------ There are two interfaces about percpu first chunk allocator, extern int __init pcpu_embed_first_chunk(size_t reserved_size, size_t dyn_size, size_t atom_size, pcpu_fc_cpu_distance_fn_t cpu_distance_fn, - pcpu_fc_alloc_fn_t alloc_fn, - pcpu_fc_free_fn_t free_fn); + pcpu_fc_cpu_to_node_fn_t cpu_to_nd_fn); extern int __init pcpu_page_first_chunk(size_t reserved_size, - pcpu_fc_alloc_fn_t alloc_fn, - pcpu_fc_free_fn_t free_fn, - pcpu_fc_populate_pte_fn_t populate_pte_fn); + pcpu_fc_cpu_to_node_fn_t cpu_to_nd_fn); The pcpu_fc_alloc_fn_t/pcpu_fc_free_fn_t is killed, we provide generic pcpu_fc_alloc() and pcpu_fc_free() function, which are called in the pcpu_embed/page_first_chunk(). 1) For pcpu_embed_first_chunk(), pcpu_fc_cpu_to_node_fn_t is needed to be provided when archs supported NUMA. 2) For pcpu_page_first_chunk(), the pcpu_fc_populate_pte_fn_t is killed too, a generic pcpu_populate_pte() which marked '__weak' is provided, if you need a different function to populate pte on the arch(like x86), please provide its own implementation. [1] https://github.com/kevin78/linux.git percpu-cleanup This patch (of 4): The HAVE_SETUP_PER_CPU_AREA/NEED_PER_CPU_EMBED_FIRST_CHUNK/ NEED_PER_CPU_PAGE_FIRST_CHUNK/USE_PERCPU_NUMA_NODE_ID configs, which have duplicate definitions on platforms that subscribe it. Move them into mm, drop these redundant definitions and instead just select it on applicable platforms. Link: https://lkml.kernel.org/r/20211216112359.103822-1-wangkefeng.wang@huawei.com Link: https://lkml.kernel.org/r/20211216112359.103822-2-wangkefeng.wang@huawei.com Signed-off-by: Kefeng Wang <wangkefeng.wang@huawei.com> Acked-by: Catalin Marinas <catalin.marinas@arm.com> [arm64] Cc: Will Deacon <will@kernel.org> Cc: Thomas Bogendoerfer <tsbogend@alpha.franken.de> Cc: Michael Ellerman <mpe@ellerman.id.au> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: Paul Mackerras <paulus@samba.org> Cc: Paul Walmsley <paul.walmsley@sifive.com> Cc: Palmer Dabbelt <palmer@dabbelt.com> Cc: Albert Ou <aou@eecs.berkeley.edu> Cc: "David S. Miller" <davem@davemloft.net> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Ingo Molnar <mingo@redhat.com> Cc: Borislav Petkov <bp@alien8.de> Cc: Dave Hansen <dave.hansen@linux.intel.com> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Christoph Lameter <cl@linux.com> Cc: Dennis Zhou <dennis@kernel.org> Cc: Greg Kroah-Hartman <gregkh@linuxfoundation.org> Cc: "Rafael J. Wysocki" <rafael@kernel.org> Cc: Tejun Heo <tj@kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2022-01-19 19:07:41 -07:00
config NEED_PER_CPU_EMBED_FIRST_CHUNK
bool
config NEED_PER_CPU_PAGE_FIRST_CHUNK
bool
config USE_PERCPU_NUMA_NODE_ID
bool
config HAVE_SETUP_PER_CPU_AREA
bool
config CMA
bool "Contiguous Memory Allocator"
mm: remove CONFIG_HAVE_MEMBLOCK All architecures use memblock for early memory management. There is no need for the CONFIG_HAVE_MEMBLOCK configuration option. [rppt@linux.vnet.ibm.com: of/fdt: fixup #ifdefs] Link: http://lkml.kernel.org/r/20180919103457.GA20545@rapoport-lnx [rppt@linux.vnet.ibm.com: csky: fixups after bootmem removal] Link: http://lkml.kernel.org/r/20180926112744.GC4628@rapoport-lnx [rppt@linux.vnet.ibm.com: remove stale #else and the code it protects] Link: http://lkml.kernel.org/r/1538067825-24835-1-git-send-email-rppt@linux.vnet.ibm.com Link: http://lkml.kernel.org/r/1536927045-23536-4-git-send-email-rppt@linux.vnet.ibm.com Signed-off-by: Mike Rapoport <rppt@linux.vnet.ibm.com> Acked-by: Michal Hocko <mhocko@suse.com> Tested-by: Jonathan Cameron <jonathan.cameron@huawei.com> Cc: Catalin Marinas <catalin.marinas@arm.com> Cc: Chris Zankel <chris@zankel.net> Cc: "David S. Miller" <davem@davemloft.net> Cc: Geert Uytterhoeven <geert@linux-m68k.org> Cc: Greentime Hu <green.hu@gmail.com> Cc: Greg Kroah-Hartman <gregkh@linuxfoundation.org> Cc: Guan Xuetao <gxt@pku.edu.cn> Cc: Ingo Molnar <mingo@redhat.com> Cc: "James E.J. Bottomley" <jejb@parisc-linux.org> Cc: Jonas Bonn <jonas@southpole.se> Cc: Jonathan Corbet <corbet@lwn.net> Cc: Ley Foon Tan <lftan@altera.com> Cc: Mark Salter <msalter@redhat.com> Cc: Martin Schwidefsky <schwidefsky@de.ibm.com> Cc: Matt Turner <mattst88@gmail.com> Cc: Michael Ellerman <mpe@ellerman.id.au> Cc: Michal Simek <monstr@monstr.eu> Cc: Palmer Dabbelt <palmer@sifive.com> Cc: Paul Burton <paul.burton@mips.com> Cc: Richard Kuo <rkuo@codeaurora.org> Cc: Richard Weinberger <richard@nod.at> Cc: Rich Felker <dalias@libc.org> Cc: Russell King <linux@armlinux.org.uk> Cc: Serge Semin <fancer.lancer@gmail.com> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Tony Luck <tony.luck@intel.com> Cc: Vineet Gupta <vgupta@synopsys.com> Cc: Yoshinori Sato <ysato@users.sourceforge.jp> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2018-10-30 15:07:44 -07:00
depends on MMU
select MIGRATION
select MEMORY_ISOLATION
help
This enables the Contiguous Memory Allocator which allows other
subsystems to allocate big physically-contiguous blocks of memory.
CMA reserves a region of memory and allows only movable pages to
be allocated from it. This way, the kernel can use the memory for
pagecache and when a subsystem requests for contiguous area, the
allocated pages are migrated away to serve the contiguous request.
If unsure, say "n".
config CMA_DEBUGFS
bool "CMA debugfs interface"
depends on CMA && DEBUG_FS
help
Turns on the DebugFS interface for CMA.
mm: cma: support sysfs Since CMA is getting used more widely, it's more important to keep monitoring CMA statistics for system health since it's directly related to user experience. This patch introduces sysfs statistics for CMA, in order to provide some basic monitoring of the CMA allocator. * the number of CMA page successful allocations * the number of CMA page allocation failures These two values allow the user to calcuate the allocation failure rate for each CMA area. e.g.) /sys/kernel/mm/cma/WIFI/alloc_pages_[success|fail] /sys/kernel/mm/cma/SENSOR/alloc_pages_[success|fail] /sys/kernel/mm/cma/BLUETOOTH/alloc_pages_[success|fail] The cma_stat was intentionally allocated by dynamic allocation to harmonize with kobject lifetime management. https://lore.kernel.org/linux-mm/YCOAmXqt6dZkCQYs@kroah.com/ Link: https://lkml.kernel.org/r/20210324230759.2213957-1-minchan@kernel.org Link: https://lore.kernel.org/linux-mm/20210316100433.17665-1-colin.king@canonical.com/ Signed-off-by: Minchan Kim <minchan@kernel.org> Signed-off-by: Colin Ian King <colin.king@canonical.com> Tested-by: Dmitry Osipenko <digetx@gmail.com> Reviewed-by: Dmitry Osipenko <digetx@gmail.com> Reviewed-by: Greg Kroah-Hartman <gregkh@linuxfoundation.org> Reviewed-by: John Hubbard <jhubbard@nvidia.com> Tested-by: Anders Roxell <anders.roxell@linaro.org> Cc: Suren Baghdasaryan <surenb@google.com> Cc: John Dias <joaodias@google.com> Cc: Matthew Wilcox (Oracle) <willy@infradead.org> Cc: Colin Ian King <colin.king@canonical.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2021-05-04 18:37:28 -07:00
config CMA_SYSFS
bool "CMA information through sysfs interface"
depends on CMA && SYSFS
help
This option exposes some sysfs attributes to get information
from CMA.
CMA: generalize CMA reserved area management functionality Currently, there are two users on CMA functionality, one is the DMA subsystem and the other is the KVM on powerpc. They have their own code to manage CMA reserved area even if they looks really similar. From my guess, it is caused by some needs on bitmap management. KVM side wants to maintain bitmap not for 1 page, but for more size. Eventually it use bitmap where one bit represents 64 pages. When I implement CMA related patches, I should change those two places to apply my change and it seem to be painful to me. I want to change this situation and reduce future code management overhead through this patch. This change could also help developer who want to use CMA in their new feature development, since they can use CMA easily without copying & pasting this reserved area management code. In previous patches, we have prepared some features to generalize CMA reserved area management and now it's time to do it. This patch moves core functions to mm/cma.c and change DMA APIs to use these functions. There is no functional change in DMA APIs. Signed-off-by: Joonsoo Kim <iamjoonsoo.kim@lge.com> Acked-by: Michal Nazarewicz <mina86@mina86.com> Acked-by: Zhang Yanfei <zhangyanfei@cn.fujitsu.com> Acked-by: Minchan Kim <minchan@kernel.org> Reviewed-by: Aneesh Kumar K.V <aneesh.kumar@linux.vnet.ibm.com> Cc: Alexander Graf <agraf@suse.de> Cc: Aneesh Kumar K.V <aneesh.kumar@linux.vnet.ibm.com> Cc: Gleb Natapov <gleb@kernel.org> Acked-by: Marek Szyprowski <m.szyprowski@samsung.com> Tested-by: Marek Szyprowski <m.szyprowski@samsung.com> Cc: Paolo Bonzini <pbonzini@redhat.com> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: Paul Mackerras <paulus@samba.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2014-08-06 16:05:25 -07:00
config CMA_AREAS
int "Maximum count of the CMA areas"
depends on CMA
default 20 if NUMA
default 8
CMA: generalize CMA reserved area management functionality Currently, there are two users on CMA functionality, one is the DMA subsystem and the other is the KVM on powerpc. They have their own code to manage CMA reserved area even if they looks really similar. From my guess, it is caused by some needs on bitmap management. KVM side wants to maintain bitmap not for 1 page, but for more size. Eventually it use bitmap where one bit represents 64 pages. When I implement CMA related patches, I should change those two places to apply my change and it seem to be painful to me. I want to change this situation and reduce future code management overhead through this patch. This change could also help developer who want to use CMA in their new feature development, since they can use CMA easily without copying & pasting this reserved area management code. In previous patches, we have prepared some features to generalize CMA reserved area management and now it's time to do it. This patch moves core functions to mm/cma.c and change DMA APIs to use these functions. There is no functional change in DMA APIs. Signed-off-by: Joonsoo Kim <iamjoonsoo.kim@lge.com> Acked-by: Michal Nazarewicz <mina86@mina86.com> Acked-by: Zhang Yanfei <zhangyanfei@cn.fujitsu.com> Acked-by: Minchan Kim <minchan@kernel.org> Reviewed-by: Aneesh Kumar K.V <aneesh.kumar@linux.vnet.ibm.com> Cc: Alexander Graf <agraf@suse.de> Cc: Aneesh Kumar K.V <aneesh.kumar@linux.vnet.ibm.com> Cc: Gleb Natapov <gleb@kernel.org> Acked-by: Marek Szyprowski <m.szyprowski@samsung.com> Tested-by: Marek Szyprowski <m.szyprowski@samsung.com> Cc: Paolo Bonzini <pbonzini@redhat.com> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: Paul Mackerras <paulus@samba.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2014-08-06 16:05:25 -07:00
help
CMA allows to create CMA areas for particular purpose, mainly,
used as device private area. This parameter sets the maximum
number of CMA area in the system.
If unsure, leave the default value "8" in UMA and "20" in NUMA.
CMA: generalize CMA reserved area management functionality Currently, there are two users on CMA functionality, one is the DMA subsystem and the other is the KVM on powerpc. They have their own code to manage CMA reserved area even if they looks really similar. From my guess, it is caused by some needs on bitmap management. KVM side wants to maintain bitmap not for 1 page, but for more size. Eventually it use bitmap where one bit represents 64 pages. When I implement CMA related patches, I should change those two places to apply my change and it seem to be painful to me. I want to change this situation and reduce future code management overhead through this patch. This change could also help developer who want to use CMA in their new feature development, since they can use CMA easily without copying & pasting this reserved area management code. In previous patches, we have prepared some features to generalize CMA reserved area management and now it's time to do it. This patch moves core functions to mm/cma.c and change DMA APIs to use these functions. There is no functional change in DMA APIs. Signed-off-by: Joonsoo Kim <iamjoonsoo.kim@lge.com> Acked-by: Michal Nazarewicz <mina86@mina86.com> Acked-by: Zhang Yanfei <zhangyanfei@cn.fujitsu.com> Acked-by: Minchan Kim <minchan@kernel.org> Reviewed-by: Aneesh Kumar K.V <aneesh.kumar@linux.vnet.ibm.com> Cc: Alexander Graf <agraf@suse.de> Cc: Aneesh Kumar K.V <aneesh.kumar@linux.vnet.ibm.com> Cc: Gleb Natapov <gleb@kernel.org> Acked-by: Marek Szyprowski <m.szyprowski@samsung.com> Tested-by: Marek Szyprowski <m.szyprowski@samsung.com> Cc: Paolo Bonzini <pbonzini@redhat.com> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: Paul Mackerras <paulus@samba.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2014-08-06 16:05:25 -07:00
config MEM_SOFT_DIRTY
bool "Track memory changes"
depends on CHECKPOINT_RESTORE && HAVE_ARCH_SOFT_DIRTY && PROC_FS
select PROC_PAGE_MONITOR
zbud: add to mm/ zbud is an special purpose allocator for storing compressed pages. It is designed to store up to two compressed pages per physical page. While this design limits storage density, it has simple and deterministic reclaim properties that make it preferable to a higher density approach when reclaim will be used. zbud works by storing compressed pages, or "zpages", together in pairs in a single memory page called a "zbud page". The first buddy is "left justifed" at the beginning of the zbud page, and the last buddy is "right justified" at the end of the zbud page. The benefit is that if either buddy is freed, the freed buddy space, coalesced with whatever slack space that existed between the buddies, results in the largest possible free region within the zbud page. zbud also provides an attractive lower bound on density. The ratio of zpages to zbud pages can not be less than 1. This ensures that zbud can never "do harm" by using more pages to store zpages than the uncompressed zpages would have used on their own. This implementation is a rewrite of the zbud allocator internally used by zcache in the driver/staging tree. The rewrite was necessary to remove some of the zcache specific elements that were ingrained throughout and provide a generic allocation interface that can later be used by zsmalloc and others. This patch adds zbud to mm/ for later use by zswap. Signed-off-by: Seth Jennings <sjenning@linux.vnet.ibm.com> Acked-by: Rik van Riel <riel@redhat.com> Cc: Greg Kroah-Hartman <gregkh@linuxfoundation.org> Cc: Nitin Gupta <ngupta@vflare.org> Cc: Minchan Kim <minchan@kernel.org> Cc: Konrad Rzeszutek Wilk <konrad.wilk@oracle.com> Cc: Dan Magenheimer <dan.magenheimer@oracle.com> Cc: Robert Jennings <rcj@linux.vnet.ibm.com> Cc: Jenifer Hopper <jhopper@us.ibm.com> Cc: Mel Gorman <mgorman@suse.de> Cc: Johannes Weiner <jweiner@redhat.com> Cc: Larry Woodman <lwoodman@redhat.com> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: Dave Hansen <dave@sr71.net> Cc: Joe Perches <joe@perches.com> Cc: Joonsoo Kim <iamjoonsoo.kim@lge.com> Cc: Cody P Schafer <cody@linux.vnet.ibm.com> Cc: Hugh Dickens <hughd@google.com> Cc: Paul Mackerras <paulus@samba.org> Cc: Bob Liu <bob.liu@oracle.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2013-07-10 16:04:55 -07:00
help
This option enables memory changes tracking by introducing a
soft-dirty bit on pte-s. This bit it set when someone writes
into a page just as regular dirty bit, but unlike the latter
it can be cleared by hands.
See Documentation/admin-guide/mm/soft-dirty.rst for more details.
zbud: add to mm/ zbud is an special purpose allocator for storing compressed pages. It is designed to store up to two compressed pages per physical page. While this design limits storage density, it has simple and deterministic reclaim properties that make it preferable to a higher density approach when reclaim will be used. zbud works by storing compressed pages, or "zpages", together in pairs in a single memory page called a "zbud page". The first buddy is "left justifed" at the beginning of the zbud page, and the last buddy is "right justified" at the end of the zbud page. The benefit is that if either buddy is freed, the freed buddy space, coalesced with whatever slack space that existed between the buddies, results in the largest possible free region within the zbud page. zbud also provides an attractive lower bound on density. The ratio of zpages to zbud pages can not be less than 1. This ensures that zbud can never "do harm" by using more pages to store zpages than the uncompressed zpages would have used on their own. This implementation is a rewrite of the zbud allocator internally used by zcache in the driver/staging tree. The rewrite was necessary to remove some of the zcache specific elements that were ingrained throughout and provide a generic allocation interface that can later be used by zsmalloc and others. This patch adds zbud to mm/ for later use by zswap. Signed-off-by: Seth Jennings <sjenning@linux.vnet.ibm.com> Acked-by: Rik van Riel <riel@redhat.com> Cc: Greg Kroah-Hartman <gregkh@linuxfoundation.org> Cc: Nitin Gupta <ngupta@vflare.org> Cc: Minchan Kim <minchan@kernel.org> Cc: Konrad Rzeszutek Wilk <konrad.wilk@oracle.com> Cc: Dan Magenheimer <dan.magenheimer@oracle.com> Cc: Robert Jennings <rcj@linux.vnet.ibm.com> Cc: Jenifer Hopper <jhopper@us.ibm.com> Cc: Mel Gorman <mgorman@suse.de> Cc: Johannes Weiner <jweiner@redhat.com> Cc: Larry Woodman <lwoodman@redhat.com> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: Dave Hansen <dave@sr71.net> Cc: Joe Perches <joe@perches.com> Cc: Joonsoo Kim <iamjoonsoo.kim@lge.com> Cc: Cody P Schafer <cody@linux.vnet.ibm.com> Cc: Hugh Dickens <hughd@google.com> Cc: Paul Mackerras <paulus@samba.org> Cc: Bob Liu <bob.liu@oracle.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2013-07-10 16:04:55 -07:00
config GENERIC_EARLY_IOREMAP
bool
config STACK_MAX_DEFAULT_SIZE_MB
int "Default maximum user stack size for 32-bit processes (MB)"
default 100
range 8 2048
depends on STACK_GROWSUP && (!64BIT || COMPAT)
help
This is the maximum stack size in Megabytes in the VM layout of 32-bit
user processes when the stack grows upwards (currently only on parisc
arch) when the RLIMIT_STACK hard limit is unlimited.
A sane initial value is 100 MB.
config DEFERRED_STRUCT_PAGE_INIT
bool "Defer initialisation of struct pages to kthreads"
mm: make DEFERRED_STRUCT_PAGE_INIT explicitly depend on SPARSEMEM The deferred memory initialization relies on section definitions, e.g PAGES_PER_SECTION, that are only available when CONFIG_SPARSEMEM=y on most architectures. Initially DEFERRED_STRUCT_PAGE_INIT depended on explicit ARCH_SUPPORTS_DEFERRED_STRUCT_PAGE_INIT configuration option, but since the commit 2e3ca40f03bb13709df4 ("mm: relax deferred struct page requirements") this requirement was relaxed and now it is possible to enable DEFERRED_STRUCT_PAGE_INIT on architectures that support DISCONTINGMEM and NO_BOOTMEM which causes build failures. For instance, setting SMP=y and DEFERRED_STRUCT_PAGE_INIT=y on arc causes the following build failure: CC mm/page_alloc.o mm/page_alloc.c: In function 'update_defer_init': mm/page_alloc.c:321:14: error: 'PAGES_PER_SECTION' undeclared (first use in this function); did you mean 'USEC_PER_SEC'? (pfn & (PAGES_PER_SECTION - 1)) == 0) { ^~~~~~~~~~~~~~~~~ USEC_PER_SEC mm/page_alloc.c:321:14: note: each undeclared identifier is reported only once for each function it appears in In file included from include/linux/cache.h:5:0, from include/linux/printk.h:9, from include/linux/kernel.h:14, from include/asm-generic/bug.h:18, from arch/arc/include/asm/bug.h:32, from include/linux/bug.h:5, from include/linux/mmdebug.h:5, from include/linux/mm.h:9, from mm/page_alloc.c:18: mm/page_alloc.c: In function 'deferred_grow_zone': mm/page_alloc.c:1624:52: error: 'PAGES_PER_SECTION' undeclared (first use in this function); did you mean 'USEC_PER_SEC'? unsigned long nr_pages_needed = ALIGN(1 << order, PAGES_PER_SECTION); ^ include/uapi/linux/kernel.h:11:47: note: in definition of macro '__ALIGN_KERNEL_MASK' #define __ALIGN_KERNEL_MASK(x, mask) (((x) + (mask)) & ~(mask)) ^~~~ include/linux/kernel.h:58:22: note: in expansion of macro '__ALIGN_KERNEL' #define ALIGN(x, a) __ALIGN_KERNEL((x), (a)) ^~~~~~~~~~~~~~ mm/page_alloc.c:1624:34: note: in expansion of macro 'ALIGN' unsigned long nr_pages_needed = ALIGN(1 << order, PAGES_PER_SECTION); ^~~~~ In file included from include/asm-generic/bug.h:18:0, from arch/arc/include/asm/bug.h:32, from include/linux/bug.h:5, from include/linux/mmdebug.h:5, from include/linux/mm.h:9, from mm/page_alloc.c:18: mm/page_alloc.c: In function 'free_area_init_node': mm/page_alloc.c:6379:50: error: 'PAGES_PER_SECTION' undeclared (first use in this function); did you mean 'USEC_PER_SEC'? pgdat->static_init_pgcnt = min_t(unsigned long, PAGES_PER_SECTION, ^ include/linux/kernel.h:812:22: note: in definition of macro '__typecheck' (!!(sizeof((typeof(x) *)1 == (typeof(y) *)1))) ^ include/linux/kernel.h:836:24: note: in expansion of macro '__safe_cmp' __builtin_choose_expr(__safe_cmp(x, y), \ ^~~~~~~~~~ include/linux/kernel.h:904:27: note: in expansion of macro '__careful_cmp' #define min_t(type, x, y) __careful_cmp((type)(x), (type)(y), <) ^~~~~~~~~~~~~ mm/page_alloc.c:6379:29: note: in expansion of macro 'min_t' pgdat->static_init_pgcnt = min_t(unsigned long, PAGES_PER_SECTION, ^~~~~ include/linux/kernel.h:836:2: error: first argument to '__builtin_choose_expr' not a constant __builtin_choose_expr(__safe_cmp(x, y), \ ^ include/linux/kernel.h:904:27: note: in expansion of macro '__careful_cmp' #define min_t(type, x, y) __careful_cmp((type)(x), (type)(y), <) ^~~~~~~~~~~~~ mm/page_alloc.c:6379:29: note: in expansion of macro 'min_t' pgdat->static_init_pgcnt = min_t(unsigned long, PAGES_PER_SECTION, ^~~~~ scripts/Makefile.build:317: recipe for target 'mm/page_alloc.o' failed Let's make the DEFERRED_STRUCT_PAGE_INIT explicitly depend on SPARSEMEM as the systems that support DISCONTIGMEM do not seem to have that huge amounts of memory that would make DEFERRED_STRUCT_PAGE_INIT relevant. Link: http://lkml.kernel.org/r/1530279308-24988-1-git-send-email-rppt@linux.vnet.ibm.com Signed-off-by: Mike Rapoport <rppt@linux.vnet.ibm.com> Acked-by: Michal Hocko <mhocko@suse.com> Reviewed-by: Pavel Tatashin <pasha.tatashin@oracle.com> Tested-by: Randy Dunlap <rdunlap@infradead.org> Cc: Pasha Tatashin <Pavel.Tatashin@microsoft.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2018-08-17 15:47:07 -07:00
depends on SPARSEMEM
mm: don't allow deferred pages with NEED_PER_CPU_KM It is unsafe to do virtual to physical translations before mm_init() is called if struct page is needed in order to determine the memory section number (see SECTION_IN_PAGE_FLAGS). This is because only in mm_init() we initialize struct pages for all the allocated memory when deferred struct pages are used. My recent fix in commit c9e97a1997 ("mm: initialize pages on demand during boot") exposed this problem, because it greatly reduced number of pages that are initialized before mm_init(), but the problem existed even before my fix, as Fengguang Wu found. Below is a more detailed explanation of the problem. We initialize struct pages in four places: 1. Early in boot a small set of struct pages is initialized to fill the first section, and lower zones. 2. During mm_init() we initialize "struct pages" for all the memory that is allocated, i.e reserved in memblock. 3. Using on-demand logic when pages are allocated after mm_init call (when memblock is finished) 4. After smp_init() when the rest free deferred pages are initialized. The problem occurs if we try to do va to phys translation of a memory between steps 1 and 2. Because we have not yet initialized struct pages for all the reserved pages, it is inherently unsafe to do va to phys if the translation itself requires access of "struct page" as in case of this combination: CONFIG_SPARSE && !CONFIG_SPARSE_VMEMMAP The following path exposes the problem: start_kernel() trap_init() setup_cpu_entry_areas() setup_cpu_entry_area(cpu) get_cpu_gdt_paddr(cpu) per_cpu_ptr_to_phys(addr) pcpu_addr_to_page(addr) virt_to_page(addr) pfn_to_page(__pa(addr) >> PAGE_SHIFT) We disable this path by not allowing NEED_PER_CPU_KM with deferred struct pages feature. The problems are discussed in these threads: http://lkml.kernel.org/r/20180418135300.inazvpxjxowogyge@wfg-t540p.sh.intel.com http://lkml.kernel.org/r/20180419013128.iurzouiqxvcnpbvz@wfg-t540p.sh.intel.com http://lkml.kernel.org/r/20180426202619.2768-1-pasha.tatashin@oracle.com Link: http://lkml.kernel.org/r/20180515175124.1770-1-pasha.tatashin@oracle.com Fixes: 3a80a7fa7989 ("mm: meminit: initialise a subset of struct pages if CONFIG_DEFERRED_STRUCT_PAGE_INIT is set") Signed-off-by: Pavel Tatashin <pasha.tatashin@oracle.com> Acked-by: Michal Hocko <mhocko@suse.com> Reviewed-by: Andrew Morton <akpm@linux-foundation.org> Cc: Steven Sistare <steven.sistare@oracle.com> Cc: Daniel Jordan <daniel.m.jordan@oracle.com> Cc: Mel Gorman <mgorman@techsingularity.net> Cc: Fengguang Wu <fengguang.wu@intel.com> Cc: Dennis Zhou <dennisszhou@gmail.com> Cc: <stable@vger.kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2018-05-18 16:09:13 -07:00
depends on !NEED_PER_CPU_KM
mm: disable deferred struct page for 32-bit arches Deferred struct page init is needed only on systems with large amount of physical memory to improve boot performance. 32-bit systems do not benefit from this feature. Jiri reported a problem where deferred struct pages do not work well with x86-32: [ 0.035162] Dentry cache hash table entries: 131072 (order: 7, 524288 bytes) [ 0.035725] Inode-cache hash table entries: 65536 (order: 6, 262144 bytes) [ 0.036269] Initializing CPU#0 [ 0.036513] Initializing HighMem for node 0 (00036ffe:0007ffe0) [ 0.038459] page:f6780000 is uninitialized and poisoned [ 0.038460] raw: ffffffff ffffffff ffffffff ffffffff ffffffff ffffffff ffffffff ffffffff [ 0.039509] page dumped because: VM_BUG_ON_PAGE(1 && PageCompound(page)) [ 0.040038] ------------[ cut here ]------------ [ 0.040399] kernel BUG at include/linux/page-flags.h:293! [ 0.040823] invalid opcode: 0000 [#1] SMP PTI [ 0.041166] CPU: 0 PID: 0 Comm: swapper Not tainted 4.19.0-rc1_pt_jiri #9 [ 0.041694] Hardware name: QEMU Standard PC (i440FX + PIIX, 1996), BIOS 1.11.0-20171110_100015-anatol 04/01/2014 [ 0.042496] EIP: free_highmem_page+0x64/0x80 [ 0.042839] Code: 13 46 d8 c1 e8 18 5d 83 e0 03 8d 04 c0 c1 e0 06 ff 80 ec 5f 44 d8 c3 8d b4 26 00 00 00 00 ba 08 65 28 d8 89 d8 e8 fc 71 02 00 <0f> 0b 8d 76 00 8d bc 27 00 00 00 00 ba d0 b1 26 d8 89 d8 e8 e4 71 [ 0.044338] EAX: 0000003c EBX: f6780000 ECX: 00000000 EDX: d856cbe8 [ 0.044868] ESI: 0007ffe0 EDI: d838df20 EBP: d838df00 ESP: d838defc [ 0.045372] DS: 007b ES: 007b FS: 00d8 GS: 00e0 SS: 0068 EFLAGS: 00210086 [ 0.045913] CR0: 80050033 CR2: 00000000 CR3: 18556000 CR4: 00040690 [ 0.046413] DR0: 00000000 DR1: 00000000 DR2: 00000000 DR3: 00000000 [ 0.046913] DR6: fffe0ff0 DR7: 00000400 [ 0.047220] Call Trace: [ 0.047419] add_highpages_with_active_regions+0xbd/0x10d [ 0.047854] set_highmem_pages_init+0x5b/0x71 [ 0.048202] mem_init+0x2b/0x1e8 [ 0.048460] start_kernel+0x1d2/0x425 [ 0.048757] i386_start_kernel+0x93/0x97 [ 0.049073] startup_32_smp+0x164/0x168 [ 0.049379] Modules linked in: [ 0.049626] ---[ end trace 337949378db0abbb ]--- We free highmem pages before their struct pages are initialized: mem_init() set_highmem_pages_init() add_highpages_with_active_regions() free_highmem_page() .. Access uninitialized struct page here.. Because there is no reason to have this feature on 32-bit systems, just disable it. Link: http://lkml.kernel.org/r/20180831150506.31246-1-pavel.tatashin@microsoft.com Fixes: 2e3ca40f03bb ("mm: relax deferred struct page requirements") Signed-off-by: Pavel Tatashin <pavel.tatashin@microsoft.com> Reported-by: Jiri Slaby <jslaby@suse.cz> Acked-by: Michal Hocko <mhocko@suse.com> Cc: <stable@vger.kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Greg Kroah-Hartman <gregkh@linuxfoundation.org>
2018-09-20 12:22:30 -07:00
depends on 64BIT
depends on !KMSAN
mm: parallelize deferred_init_memmap() Deferred struct page init is a significant bottleneck in kernel boot. Optimizing it maximizes availability for large-memory systems and allows spinning up short-lived VMs as needed without having to leave them running. It also benefits bare metal machines hosting VMs that are sensitive to downtime. In projects such as VMM Fast Restart[1], where guest state is preserved across kexec reboot, it helps prevent application and network timeouts in the guests. Multithread to take full advantage of system memory bandwidth. The maximum number of threads is capped at the number of CPUs on the node because speedups always improve with additional threads on every system tested, and at this phase of boot, the system is otherwise idle and waiting on page init to finish. Helper threads operate on section-aligned ranges to both avoid false sharing when setting the pageblock's migrate type and to avoid accessing uninitialized buddy pages, though max order alignment is enough for the latter. The minimum chunk size is also a section. There was benefit to using multiple threads even on relatively small memory (1G) systems, and this is the smallest size that the alignment allows. The time (milliseconds) is the slowest node to initialize since boot blocks until all nodes finish. intel_pstate is loaded in active mode without hwp and with turbo enabled, and intel_idle is active as well. Intel(R) Xeon(R) Platinum 8167M CPU @ 2.00GHz (Skylake, bare metal) 2 nodes * 26 cores * 2 threads = 104 CPUs 384G/node = 768G memory kernel boot deferred init ------------------------ ------------------------ node% (thr) speedup time_ms (stdev) speedup time_ms (stdev) ( 0) -- 4089.7 ( 8.1) -- 1785.7 ( 7.6) 2% ( 1) 1.7% 4019.3 ( 1.5) 3.8% 1717.7 ( 11.8) 12% ( 6) 34.9% 2662.7 ( 2.9) 79.9% 359.3 ( 0.6) 25% ( 13) 39.9% 2459.0 ( 3.6) 91.2% 157.0 ( 0.0) 37% ( 19) 39.2% 2485.0 ( 29.7) 90.4% 172.0 ( 28.6) 50% ( 26) 39.3% 2482.7 ( 25.7) 90.3% 173.7 ( 30.0) 75% ( 39) 39.0% 2495.7 ( 5.5) 89.4% 190.0 ( 1.0) 100% ( 52) 40.2% 2443.7 ( 3.8) 92.3% 138.0 ( 1.0) Intel(R) Xeon(R) CPU E5-2699C v4 @ 2.20GHz (Broadwell, kvm guest) 1 node * 16 cores * 2 threads = 32 CPUs 192G/node = 192G memory kernel boot deferred init ------------------------ ------------------------ node% (thr) speedup time_ms (stdev) speedup time_ms (stdev) ( 0) -- 1988.7 ( 9.6) -- 1096.0 ( 11.5) 3% ( 1) 1.1% 1967.0 ( 17.6) 0.3% 1092.7 ( 11.0) 12% ( 4) 41.1% 1170.3 ( 14.2) 73.8% 287.0 ( 3.6) 25% ( 8) 47.1% 1052.7 ( 21.9) 83.9% 177.0 ( 13.5) 38% ( 12) 48.9% 1016.3 ( 12.1) 86.8% 144.7 ( 1.5) 50% ( 16) 48.9% 1015.7 ( 8.1) 87.8% 134.0 ( 4.4) 75% ( 24) 49.1% 1012.3 ( 3.1) 88.1% 130.3 ( 2.3) 100% ( 32) 49.5% 1004.0 ( 5.3) 88.5% 125.7 ( 2.1) Intel(R) Xeon(R) CPU E5-2699 v3 @ 2.30GHz (Haswell, bare metal) 2 nodes * 18 cores * 2 threads = 72 CPUs 128G/node = 256G memory kernel boot deferred init ------------------------ ------------------------ node% (thr) speedup time_ms (stdev) speedup time_ms (stdev) ( 0) -- 1680.0 ( 4.6) -- 627.0 ( 4.0) 3% ( 1) 0.3% 1675.7 ( 4.5) -0.2% 628.0 ( 3.6) 11% ( 4) 25.6% 1250.7 ( 2.1) 67.9% 201.0 ( 0.0) 25% ( 9) 30.7% 1164.0 ( 17.3) 81.8% 114.3 ( 17.7) 36% ( 13) 31.4% 1152.7 ( 10.8) 84.0% 100.3 ( 17.9) 50% ( 18) 31.5% 1150.7 ( 9.3) 83.9% 101.0 ( 14.1) 75% ( 27) 31.7% 1148.0 ( 5.6) 84.5% 97.3 ( 6.4) 100% ( 36) 32.0% 1142.3 ( 4.0) 85.6% 90.0 ( 1.0) AMD EPYC 7551 32-Core Processor (Zen, kvm guest) 1 node * 8 cores * 2 threads = 16 CPUs 64G/node = 64G memory kernel boot deferred init ------------------------ ------------------------ node% (thr) speedup time_ms (stdev) speedup time_ms (stdev) ( 0) -- 1029.3 ( 25.1) -- 240.7 ( 1.5) 6% ( 1) -0.6% 1036.0 ( 7.8) -2.2% 246.0 ( 0.0) 12% ( 2) 11.8% 907.7 ( 8.6) 44.7% 133.0 ( 1.0) 25% ( 4) 13.9% 886.0 ( 10.6) 62.6% 90.0 ( 6.0) 38% ( 6) 17.8% 845.7 ( 14.2) 69.1% 74.3 ( 3.8) 50% ( 8) 16.8% 856.0 ( 22.1) 72.9% 65.3 ( 5.7) 75% ( 12) 15.4% 871.0 ( 29.2) 79.8% 48.7 ( 7.4) 100% ( 16) 21.0% 813.7 ( 21.0) 80.5% 47.0 ( 5.2) Server-oriented distros that enable deferred page init sometimes run in small VMs, and they still benefit even though the fraction of boot time saved is smaller: AMD EPYC 7551 32-Core Processor (Zen, kvm guest) 1 node * 2 cores * 2 threads = 4 CPUs 16G/node = 16G memory kernel boot deferred init ------------------------ ------------------------ node% (thr) speedup time_ms (stdev) speedup time_ms (stdev) ( 0) -- 716.0 ( 14.0) -- 49.7 ( 0.6) 25% ( 1) 1.8% 703.0 ( 5.3) -4.0% 51.7 ( 0.6) 50% ( 2) 1.6% 704.7 ( 1.2) 43.0% 28.3 ( 0.6) 75% ( 3) 2.7% 696.7 ( 13.1) 49.7% 25.0 ( 0.0) 100% ( 4) 4.1% 687.0 ( 10.4) 55.7% 22.0 ( 0.0) Intel(R) Xeon(R) CPU E5-2699 v3 @ 2.30GHz (Haswell, kvm guest) 1 node * 2 cores * 2 threads = 4 CPUs 14G/node = 14G memory kernel boot deferred init ------------------------ ------------------------ node% (thr) speedup time_ms (stdev) speedup time_ms (stdev) ( 0) -- 787.7 ( 6.4) -- 122.3 ( 0.6) 25% ( 1) 0.2% 786.3 ( 10.8) -2.5% 125.3 ( 2.1) 50% ( 2) 5.9% 741.0 ( 13.9) 37.6% 76.3 ( 19.7) 75% ( 3) 8.3% 722.0 ( 19.0) 49.9% 61.3 ( 3.2) 100% ( 4) 9.3% 714.7 ( 9.5) 56.4% 53.3 ( 1.5) On Josh's 96-CPU and 192G memory system: Without this patch series: [ 0.487132] node 0 initialised, 23398907 pages in 292ms [ 0.499132] node 1 initialised, 24189223 pages in 304ms ... [ 0.629376] Run /sbin/init as init process With this patch series: [ 0.231435] node 1 initialised, 24189223 pages in 32ms [ 0.236718] node 0 initialised, 23398907 pages in 36ms [1] https://static.sched.com/hosted_files/kvmforum2019/66/VMM-fast-restart_kvmforum2019.pdf Signed-off-by: Daniel Jordan <daniel.m.jordan@oracle.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Tested-by: Josh Triplett <josh@joshtriplett.org> Reviewed-by: Alexander Duyck <alexander.h.duyck@linux.intel.com> Cc: Alex Williamson <alex.williamson@redhat.com> Cc: Dan Williams <dan.j.williams@intel.com> Cc: Dave Hansen <dave.hansen@linux.intel.com> Cc: David Hildenbrand <david@redhat.com> Cc: Herbert Xu <herbert@gondor.apana.org.au> Cc: Jason Gunthorpe <jgg@ziepe.ca> Cc: Jonathan Corbet <corbet@lwn.net> Cc: Kirill Tkhai <ktkhai@virtuozzo.com> Cc: Michal Hocko <mhocko@kernel.org> Cc: Pavel Machek <pavel@ucw.cz> Cc: Pavel Tatashin <pasha.tatashin@soleen.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Randy Dunlap <rdunlap@infradead.org> Cc: Robert Elliott <elliott@hpe.com> Cc: Shile Zhang <shile.zhang@linux.alibaba.com> Cc: Steffen Klassert <steffen.klassert@secunet.com> Cc: Steven Sistare <steven.sistare@oracle.com> Cc: Tejun Heo <tj@kernel.org> Cc: Zi Yan <ziy@nvidia.com> Link: http://lkml.kernel.org/r/20200527173608.2885243-7-daniel.m.jordan@oracle.com Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2020-06-03 15:59:51 -07:00
select PADATA
help
Ordinarily all struct pages are initialised during early boot in a
single thread. On very large machines this can take a considerable
amount of time. If this option is set, large machines will bring up
mm: parallelize deferred_init_memmap() Deferred struct page init is a significant bottleneck in kernel boot. Optimizing it maximizes availability for large-memory systems and allows spinning up short-lived VMs as needed without having to leave them running. It also benefits bare metal machines hosting VMs that are sensitive to downtime. In projects such as VMM Fast Restart[1], where guest state is preserved across kexec reboot, it helps prevent application and network timeouts in the guests. Multithread to take full advantage of system memory bandwidth. The maximum number of threads is capped at the number of CPUs on the node because speedups always improve with additional threads on every system tested, and at this phase of boot, the system is otherwise idle and waiting on page init to finish. Helper threads operate on section-aligned ranges to both avoid false sharing when setting the pageblock's migrate type and to avoid accessing uninitialized buddy pages, though max order alignment is enough for the latter. The minimum chunk size is also a section. There was benefit to using multiple threads even on relatively small memory (1G) systems, and this is the smallest size that the alignment allows. The time (milliseconds) is the slowest node to initialize since boot blocks until all nodes finish. intel_pstate is loaded in active mode without hwp and with turbo enabled, and intel_idle is active as well. Intel(R) Xeon(R) Platinum 8167M CPU @ 2.00GHz (Skylake, bare metal) 2 nodes * 26 cores * 2 threads = 104 CPUs 384G/node = 768G memory kernel boot deferred init ------------------------ ------------------------ node% (thr) speedup time_ms (stdev) speedup time_ms (stdev) ( 0) -- 4089.7 ( 8.1) -- 1785.7 ( 7.6) 2% ( 1) 1.7% 4019.3 ( 1.5) 3.8% 1717.7 ( 11.8) 12% ( 6) 34.9% 2662.7 ( 2.9) 79.9% 359.3 ( 0.6) 25% ( 13) 39.9% 2459.0 ( 3.6) 91.2% 157.0 ( 0.0) 37% ( 19) 39.2% 2485.0 ( 29.7) 90.4% 172.0 ( 28.6) 50% ( 26) 39.3% 2482.7 ( 25.7) 90.3% 173.7 ( 30.0) 75% ( 39) 39.0% 2495.7 ( 5.5) 89.4% 190.0 ( 1.0) 100% ( 52) 40.2% 2443.7 ( 3.8) 92.3% 138.0 ( 1.0) Intel(R) Xeon(R) CPU E5-2699C v4 @ 2.20GHz (Broadwell, kvm guest) 1 node * 16 cores * 2 threads = 32 CPUs 192G/node = 192G memory kernel boot deferred init ------------------------ ------------------------ node% (thr) speedup time_ms (stdev) speedup time_ms (stdev) ( 0) -- 1988.7 ( 9.6) -- 1096.0 ( 11.5) 3% ( 1) 1.1% 1967.0 ( 17.6) 0.3% 1092.7 ( 11.0) 12% ( 4) 41.1% 1170.3 ( 14.2) 73.8% 287.0 ( 3.6) 25% ( 8) 47.1% 1052.7 ( 21.9) 83.9% 177.0 ( 13.5) 38% ( 12) 48.9% 1016.3 ( 12.1) 86.8% 144.7 ( 1.5) 50% ( 16) 48.9% 1015.7 ( 8.1) 87.8% 134.0 ( 4.4) 75% ( 24) 49.1% 1012.3 ( 3.1) 88.1% 130.3 ( 2.3) 100% ( 32) 49.5% 1004.0 ( 5.3) 88.5% 125.7 ( 2.1) Intel(R) Xeon(R) CPU E5-2699 v3 @ 2.30GHz (Haswell, bare metal) 2 nodes * 18 cores * 2 threads = 72 CPUs 128G/node = 256G memory kernel boot deferred init ------------------------ ------------------------ node% (thr) speedup time_ms (stdev) speedup time_ms (stdev) ( 0) -- 1680.0 ( 4.6) -- 627.0 ( 4.0) 3% ( 1) 0.3% 1675.7 ( 4.5) -0.2% 628.0 ( 3.6) 11% ( 4) 25.6% 1250.7 ( 2.1) 67.9% 201.0 ( 0.0) 25% ( 9) 30.7% 1164.0 ( 17.3) 81.8% 114.3 ( 17.7) 36% ( 13) 31.4% 1152.7 ( 10.8) 84.0% 100.3 ( 17.9) 50% ( 18) 31.5% 1150.7 ( 9.3) 83.9% 101.0 ( 14.1) 75% ( 27) 31.7% 1148.0 ( 5.6) 84.5% 97.3 ( 6.4) 100% ( 36) 32.0% 1142.3 ( 4.0) 85.6% 90.0 ( 1.0) AMD EPYC 7551 32-Core Processor (Zen, kvm guest) 1 node * 8 cores * 2 threads = 16 CPUs 64G/node = 64G memory kernel boot deferred init ------------------------ ------------------------ node% (thr) speedup time_ms (stdev) speedup time_ms (stdev) ( 0) -- 1029.3 ( 25.1) -- 240.7 ( 1.5) 6% ( 1) -0.6% 1036.0 ( 7.8) -2.2% 246.0 ( 0.0) 12% ( 2) 11.8% 907.7 ( 8.6) 44.7% 133.0 ( 1.0) 25% ( 4) 13.9% 886.0 ( 10.6) 62.6% 90.0 ( 6.0) 38% ( 6) 17.8% 845.7 ( 14.2) 69.1% 74.3 ( 3.8) 50% ( 8) 16.8% 856.0 ( 22.1) 72.9% 65.3 ( 5.7) 75% ( 12) 15.4% 871.0 ( 29.2) 79.8% 48.7 ( 7.4) 100% ( 16) 21.0% 813.7 ( 21.0) 80.5% 47.0 ( 5.2) Server-oriented distros that enable deferred page init sometimes run in small VMs, and they still benefit even though the fraction of boot time saved is smaller: AMD EPYC 7551 32-Core Processor (Zen, kvm guest) 1 node * 2 cores * 2 threads = 4 CPUs 16G/node = 16G memory kernel boot deferred init ------------------------ ------------------------ node% (thr) speedup time_ms (stdev) speedup time_ms (stdev) ( 0) -- 716.0 ( 14.0) -- 49.7 ( 0.6) 25% ( 1) 1.8% 703.0 ( 5.3) -4.0% 51.7 ( 0.6) 50% ( 2) 1.6% 704.7 ( 1.2) 43.0% 28.3 ( 0.6) 75% ( 3) 2.7% 696.7 ( 13.1) 49.7% 25.0 ( 0.0) 100% ( 4) 4.1% 687.0 ( 10.4) 55.7% 22.0 ( 0.0) Intel(R) Xeon(R) CPU E5-2699 v3 @ 2.30GHz (Haswell, kvm guest) 1 node * 2 cores * 2 threads = 4 CPUs 14G/node = 14G memory kernel boot deferred init ------------------------ ------------------------ node% (thr) speedup time_ms (stdev) speedup time_ms (stdev) ( 0) -- 787.7 ( 6.4) -- 122.3 ( 0.6) 25% ( 1) 0.2% 786.3 ( 10.8) -2.5% 125.3 ( 2.1) 50% ( 2) 5.9% 741.0 ( 13.9) 37.6% 76.3 ( 19.7) 75% ( 3) 8.3% 722.0 ( 19.0) 49.9% 61.3 ( 3.2) 100% ( 4) 9.3% 714.7 ( 9.5) 56.4% 53.3 ( 1.5) On Josh's 96-CPU and 192G memory system: Without this patch series: [ 0.487132] node 0 initialised, 23398907 pages in 292ms [ 0.499132] node 1 initialised, 24189223 pages in 304ms ... [ 0.629376] Run /sbin/init as init process With this patch series: [ 0.231435] node 1 initialised, 24189223 pages in 32ms [ 0.236718] node 0 initialised, 23398907 pages in 36ms [1] https://static.sched.com/hosted_files/kvmforum2019/66/VMM-fast-restart_kvmforum2019.pdf Signed-off-by: Daniel Jordan <daniel.m.jordan@oracle.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Tested-by: Josh Triplett <josh@joshtriplett.org> Reviewed-by: Alexander Duyck <alexander.h.duyck@linux.intel.com> Cc: Alex Williamson <alex.williamson@redhat.com> Cc: Dan Williams <dan.j.williams@intel.com> Cc: Dave Hansen <dave.hansen@linux.intel.com> Cc: David Hildenbrand <david@redhat.com> Cc: Herbert Xu <herbert@gondor.apana.org.au> Cc: Jason Gunthorpe <jgg@ziepe.ca> Cc: Jonathan Corbet <corbet@lwn.net> Cc: Kirill Tkhai <ktkhai@virtuozzo.com> Cc: Michal Hocko <mhocko@kernel.org> Cc: Pavel Machek <pavel@ucw.cz> Cc: Pavel Tatashin <pasha.tatashin@soleen.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Randy Dunlap <rdunlap@infradead.org> Cc: Robert Elliott <elliott@hpe.com> Cc: Shile Zhang <shile.zhang@linux.alibaba.com> Cc: Steffen Klassert <steffen.klassert@secunet.com> Cc: Steven Sistare <steven.sistare@oracle.com> Cc: Tejun Heo <tj@kernel.org> Cc: Zi Yan <ziy@nvidia.com> Link: http://lkml.kernel.org/r/20200527173608.2885243-7-daniel.m.jordan@oracle.com Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2020-06-03 15:59:51 -07:00
a subset of memmap at boot and then initialise the rest in parallel.
This has a potential performance impact on tasks running early in the
lifetime of the system until these kthreads finish the
initialisation.
mm/idle_page_tracking: make PG_idle reusable PG_idle and PG_young allow the two PTE Accessed bit users, Idle Page Tracking and the reclaim logic concurrently work while not interfering with each other. That is, when they need to clear the Accessed bit, they set PG_young to represent the previous state of the bit, respectively. And when they need to read the bit, if the bit is cleared, they further read the PG_young to know whether the other has cleared the bit meanwhile or not. For yet another user of the PTE Accessed bit, we could add another page flag, or extend the mechanism to use the flags. For the DAMON usecase, however, we don't need to do that just yet. IDLE_PAGE_TRACKING and DAMON are mutually exclusive, so there's only ever going to be one user of the current set of flags. In this commit, we split out the CONFIG options to allow for the use of PG_young and PG_idle outside of idle page tracking. In the next commit, DAMON's reference implementation of the virtual memory address space monitoring primitives will use it. [sjpark@amazon.de: set PAGE_EXTENSION for non-64BIT] Link: https://lkml.kernel.org/r/20210806095153.6444-1-sj38.park@gmail.com [akpm@linux-foundation.org: tweak Kconfig text] [sjpark@amazon.de: hide PAGE_IDLE_FLAG from users] Link: https://lkml.kernel.org/r/20210813081238.34705-1-sj38.park@gmail.com Link: https://lkml.kernel.org/r/20210716081449.22187-5-sj38.park@gmail.com Signed-off-by: SeongJae Park <sjpark@amazon.de> Reviewed-by: Shakeel Butt <shakeelb@google.com> Reviewed-by: Fernand Sieber <sieberf@amazon.com> Cc: Alexander Shishkin <alexander.shishkin@linux.intel.com> Cc: Amit Shah <amit@kernel.org> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: Brendan Higgins <brendanhiggins@google.com> Cc: David Hildenbrand <david@redhat.com> Cc: David Rientjes <rientjes@google.com> Cc: David Woodhouse <dwmw@amazon.com> Cc: Fan Du <fan.du@intel.com> Cc: Greg Kroah-Hartman <greg@kroah.com> Cc: Greg Thelen <gthelen@google.com> Cc: Ingo Molnar <mingo@redhat.com> Cc: Joe Perches <joe@perches.com> Cc: Jonathan Cameron <Jonathan.Cameron@huawei.com> Cc: Jonathan Corbet <corbet@lwn.net> Cc: Leonard Foerster <foersleo@amazon.de> Cc: Marco Elver <elver@google.com> Cc: Markus Boehme <markubo@amazon.de> Cc: Maximilian Heyne <mheyne@amazon.de> Cc: Mel Gorman <mgorman@suse.de> Cc: Minchan Kim <minchan@kernel.org> Cc: Namhyung Kim <namhyung@kernel.org> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Rik van Riel <riel@surriel.com> Cc: Shuah Khan <shuah@kernel.org> Cc: Steven Rostedt (VMware) <rostedt@goodmis.org> Cc: Vladimir Davydov <vdavydov.dev@gmail.com> Cc: Vlastimil Babka <vbabka@suse.cz> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2021-09-07 19:56:40 -07:00
config PAGE_IDLE_FLAG
bool
select PAGE_EXTENSION if !64BIT
help
This adds PG_idle and PG_young flags to 'struct page'. PTE Accessed
bit writers can set the state of the bit in the flags so that PTE
Accessed bit readers may avoid disturbance.
mm: introduce idle page tracking Knowing the portion of memory that is not used by a certain application or memory cgroup (idle memory) can be useful for partitioning the system efficiently, e.g. by setting memory cgroup limits appropriately. Currently, the only means to estimate the amount of idle memory provided by the kernel is /proc/PID/{clear_refs,smaps}: the user can clear the access bit for all pages mapped to a particular process by writing 1 to clear_refs, wait for some time, and then count smaps:Referenced. However, this method has two serious shortcomings: - it does not count unmapped file pages - it affects the reclaimer logic To overcome these drawbacks, this patch introduces two new page flags, Idle and Young, and a new sysfs file, /sys/kernel/mm/page_idle/bitmap. A page's Idle flag can only be set from userspace by setting bit in /sys/kernel/mm/page_idle/bitmap at the offset corresponding to the page, and it is cleared whenever the page is accessed either through page tables (it is cleared in page_referenced() in this case) or using the read(2) system call (mark_page_accessed()). Thus by setting the Idle flag for pages of a particular workload, which can be found e.g. by reading /proc/PID/pagemap, waiting for some time to let the workload access its working set, and then reading the bitmap file, one can estimate the amount of pages that are not used by the workload. The Young page flag is used to avoid interference with the memory reclaimer. A page's Young flag is set whenever the Access bit of a page table entry pointing to the page is cleared by writing to the bitmap file. If page_referenced() is called on a Young page, it will add 1 to its return value, therefore concealing the fact that the Access bit was cleared. Note, since there is no room for extra page flags on 32 bit, this feature uses extended page flags when compiled on 32 bit. [akpm@linux-foundation.org: fix build] [akpm@linux-foundation.org: kpageidle requires an MMU] [akpm@linux-foundation.org: decouple from page-flags rework] Signed-off-by: Vladimir Davydov <vdavydov@parallels.com> Reviewed-by: Andres Lagar-Cavilla <andreslc@google.com> Cc: Minchan Kim <minchan@kernel.org> Cc: Raghavendra K T <raghavendra.kt@linux.vnet.ibm.com> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Michal Hocko <mhocko@suse.cz> Cc: Greg Thelen <gthelen@google.com> Cc: Michel Lespinasse <walken@google.com> Cc: David Rientjes <rientjes@google.com> Cc: Pavel Emelyanov <xemul@parallels.com> Cc: Cyrill Gorcunov <gorcunov@openvz.org> Cc: Jonathan Corbet <corbet@lwn.net> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-09-09 15:35:45 -07:00
config IDLE_PAGE_TRACKING
bool "Enable idle page tracking"
depends on SYSFS && MMU
mm/idle_page_tracking: make PG_idle reusable PG_idle and PG_young allow the two PTE Accessed bit users, Idle Page Tracking and the reclaim logic concurrently work while not interfering with each other. That is, when they need to clear the Accessed bit, they set PG_young to represent the previous state of the bit, respectively. And when they need to read the bit, if the bit is cleared, they further read the PG_young to know whether the other has cleared the bit meanwhile or not. For yet another user of the PTE Accessed bit, we could add another page flag, or extend the mechanism to use the flags. For the DAMON usecase, however, we don't need to do that just yet. IDLE_PAGE_TRACKING and DAMON are mutually exclusive, so there's only ever going to be one user of the current set of flags. In this commit, we split out the CONFIG options to allow for the use of PG_young and PG_idle outside of idle page tracking. In the next commit, DAMON's reference implementation of the virtual memory address space monitoring primitives will use it. [sjpark@amazon.de: set PAGE_EXTENSION for non-64BIT] Link: https://lkml.kernel.org/r/20210806095153.6444-1-sj38.park@gmail.com [akpm@linux-foundation.org: tweak Kconfig text] [sjpark@amazon.de: hide PAGE_IDLE_FLAG from users] Link: https://lkml.kernel.org/r/20210813081238.34705-1-sj38.park@gmail.com Link: https://lkml.kernel.org/r/20210716081449.22187-5-sj38.park@gmail.com Signed-off-by: SeongJae Park <sjpark@amazon.de> Reviewed-by: Shakeel Butt <shakeelb@google.com> Reviewed-by: Fernand Sieber <sieberf@amazon.com> Cc: Alexander Shishkin <alexander.shishkin@linux.intel.com> Cc: Amit Shah <amit@kernel.org> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: Brendan Higgins <brendanhiggins@google.com> Cc: David Hildenbrand <david@redhat.com> Cc: David Rientjes <rientjes@google.com> Cc: David Woodhouse <dwmw@amazon.com> Cc: Fan Du <fan.du@intel.com> Cc: Greg Kroah-Hartman <greg@kroah.com> Cc: Greg Thelen <gthelen@google.com> Cc: Ingo Molnar <mingo@redhat.com> Cc: Joe Perches <joe@perches.com> Cc: Jonathan Cameron <Jonathan.Cameron@huawei.com> Cc: Jonathan Corbet <corbet@lwn.net> Cc: Leonard Foerster <foersleo@amazon.de> Cc: Marco Elver <elver@google.com> Cc: Markus Boehme <markubo@amazon.de> Cc: Maximilian Heyne <mheyne@amazon.de> Cc: Mel Gorman <mgorman@suse.de> Cc: Minchan Kim <minchan@kernel.org> Cc: Namhyung Kim <namhyung@kernel.org> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Rik van Riel <riel@surriel.com> Cc: Shuah Khan <shuah@kernel.org> Cc: Steven Rostedt (VMware) <rostedt@goodmis.org> Cc: Vladimir Davydov <vdavydov.dev@gmail.com> Cc: Vlastimil Babka <vbabka@suse.cz> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2021-09-07 19:56:40 -07:00
select PAGE_IDLE_FLAG
mm: introduce idle page tracking Knowing the portion of memory that is not used by a certain application or memory cgroup (idle memory) can be useful for partitioning the system efficiently, e.g. by setting memory cgroup limits appropriately. Currently, the only means to estimate the amount of idle memory provided by the kernel is /proc/PID/{clear_refs,smaps}: the user can clear the access bit for all pages mapped to a particular process by writing 1 to clear_refs, wait for some time, and then count smaps:Referenced. However, this method has two serious shortcomings: - it does not count unmapped file pages - it affects the reclaimer logic To overcome these drawbacks, this patch introduces two new page flags, Idle and Young, and a new sysfs file, /sys/kernel/mm/page_idle/bitmap. A page's Idle flag can only be set from userspace by setting bit in /sys/kernel/mm/page_idle/bitmap at the offset corresponding to the page, and it is cleared whenever the page is accessed either through page tables (it is cleared in page_referenced() in this case) or using the read(2) system call (mark_page_accessed()). Thus by setting the Idle flag for pages of a particular workload, which can be found e.g. by reading /proc/PID/pagemap, waiting for some time to let the workload access its working set, and then reading the bitmap file, one can estimate the amount of pages that are not used by the workload. The Young page flag is used to avoid interference with the memory reclaimer. A page's Young flag is set whenever the Access bit of a page table entry pointing to the page is cleared by writing to the bitmap file. If page_referenced() is called on a Young page, it will add 1 to its return value, therefore concealing the fact that the Access bit was cleared. Note, since there is no room for extra page flags on 32 bit, this feature uses extended page flags when compiled on 32 bit. [akpm@linux-foundation.org: fix build] [akpm@linux-foundation.org: kpageidle requires an MMU] [akpm@linux-foundation.org: decouple from page-flags rework] Signed-off-by: Vladimir Davydov <vdavydov@parallels.com> Reviewed-by: Andres Lagar-Cavilla <andreslc@google.com> Cc: Minchan Kim <minchan@kernel.org> Cc: Raghavendra K T <raghavendra.kt@linux.vnet.ibm.com> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Michal Hocko <mhocko@suse.cz> Cc: Greg Thelen <gthelen@google.com> Cc: Michel Lespinasse <walken@google.com> Cc: David Rientjes <rientjes@google.com> Cc: Pavel Emelyanov <xemul@parallels.com> Cc: Cyrill Gorcunov <gorcunov@openvz.org> Cc: Jonathan Corbet <corbet@lwn.net> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-09-09 15:35:45 -07:00
help
This feature allows to estimate the amount of user pages that have
not been touched during a given period of time. This information can
be useful to tune memory cgroup limits and/or for job placement
within a compute cluster.
See Documentation/admin-guide/mm/idle_page_tracking.rst for
more details.
mm: introduce idle page tracking Knowing the portion of memory that is not used by a certain application or memory cgroup (idle memory) can be useful for partitioning the system efficiently, e.g. by setting memory cgroup limits appropriately. Currently, the only means to estimate the amount of idle memory provided by the kernel is /proc/PID/{clear_refs,smaps}: the user can clear the access bit for all pages mapped to a particular process by writing 1 to clear_refs, wait for some time, and then count smaps:Referenced. However, this method has two serious shortcomings: - it does not count unmapped file pages - it affects the reclaimer logic To overcome these drawbacks, this patch introduces two new page flags, Idle and Young, and a new sysfs file, /sys/kernel/mm/page_idle/bitmap. A page's Idle flag can only be set from userspace by setting bit in /sys/kernel/mm/page_idle/bitmap at the offset corresponding to the page, and it is cleared whenever the page is accessed either through page tables (it is cleared in page_referenced() in this case) or using the read(2) system call (mark_page_accessed()). Thus by setting the Idle flag for pages of a particular workload, which can be found e.g. by reading /proc/PID/pagemap, waiting for some time to let the workload access its working set, and then reading the bitmap file, one can estimate the amount of pages that are not used by the workload. The Young page flag is used to avoid interference with the memory reclaimer. A page's Young flag is set whenever the Access bit of a page table entry pointing to the page is cleared by writing to the bitmap file. If page_referenced() is called on a Young page, it will add 1 to its return value, therefore concealing the fact that the Access bit was cleared. Note, since there is no room for extra page flags on 32 bit, this feature uses extended page flags when compiled on 32 bit. [akpm@linux-foundation.org: fix build] [akpm@linux-foundation.org: kpageidle requires an MMU] [akpm@linux-foundation.org: decouple from page-flags rework] Signed-off-by: Vladimir Davydov <vdavydov@parallels.com> Reviewed-by: Andres Lagar-Cavilla <andreslc@google.com> Cc: Minchan Kim <minchan@kernel.org> Cc: Raghavendra K T <raghavendra.kt@linux.vnet.ibm.com> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Michal Hocko <mhocko@suse.cz> Cc: Greg Thelen <gthelen@google.com> Cc: Michel Lespinasse <walken@google.com> Cc: David Rientjes <rientjes@google.com> Cc: Pavel Emelyanov <xemul@parallels.com> Cc: Cyrill Gorcunov <gorcunov@openvz.org> Cc: Jonathan Corbet <corbet@lwn.net> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-09-09 15:35:45 -07:00
Introduce cpu_dcache_is_aliasing() across all architectures Introduce a generic way to query whether the data cache is virtually aliased on all architectures. Its purpose is to ensure that subsystems which are incompatible with virtually aliased data caches (e.g. FS_DAX) can reliably query this. For data cache aliasing, there are three scenarios dependending on the architecture. Here is a breakdown based on my understanding: A) The data cache is always aliasing: * arc * csky * m68k (note: shared memory mappings are incoherent ? SHMLBA is missing there.) * sh * parisc B) The data cache aliasing is statically known or depends on querying CPU state at runtime: * arm (cache_is_vivt() || cache_is_vipt_aliasing()) * mips (cpu_has_dc_aliases) * nios2 (NIOS2_DCACHE_SIZE > PAGE_SIZE) * sparc32 (vac_cache_size > PAGE_SIZE) * sparc64 (L1DCACHE_SIZE > PAGE_SIZE) * xtensa (DCACHE_WAY_SIZE > PAGE_SIZE) C) The data cache is never aliasing: * alpha * arm64 (aarch64) * hexagon * loongarch (but with incoherent write buffers, which are disabled since commit d23b7795 ("LoongArch: Change SHMLBA from SZ_64K to PAGE_SIZE")) * microblaze * openrisc * powerpc * riscv * s390 * um * x86 Require architectures in A) and B) to select ARCH_HAS_CPU_CACHE_ALIASING and implement "cpu_dcache_is_aliasing()". Architectures in C) don't select ARCH_HAS_CPU_CACHE_ALIASING, and thus cpu_dcache_is_aliasing() simply evaluates to "false". Note that this leaves "cpu_icache_is_aliasing()" to be implemented as future work. This would be useful to gate features like XIP on architectures which have aliasing CPU dcache-icache but not CPU dcache-dcache. Use "cpu_dcache" and "cpu_cache" rather than just "dcache" and "cache" to clarify that we really mean "CPU data cache" and "CPU cache" to eliminate any possible confusion with VFS "dentry cache" and "page cache". Link: https://lore.kernel.org/lkml/20030910210416.GA24258@mail.jlokier.co.uk/ Link: https://lkml.kernel.org/r/20240215144633.96437-9-mathieu.desnoyers@efficios.com Fixes: d92576f1167c ("dax: does not work correctly with virtual aliasing caches") Signed-off-by: Mathieu Desnoyers <mathieu.desnoyers@efficios.com> Cc: Dan Williams <dan.j.williams@intel.com> Cc: Vishal Verma <vishal.l.verma@intel.com> Cc: Dave Jiang <dave.jiang@intel.com> Cc: Matthew Wilcox <willy@infradead.org> Cc: Arnd Bergmann <arnd@arndb.de> Cc: Russell King <linux@armlinux.org.uk> Cc: Alasdair Kergon <agk@redhat.com> Cc: Christoph Hellwig <hch@lst.de> Cc: Dave Chinner <david@fromorbit.com> Cc: Heiko Carstens <hca@linux.ibm.com> Cc: kernel test robot <lkp@intel.com> Cc: Michael Sclafani <dm-devel@lists.linux.dev> Cc: Mike Snitzer <snitzer@kernel.org> Cc: Mikulas Patocka <mpatocka@redhat.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2024-02-15 07:46:32 -07:00
# Architectures which implement cpu_dcache_is_aliasing() to query
# whether the data caches are aliased (VIVT or VIPT with dcache
# aliasing) need to select this.
config ARCH_HAS_CPU_CACHE_ALIASING
bool
mm: generalize ARCH_HAS_CACHE_LINE_SIZE Patch series "mm: some config cleanups", v2. This series contains config cleanup patches which reduces code duplication across platforms and also improves maintainability. There is no functional change intended with this series. This patch (of 6): ARCH_HAS_CACHE_LINE_SIZE config has duplicate definitions on platforms that subscribe it. Instead, just make it a generic option which can be selected on applicable platforms. This change reduces code duplication and makes it cleaner. Link: https://lkml.kernel.org/r/1617259448-22529-1-git-send-email-anshuman.khandual@arm.com Link: https://lkml.kernel.org/r/1617259448-22529-2-git-send-email-anshuman.khandual@arm.com Signed-off-by: Anshuman Khandual <anshuman.khandual@arm.com> Acked-by: Catalin Marinas <catalin.marinas@arm.com> [arm64] Acked-by: Vineet Gupta <vgupta@synopsys.com> [arc] Cc: Will Deacon <will@kernel.org> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Ingo Molnar <mingo@redhat.com> Cc: Borislav Petkov <bp@alien8.de> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Albert Ou <aou@eecs.berkeley.edu> Cc: Alexander Viro <viro@zeniv.linux.org.uk> Cc: Arnd Bergmann <arnd@arndb.de> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: Christian Borntraeger <borntraeger@de.ibm.com> Cc: Heiko Carstens <hca@linux.ibm.com> Cc: Helge Deller <deller@gmx.de> Cc: "James E.J. Bottomley" <James.Bottomley@HansenPartnership.com> Cc: Michael Ellerman <mpe@ellerman.id.au> Cc: Palmer Dabbelt <palmerdabbelt@google.com> Cc: Paul Mackerras <paulus@samba.org> Cc: Paul Walmsley <paul.walmsley@sifive.com> Cc: Rich Felker <dalias@libc.org> Cc: Russell King <linux@armlinux.org.uk> Cc: Thomas Bogendoerfer <tsbogend@alpha.franken.de> Cc: Vasily Gorbik <gor@linux.ibm.com> Cc: Yoshinori Sato <ysato@users.sourceforge.jp> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2021-05-04 18:38:09 -07:00
config ARCH_HAS_CACHE_LINE_SIZE
bool
usercopy: Check valid lifetime via stack depth One of the things that CONFIG_HARDENED_USERCOPY sanity-checks is whether an object that is about to be copied to/from userspace is overlapping the stack at all. If it is, it performs a number of inexpensive bounds checks. One of the finer-grained checks is whether an object crosses stack frames within the stack region. Doing this on x86 with CONFIG_FRAME_POINTER was cheap/easy. Doing it with ORC was deemed too heavy, and was left out (a while ago), leaving the courser whole-stack check. The LKDTM tests USERCOPY_STACK_FRAME_TO and USERCOPY_STACK_FRAME_FROM try to exercise these cross-frame cases to validate the defense is working. They have been failing ever since ORC was added (which was expected). While Muhammad was investigating various LKDTM failures[1], he asked me for additional details on them, and I realized that when exact stack frame boundary checking is not available (i.e. everything except x86 with FRAME_POINTER), it could check if a stack object is at least "current depth valid", in the sense that any object within the stack region but not between start-of-stack and current_stack_pointer should be considered unavailable (i.e. its lifetime is from a call no longer present on the stack). Introduce ARCH_HAS_CURRENT_STACK_POINTER to track which architectures have actually implemented the common global register alias. Additionally report usercopy bounds checking failures with an offset from current_stack_pointer, which may assist with diagnosing failures. The LKDTM USERCOPY_STACK_FRAME_TO and USERCOPY_STACK_FRAME_FROM tests (once slightly adjusted in a separate patch) pass again with this fixed. [1] https://github.com/kernelci/kernelci-project/issues/84 Cc: Matthew Wilcox (Oracle) <willy@infradead.org> Cc: Josh Poimboeuf <jpoimboe@redhat.com> Cc: Andrew Morton <akpm@linux-foundation.org> Cc: linux-mm@kvack.org Reported-by: Muhammad Usama Anjum <usama.anjum@collabora.com> Signed-off-by: Kees Cook <keescook@chromium.org> --- v1: https://lore.kernel.org/lkml/20220216201449.2087956-1-keescook@chromium.org v2: https://lore.kernel.org/lkml/20220224060342.1855457-1-keescook@chromium.org v3: https://lore.kernel.org/lkml/20220225173345.3358109-1-keescook@chromium.org v4: - improve commit log (akpm)
2022-02-16 13:05:28 -07:00
config ARCH_HAS_CURRENT_STACK_POINTER
bool
help
In support of HARDENED_USERCOPY performing stack variable lifetime
checking, an architecture-agnostic way to find the stack pointer
is needed. Once an architecture defines an unsigned long global
register alias named "current_stack_pointer", this config can be
selected.
config ARCH_HAS_PTE_DEVMAP
bool
config ARCH_HAS_ZONE_DMA_SET
bool
config ZONE_DMA
bool "Support DMA zone" if ARCH_HAS_ZONE_DMA_SET
default y if ARM64 || X86
config ZONE_DMA32
bool "Support DMA32 zone" if ARCH_HAS_ZONE_DMA_SET
depends on !X86_32
default y if ARM64
config ZONE_DEVICE
mm/ZONE_DEVICE: new type of ZONE_DEVICE for unaddressable memory HMM (heterogeneous memory management) need struct page to support migration from system main memory to device memory. Reasons for HMM and migration to device memory is explained with HMM core patch. This patch deals with device memory that is un-addressable memory (ie CPU can not access it). Hence we do not want those struct page to be manage like regular memory. That is why we extend ZONE_DEVICE to support different types of memory. A persistent memory type is define for existing user of ZONE_DEVICE and a new device un-addressable type is added for the un-addressable memory type. There is a clear separation between what is expected from each memory type and existing user of ZONE_DEVICE are un-affected by new requirement and new use of the un-addressable type. All specific code path are protect with test against the memory type. Because memory is un-addressable we use a new special swap type for when a page is migrated to device memory (this reduces the number of maximum swap file). The main two additions beside memory type to ZONE_DEVICE is two callbacks. First one, page_free() is call whenever page refcount reach 1 (which means the page is free as ZONE_DEVICE page never reach a refcount of 0). This allow device driver to manage its memory and associated struct page. The second callback page_fault() happens when there is a CPU access to an address that is back by a device page (which are un-addressable by the CPU). This callback is responsible to migrate the page back to system main memory. Device driver can not block migration back to system memory, HMM make sure that such page can not be pin into device memory. If device is in some error condition and can not migrate memory back then a CPU page fault to device memory should end with SIGBUS. [arnd@arndb.de: fix warning] Link: http://lkml.kernel.org/r/20170823133213.712917-1-arnd@arndb.de Link: http://lkml.kernel.org/r/20170817000548.32038-8-jglisse@redhat.com Signed-off-by: Jérôme Glisse <jglisse@redhat.com> Signed-off-by: Arnd Bergmann <arnd@arndb.de> Acked-by: Dan Williams <dan.j.williams@intel.com> Cc: Ross Zwisler <ross.zwisler@linux.intel.com> Cc: Aneesh Kumar <aneesh.kumar@linux.vnet.ibm.com> Cc: Balbir Singh <bsingharora@gmail.com> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: David Nellans <dnellans@nvidia.com> Cc: Evgeny Baskakov <ebaskakov@nvidia.com> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: John Hubbard <jhubbard@nvidia.com> Cc: Kirill A. Shutemov <kirill.shutemov@linux.intel.com> Cc: Mark Hairgrove <mhairgrove@nvidia.com> Cc: Michal Hocko <mhocko@kernel.org> Cc: Paul E. McKenney <paulmck@linux.vnet.ibm.com> Cc: Sherry Cheung <SCheung@nvidia.com> Cc: Subhash Gutti <sgutti@nvidia.com> Cc: Vladimir Davydov <vdavydov.dev@gmail.com> Cc: Bob Liu <liubo95@huawei.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2017-09-08 16:11:43 -07:00
bool "Device memory (pmem, HMM, etc...) hotplug support"
depends on MEMORY_HOTPLUG
depends on MEMORY_HOTREMOVE
depends on SPARSEMEM_VMEMMAP
depends on ARCH_HAS_PTE_DEVMAP
select XARRAY_MULTI
help
Device memory hotplug support allows for establishing pmem,
or other device driver discovered memory regions, in the
memmap. This allows pfn_to_page() lookups of otherwise
"device-physical" addresses which is needed for using a DAX
mapping in an O_DIRECT operation, among other things.
If FS_DAX is enabled, then say Y.
media updates for v4.3-rc1 -----BEGIN PGP SIGNATURE----- Version: GnuPG v1 iQIcBAABAgAGBQJV8WvjAAoJEAhfPr2O5OEV5wIP/AjmqOau99ms4FvOQ932sO57 kKDM4CYeTBkYY2Xz2eGStgxhcEj538JTf6SXdrceEEYJHb/GNCb2iBM1TnB4YciF rqhFv+n3R8h4Yn5KmhEhYzEfO7HUoyHPrOhcmTLzDoTO5wyrhAlPZxDWHohmfU84 uQ8WyGPYLxwm8hdZ+/NkB8PXsGbWN65EoKzN6tt2kA6HUP52UxE0Cw7Qu7Iu5zmO y/x03mMbjhCBFFE41EeM76J+xKBhuaS4cyf8g08DJy5Zpf6ic8bKFmVg1tAFOZRD mCETLrUlPYhglHqOoVS25bCI5kCw9xTAyjPZdQnwCTwgHl5gG3E4oJYKASrmZlps igMSmLJEpQilsLy1Ze+K+Ci8EILmZzwbi21X0sbjq74Jd+tJZ+C8ZuWHVmPEF9j7 iHtZNIRzkzufNBJZn3DsmlGBb/Xc/UqfZVnJAB9gu3Ktav6dmtEIHrGRPpL19iYH WtJWLt/Bpyb318K+fnxL8SzUqUxZJ4+8DrMtlgTqHmIRwVQ4CczyeWi0utQmBXEF CaNp00S2V9N1hn8OIc+gaf7LTYJn0LkHFsskoiUZ5aZQd9ai0ql0IT1xLe0r8lMi +ieB0Vp4wJtaodWIXOPeFugDqQXIb0Mh2M8J9FIJ116FLIai6btzO2iyVCtlR9Bg 1uPztCfJ/nusPPHnE26R =TEFw -----END PGP SIGNATURE----- Merge tag 'media/v4.3-2' of git://git.kernel.org/pub/scm/linux/kernel/git/mchehab/linux-media Pull media updates from Mauro Carvalho Chehab: "A series of patches that move part of the code used to allocate memory from the media subsystem to the mm subsystem" [ The mm parts have been acked by VM people, and the series was apparently in -mm for a while - Linus ] * tag 'media/v4.3-2' of git://git.kernel.org/pub/scm/linux/kernel/git/mchehab/linux-media: [media] drm/exynos: Convert g2d_userptr_get_dma_addr() to use get_vaddr_frames() [media] media: vb2: Remove unused functions [media] media: vb2: Convert vb2_dc_get_userptr() to use frame vector [media] media: vb2: Convert vb2_vmalloc_get_userptr() to use frame vector [media] media: vb2: Convert vb2_dma_sg_get_userptr() to use frame vector [media] vb2: Provide helpers for mapping virtual addresses [media] media: omap_vout: Convert omap_vout_uservirt_to_phys() to use get_vaddr_pfns() [media] mm: Provide new get_vaddr_frames() helper [media] vb2: Push mmap_sem down to memops
2015-09-11 16:42:39 -07:00
#
# Helpers to mirror range of the CPU page tables of a process into device page
# tables.
#
mm/hmm/mirror: mirror process address space on device with HMM helpers This is a heterogeneous memory management (HMM) process address space mirroring. In a nutshell this provide an API to mirror process address space on a device. This boils down to keeping CPU and device page table synchronize (we assume that both device and CPU are cache coherent like PCIe device can be). This patch provide a simple API for device driver to achieve address space mirroring thus avoiding each device driver to grow its own CPU page table walker and its own CPU page table synchronization mechanism. This is useful for NVidia GPU >= Pascal, Mellanox IB >= mlx5 and more hardware in the future. [jglisse@redhat.com: fix hmm for "mmu_notifier kill invalidate_page callback"] Link: http://lkml.kernel.org/r/20170830231955.GD9445@redhat.com Link: http://lkml.kernel.org/r/20170817000548.32038-4-jglisse@redhat.com Signed-off-by: Jérôme Glisse <jglisse@redhat.com> Signed-off-by: Evgeny Baskakov <ebaskakov@nvidia.com> Signed-off-by: John Hubbard <jhubbard@nvidia.com> Signed-off-by: Mark Hairgrove <mhairgrove@nvidia.com> Signed-off-by: Sherry Cheung <SCheung@nvidia.com> Signed-off-by: Subhash Gutti <sgutti@nvidia.com> Cc: Aneesh Kumar <aneesh.kumar@linux.vnet.ibm.com> Cc: Balbir Singh <bsingharora@gmail.com> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: Dan Williams <dan.j.williams@intel.com> Cc: David Nellans <dnellans@nvidia.com> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Kirill A. Shutemov <kirill.shutemov@linux.intel.com> Cc: Michal Hocko <mhocko@kernel.org> Cc: Paul E. McKenney <paulmck@linux.vnet.ibm.com> Cc: Ross Zwisler <ross.zwisler@linux.intel.com> Cc: Vladimir Davydov <vdavydov.dev@gmail.com> Cc: Bob Liu <liubo95@huawei.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2017-09-08 16:11:27 -07:00
config HMM_MIRROR
bool
depends on MMU
mm/hmm/mirror: mirror process address space on device with HMM helpers This is a heterogeneous memory management (HMM) process address space mirroring. In a nutshell this provide an API to mirror process address space on a device. This boils down to keeping CPU and device page table synchronize (we assume that both device and CPU are cache coherent like PCIe device can be). This patch provide a simple API for device driver to achieve address space mirroring thus avoiding each device driver to grow its own CPU page table walker and its own CPU page table synchronization mechanism. This is useful for NVidia GPU >= Pascal, Mellanox IB >= mlx5 and more hardware in the future. [jglisse@redhat.com: fix hmm for "mmu_notifier kill invalidate_page callback"] Link: http://lkml.kernel.org/r/20170830231955.GD9445@redhat.com Link: http://lkml.kernel.org/r/20170817000548.32038-4-jglisse@redhat.com Signed-off-by: Jérôme Glisse <jglisse@redhat.com> Signed-off-by: Evgeny Baskakov <ebaskakov@nvidia.com> Signed-off-by: John Hubbard <jhubbard@nvidia.com> Signed-off-by: Mark Hairgrove <mhairgrove@nvidia.com> Signed-off-by: Sherry Cheung <SCheung@nvidia.com> Signed-off-by: Subhash Gutti <sgutti@nvidia.com> Cc: Aneesh Kumar <aneesh.kumar@linux.vnet.ibm.com> Cc: Balbir Singh <bsingharora@gmail.com> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: Dan Williams <dan.j.williams@intel.com> Cc: David Nellans <dnellans@nvidia.com> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Kirill A. Shutemov <kirill.shutemov@linux.intel.com> Cc: Michal Hocko <mhocko@kernel.org> Cc: Paul E. McKenney <paulmck@linux.vnet.ibm.com> Cc: Ross Zwisler <ross.zwisler@linux.intel.com> Cc: Vladimir Davydov <vdavydov.dev@gmail.com> Cc: Bob Liu <liubo95@huawei.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2017-09-08 16:11:27 -07:00
resource: Introduce alloc_free_mem_region() The core of devm_request_free_mem_region() is a helper that searches for free space in iomem_resource and performs __request_region_locked() on the result of that search. The policy choices of the implementation conform to what CONFIG_DEVICE_PRIVATE users want which is memory that is immediately marked busy, and a preference to search for the first-fit free range in descending order from the top of the physical address space. CXL has a need for a similar allocator, but with the following tweaks: 1/ Search for free space in ascending order 2/ Search for free space relative to a given CXL window 3/ 'insert' rather than 'request' the new resource given downstream drivers from the CXL Region driver (like the pmem or dax drivers) are responsible for request_mem_region() when they activate the memory range. Rework __request_free_mem_region() into get_free_mem_region() which takes a set of GFR_* (Get Free Region) flags to control the allocation policy (ascending vs descending), and "busy" policy (insert_resource() vs request_region()). As part of the consolidation of the legacy GFR_REQUEST_REGION case with the new default of just inserting a new resource into the free space some minor cleanups like not checking for NULL before calling devres_free() (which does its own check) is included. Suggested-by: Jason Gunthorpe <jgg@nvidia.com> Link: https://lore.kernel.org/linux-cxl/20220420143406.GY2120790@nvidia.com/ Cc: Matthew Wilcox <willy@infradead.org> Cc: Christoph Hellwig <hch@lst.de> Reviewed-by: Jonathan Cameron <Jonathan.Cameron@huawei.com> Link: https://lore.kernel.org/r/165784333333.1758207.13703329337805274043.stgit@dwillia2-xfh.jf.intel.com Signed-off-by: Dan Williams <dan.j.williams@intel.com>
2022-05-20 13:41:24 -07:00
config GET_FREE_REGION
bool
mm/ZONE_DEVICE: new type of ZONE_DEVICE for unaddressable memory HMM (heterogeneous memory management) need struct page to support migration from system main memory to device memory. Reasons for HMM and migration to device memory is explained with HMM core patch. This patch deals with device memory that is un-addressable memory (ie CPU can not access it). Hence we do not want those struct page to be manage like regular memory. That is why we extend ZONE_DEVICE to support different types of memory. A persistent memory type is define for existing user of ZONE_DEVICE and a new device un-addressable type is added for the un-addressable memory type. There is a clear separation between what is expected from each memory type and existing user of ZONE_DEVICE are un-affected by new requirement and new use of the un-addressable type. All specific code path are protect with test against the memory type. Because memory is un-addressable we use a new special swap type for when a page is migrated to device memory (this reduces the number of maximum swap file). The main two additions beside memory type to ZONE_DEVICE is two callbacks. First one, page_free() is call whenever page refcount reach 1 (which means the page is free as ZONE_DEVICE page never reach a refcount of 0). This allow device driver to manage its memory and associated struct page. The second callback page_fault() happens when there is a CPU access to an address that is back by a device page (which are un-addressable by the CPU). This callback is responsible to migrate the page back to system main memory. Device driver can not block migration back to system memory, HMM make sure that such page can not be pin into device memory. If device is in some error condition and can not migrate memory back then a CPU page fault to device memory should end with SIGBUS. [arnd@arndb.de: fix warning] Link: http://lkml.kernel.org/r/20170823133213.712917-1-arnd@arndb.de Link: http://lkml.kernel.org/r/20170817000548.32038-8-jglisse@redhat.com Signed-off-by: Jérôme Glisse <jglisse@redhat.com> Signed-off-by: Arnd Bergmann <arnd@arndb.de> Acked-by: Dan Williams <dan.j.williams@intel.com> Cc: Ross Zwisler <ross.zwisler@linux.intel.com> Cc: Aneesh Kumar <aneesh.kumar@linux.vnet.ibm.com> Cc: Balbir Singh <bsingharora@gmail.com> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: David Nellans <dnellans@nvidia.com> Cc: Evgeny Baskakov <ebaskakov@nvidia.com> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: John Hubbard <jhubbard@nvidia.com> Cc: Kirill A. Shutemov <kirill.shutemov@linux.intel.com> Cc: Mark Hairgrove <mhairgrove@nvidia.com> Cc: Michal Hocko <mhocko@kernel.org> Cc: Paul E. McKenney <paulmck@linux.vnet.ibm.com> Cc: Sherry Cheung <SCheung@nvidia.com> Cc: Subhash Gutti <sgutti@nvidia.com> Cc: Vladimir Davydov <vdavydov.dev@gmail.com> Cc: Bob Liu <liubo95@huawei.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2017-09-08 16:11:43 -07:00
config DEVICE_PRIVATE
bool "Unaddressable device memory (GPU memory, ...)"
depends on ZONE_DEVICE
resource: Introduce alloc_free_mem_region() The core of devm_request_free_mem_region() is a helper that searches for free space in iomem_resource and performs __request_region_locked() on the result of that search. The policy choices of the implementation conform to what CONFIG_DEVICE_PRIVATE users want which is memory that is immediately marked busy, and a preference to search for the first-fit free range in descending order from the top of the physical address space. CXL has a need for a similar allocator, but with the following tweaks: 1/ Search for free space in ascending order 2/ Search for free space relative to a given CXL window 3/ 'insert' rather than 'request' the new resource given downstream drivers from the CXL Region driver (like the pmem or dax drivers) are responsible for request_mem_region() when they activate the memory range. Rework __request_free_mem_region() into get_free_mem_region() which takes a set of GFR_* (Get Free Region) flags to control the allocation policy (ascending vs descending), and "busy" policy (insert_resource() vs request_region()). As part of the consolidation of the legacy GFR_REQUEST_REGION case with the new default of just inserting a new resource into the free space some minor cleanups like not checking for NULL before calling devres_free() (which does its own check) is included. Suggested-by: Jason Gunthorpe <jgg@nvidia.com> Link: https://lore.kernel.org/linux-cxl/20220420143406.GY2120790@nvidia.com/ Cc: Matthew Wilcox <willy@infradead.org> Cc: Christoph Hellwig <hch@lst.de> Reviewed-by: Jonathan Cameron <Jonathan.Cameron@huawei.com> Link: https://lore.kernel.org/r/165784333333.1758207.13703329337805274043.stgit@dwillia2-xfh.jf.intel.com Signed-off-by: Dan Williams <dan.j.williams@intel.com>
2022-05-20 13:41:24 -07:00
select GET_FREE_REGION
mm/ZONE_DEVICE: new type of ZONE_DEVICE for unaddressable memory HMM (heterogeneous memory management) need struct page to support migration from system main memory to device memory. Reasons for HMM and migration to device memory is explained with HMM core patch. This patch deals with device memory that is un-addressable memory (ie CPU can not access it). Hence we do not want those struct page to be manage like regular memory. That is why we extend ZONE_DEVICE to support different types of memory. A persistent memory type is define for existing user of ZONE_DEVICE and a new device un-addressable type is added for the un-addressable memory type. There is a clear separation between what is expected from each memory type and existing user of ZONE_DEVICE are un-affected by new requirement and new use of the un-addressable type. All specific code path are protect with test against the memory type. Because memory is un-addressable we use a new special swap type for when a page is migrated to device memory (this reduces the number of maximum swap file). The main two additions beside memory type to ZONE_DEVICE is two callbacks. First one, page_free() is call whenever page refcount reach 1 (which means the page is free as ZONE_DEVICE page never reach a refcount of 0). This allow device driver to manage its memory and associated struct page. The second callback page_fault() happens when there is a CPU access to an address that is back by a device page (which are un-addressable by the CPU). This callback is responsible to migrate the page back to system main memory. Device driver can not block migration back to system memory, HMM make sure that such page can not be pin into device memory. If device is in some error condition and can not migrate memory back then a CPU page fault to device memory should end with SIGBUS. [arnd@arndb.de: fix warning] Link: http://lkml.kernel.org/r/20170823133213.712917-1-arnd@arndb.de Link: http://lkml.kernel.org/r/20170817000548.32038-8-jglisse@redhat.com Signed-off-by: Jérôme Glisse <jglisse@redhat.com> Signed-off-by: Arnd Bergmann <arnd@arndb.de> Acked-by: Dan Williams <dan.j.williams@intel.com> Cc: Ross Zwisler <ross.zwisler@linux.intel.com> Cc: Aneesh Kumar <aneesh.kumar@linux.vnet.ibm.com> Cc: Balbir Singh <bsingharora@gmail.com> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: David Nellans <dnellans@nvidia.com> Cc: Evgeny Baskakov <ebaskakov@nvidia.com> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: John Hubbard <jhubbard@nvidia.com> Cc: Kirill A. Shutemov <kirill.shutemov@linux.intel.com> Cc: Mark Hairgrove <mhairgrove@nvidia.com> Cc: Michal Hocko <mhocko@kernel.org> Cc: Paul E. McKenney <paulmck@linux.vnet.ibm.com> Cc: Sherry Cheung <SCheung@nvidia.com> Cc: Subhash Gutti <sgutti@nvidia.com> Cc: Vladimir Davydov <vdavydov.dev@gmail.com> Cc: Bob Liu <liubo95@huawei.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2017-09-08 16:11:43 -07:00
help
Allows creation of struct pages to represent unaddressable device
memory; i.e., memory that is only accessible from the device (or
group of devices). You likely also want to select HMM_MIRROR.
config VMAP_PFN
bool
mm/core, x86/mm/pkeys: Store protection bits in high VMA flags vma->vm_flags is an 'unsigned long', so has space for 32 flags on 32-bit architectures. The high 32 bits are unused on 64-bit platforms. We've steered away from using the unused high VMA bits for things because we would have difficulty supporting it on 32-bit. Protection Keys are not available in 32-bit mode, so there is no concern about supporting this feature in 32-bit mode or on 32-bit CPUs. This patch carves out 4 bits from the high half of vma->vm_flags and allows architectures to set config option to make them available. Sparse complains about these constants unless we explicitly call them "UL". Signed-off-by: Dave Hansen <dave.hansen@linux.intel.com> Reviewed-by: Thomas Gleixner <tglx@linutronix.de> Cc: Andrew Morton <akpm@linux-foundation.org> Cc: Andy Lutomirski <luto@amacapital.net> Cc: Borislav Petkov <bp@alien8.de> Cc: Brian Gerst <brgerst@gmail.com> Cc: Dan Williams <dan.j.williams@intel.com> Cc: Dave Hansen <dave@sr71.net> Cc: Denys Vlasenko <dvlasenk@redhat.com> Cc: H. Peter Anvin <hpa@zytor.com> Cc: Jan Kara <jack@suse.cz> Cc: Kirill A. Shutemov <kirill.shutemov@linux.intel.com> Cc: Konstantin Khlebnikov <koct9i@gmail.com> Cc: Linus Torvalds <torvalds@linux-foundation.org> Cc: Mel Gorman <mgorman@suse.de> Cc: Michal Hocko <mhocko@suse.com> Cc: Oleg Nesterov <oleg@redhat.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Rik van Riel <riel@redhat.com> Cc: Sasha Levin <sasha.levin@oracle.com> Cc: Valentin Rothberg <valentinrothberg@gmail.com> Cc: Vladimir Davydov <vdavydov@parallels.com> Cc: Vlastimil Babka <vbabka@suse.cz> Cc: Xie XiuQi <xiexiuqi@huawei.com> Cc: linux-kernel@vger.kernel.org Cc: linux-mm@kvack.org Link: http://lkml.kernel.org/r/20160212210208.81AF00D5@viggo.jf.intel.com Signed-off-by: Ingo Molnar <mingo@kernel.org>
2016-02-12 14:02:08 -07:00
config ARCH_USES_HIGH_VMA_FLAGS
bool
config ARCH_HAS_PKEYS
bool
config ARCH_USES_PG_ARCH_2
bool
config ARCH_USES_PG_ARCH_3
bool
config VM_EVENT_COUNTERS
default y
bool "Enable VM event counters for /proc/vmstat" if EXPERT
help
VM event counters are needed for event counts to be shown.
This option allows the disabling of the VM event counters
on EXPERT systems. /proc/vmstat will only show page counts
if VM event counters are disabled.
config PERCPU_STATS
bool "Collect percpu memory statistics"
help
This feature collects and exposes statistics via debugfs. The
information includes global and per chunk statistics, which can
be used to help understand percpu memory usage.
mm/gup_benchmark: rename to mm/gup_test Patch series "selftests/vm: gup_test, hmm-tests, assorted improvements", v3. Summary: This series provides two main things, and a number of smaller supporting goodies. The two main points are: 1) Add a new sub-test to gup_test, which in turn is a renamed version of gup_benchmark. This sub-test allows nicer testing of dump_pages(), at least on user-space pages. For quite a while, I was doing a quick hack to gup_test.c whenever I wanted to try out changes to dump_page(). Then Matthew Wilcox asked me what I meant when I said "I used my dump_page() unit test", and I realized that it might be nice to check in a polished up version of that. Details about how it works and how to use it are in the commit description for patch #6 ("selftests/vm: gup_test: introduce the dump_pages() sub-test"). 2) Fixes a limitation of hmm-tests: these tests are incredibly useful, but only if people actually build and run them. And it turns out that libhugetlbfs is a little too effective at throwing a wrench in the works, there. So I've added a little configuration check that removes just two of the 21 hmm-tests, if libhugetlbfs is not available. Further details in the commit description of patch #8 ("selftests/vm: hmm-tests: remove the libhugetlbfs dependency"). Other smaller things that this series does: a) Remove code duplication by creating gup_test.h. b) Clear up the sub-test organization, and their invocation within run_vmtests.sh. c) Other minor assorted improvements. [1] v2 is here: https://lore.kernel.org/linux-doc/20200929212747.251804-1-jhubbard@nvidia.com/ [2] https://lore.kernel.org/r/CAHk-=wgh-TMPHLY3jueHX7Y2fWh3D+nMBqVS__AZm6-oorquWA@mail.gmail.com This patch (of 9): Rename nearly every "gup_benchmark" reference and file name to "gup_test". The one exception is for the actual gup benchmark test itself. The current code already does a *little* bit more than benchmarking, and definitely covers more than get_user_pages_fast(). More importantly, however, subsequent patches are about to add some functionality that is non-benchmark related. Closely related changes: * Kconfig: in addition to renaming the options from GUP_BENCHMARK to GUP_TEST, update the help text to reflect that it's no longer a benchmark-only test. Link: https://lkml.kernel.org/r/20201026064021.3545418-1-jhubbard@nvidia.com Link: https://lkml.kernel.org/r/20201026064021.3545418-2-jhubbard@nvidia.com Signed-off-by: John Hubbard <jhubbard@nvidia.com> Cc: Jonathan Corbet <corbet@lwn.net> Cc: Jérôme Glisse <jglisse@redhat.com> Cc: Ralph Campbell <rcampbell@nvidia.com> Cc: Shuah Khan <shuah@kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2020-12-14 20:05:05 -07:00
config GUP_TEST
bool "Enable infrastructure for get_user_pages()-related unit tests"
depends on DEBUG_FS
help
mm/gup_benchmark: rename to mm/gup_test Patch series "selftests/vm: gup_test, hmm-tests, assorted improvements", v3. Summary: This series provides two main things, and a number of smaller supporting goodies. The two main points are: 1) Add a new sub-test to gup_test, which in turn is a renamed version of gup_benchmark. This sub-test allows nicer testing of dump_pages(), at least on user-space pages. For quite a while, I was doing a quick hack to gup_test.c whenever I wanted to try out changes to dump_page(). Then Matthew Wilcox asked me what I meant when I said "I used my dump_page() unit test", and I realized that it might be nice to check in a polished up version of that. Details about how it works and how to use it are in the commit description for patch #6 ("selftests/vm: gup_test: introduce the dump_pages() sub-test"). 2) Fixes a limitation of hmm-tests: these tests are incredibly useful, but only if people actually build and run them. And it turns out that libhugetlbfs is a little too effective at throwing a wrench in the works, there. So I've added a little configuration check that removes just two of the 21 hmm-tests, if libhugetlbfs is not available. Further details in the commit description of patch #8 ("selftests/vm: hmm-tests: remove the libhugetlbfs dependency"). Other smaller things that this series does: a) Remove code duplication by creating gup_test.h. b) Clear up the sub-test organization, and their invocation within run_vmtests.sh. c) Other minor assorted improvements. [1] v2 is here: https://lore.kernel.org/linux-doc/20200929212747.251804-1-jhubbard@nvidia.com/ [2] https://lore.kernel.org/r/CAHk-=wgh-TMPHLY3jueHX7Y2fWh3D+nMBqVS__AZm6-oorquWA@mail.gmail.com This patch (of 9): Rename nearly every "gup_benchmark" reference and file name to "gup_test". The one exception is for the actual gup benchmark test itself. The current code already does a *little* bit more than benchmarking, and definitely covers more than get_user_pages_fast(). More importantly, however, subsequent patches are about to add some functionality that is non-benchmark related. Closely related changes: * Kconfig: in addition to renaming the options from GUP_BENCHMARK to GUP_TEST, update the help text to reflect that it's no longer a benchmark-only test. Link: https://lkml.kernel.org/r/20201026064021.3545418-1-jhubbard@nvidia.com Link: https://lkml.kernel.org/r/20201026064021.3545418-2-jhubbard@nvidia.com Signed-off-by: John Hubbard <jhubbard@nvidia.com> Cc: Jonathan Corbet <corbet@lwn.net> Cc: Jérôme Glisse <jglisse@redhat.com> Cc: Ralph Campbell <rcampbell@nvidia.com> Cc: Shuah Khan <shuah@kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2020-12-14 20:05:05 -07:00
Provides /sys/kernel/debug/gup_test, which in turn provides a way
to make ioctl calls that can launch kernel-based unit tests for
the get_user_pages*() and pin_user_pages*() family of API calls.
mm/gup_benchmark: rename to mm/gup_test Patch series "selftests/vm: gup_test, hmm-tests, assorted improvements", v3. Summary: This series provides two main things, and a number of smaller supporting goodies. The two main points are: 1) Add a new sub-test to gup_test, which in turn is a renamed version of gup_benchmark. This sub-test allows nicer testing of dump_pages(), at least on user-space pages. For quite a while, I was doing a quick hack to gup_test.c whenever I wanted to try out changes to dump_page(). Then Matthew Wilcox asked me what I meant when I said "I used my dump_page() unit test", and I realized that it might be nice to check in a polished up version of that. Details about how it works and how to use it are in the commit description for patch #6 ("selftests/vm: gup_test: introduce the dump_pages() sub-test"). 2) Fixes a limitation of hmm-tests: these tests are incredibly useful, but only if people actually build and run them. And it turns out that libhugetlbfs is a little too effective at throwing a wrench in the works, there. So I've added a little configuration check that removes just two of the 21 hmm-tests, if libhugetlbfs is not available. Further details in the commit description of patch #8 ("selftests/vm: hmm-tests: remove the libhugetlbfs dependency"). Other smaller things that this series does: a) Remove code duplication by creating gup_test.h. b) Clear up the sub-test organization, and their invocation within run_vmtests.sh. c) Other minor assorted improvements. [1] v2 is here: https://lore.kernel.org/linux-doc/20200929212747.251804-1-jhubbard@nvidia.com/ [2] https://lore.kernel.org/r/CAHk-=wgh-TMPHLY3jueHX7Y2fWh3D+nMBqVS__AZm6-oorquWA@mail.gmail.com This patch (of 9): Rename nearly every "gup_benchmark" reference and file name to "gup_test". The one exception is for the actual gup benchmark test itself. The current code already does a *little* bit more than benchmarking, and definitely covers more than get_user_pages_fast(). More importantly, however, subsequent patches are about to add some functionality that is non-benchmark related. Closely related changes: * Kconfig: in addition to renaming the options from GUP_BENCHMARK to GUP_TEST, update the help text to reflect that it's no longer a benchmark-only test. Link: https://lkml.kernel.org/r/20201026064021.3545418-1-jhubbard@nvidia.com Link: https://lkml.kernel.org/r/20201026064021.3545418-2-jhubbard@nvidia.com Signed-off-by: John Hubbard <jhubbard@nvidia.com> Cc: Jonathan Corbet <corbet@lwn.net> Cc: Jérôme Glisse <jglisse@redhat.com> Cc: Ralph Campbell <rcampbell@nvidia.com> Cc: Shuah Khan <shuah@kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2020-12-14 20:05:05 -07:00
These tests include benchmark testing of the _fast variants of
get_user_pages*() and pin_user_pages*(), as well as smoke tests of
the non-_fast variants.
selftests/vm: gup_test: introduce the dump_pages() sub-test For quite a while, I was doing a quick hack to gup_test.c (previously, gup_benchmark.c) whenever I wanted to try out my changes to dump_page(). This makes that hack unnecessary, and instead allows anyone to easily get the same coverage from a user space program. That saves a lot of time because you don't have to change the kernel, in order to test different pages and options. The new sub-test takes advantage of the existing gup_test infrastructure, which already provides a simple user space program, some allocated user space pages, an ioctl call, pinning of those pages (via either get_user_pages or pin_user_pages) and a corresponding kernel-side test invocation. There's not much more required, mainly just a couple of inputs from the user. In fact, the new test re-uses the existing command line options in order to get various helpful combinations (THP or normal, _fast or slow gup, gup vs. pup, and more). New command line options are: which pages to dump, and what type of "get/pin" to use. In order to figure out which pages to dump, the logic is: * If the user doesn't specify anything, the page 0 (the first page in the address range that the program sets up for testing) is dumped. * Or, the user can type up to 8 page indices anywhere on the command line. If you type more than 8, then it uses the first 8 and ignores the remaining items. For example: ./gup_test -ct -F 1 0 19 0x1000 Meaning: -c: dump pages sub-test -t: use THP pages -F 1: use pin_user_pages() instead of get_user_pages() 0 19 0x1000: dump pages 0, 19, and 4096 Link: https://lkml.kernel.org/r/20201026064021.3545418-7-jhubbard@nvidia.com Signed-off-by: John Hubbard <jhubbard@nvidia.com> Cc: Jérôme Glisse <jglisse@redhat.com> Cc: Jonathan Corbet <corbet@lwn.net> Cc: Ralph Campbell <rcampbell@nvidia.com> Cc: Shuah Khan <shuah@kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2020-12-14 20:05:21 -07:00
There is also a sub-test that allows running dump_page() on any
of up to eight pages (selected by command line args) within the
range of user-space addresses. These pages are either pinned via
pin_user_pages*(), or pinned via get_user_pages*(), as specified
by other command line arguments.
See tools/testing/selftests/mm/gup_test.c
mm: introduce ARCH_HAS_PTE_SPECIAL Currently the PTE special supports is turned on in per architecture header files. Most of the time, it is defined in arch/*/include/asm/pgtable.h depending or not on some other per architecture static definition. This patch introduce a new configuration variable to manage this directly in the Kconfig files. It would later replace __HAVE_ARCH_PTE_SPECIAL. Here notes for some architecture where the definition of __HAVE_ARCH_PTE_SPECIAL is not obvious: arm __HAVE_ARCH_PTE_SPECIAL which is currently defined in arch/arm/include/asm/pgtable-3level.h which is included by arch/arm/include/asm/pgtable.h when CONFIG_ARM_LPAE is set. So select ARCH_HAS_PTE_SPECIAL if ARM_LPAE. powerpc __HAVE_ARCH_PTE_SPECIAL is defined in 2 files: - arch/powerpc/include/asm/book3s/64/pgtable.h - arch/powerpc/include/asm/pte-common.h The first one is included if (PPC_BOOK3S & PPC64) while the second is included in all the other cases. So select ARCH_HAS_PTE_SPECIAL all the time. sparc: __HAVE_ARCH_PTE_SPECIAL is defined if defined(__sparc__) && defined(__arch64__) which are defined through the compiler in sparc/Makefile if !SPARC32 which I assume to be if SPARC64. So select ARCH_HAS_PTE_SPECIAL if SPARC64 There is no functional change introduced by this patch. Link: http://lkml.kernel.org/r/1523433816-14460-2-git-send-email-ldufour@linux.vnet.ibm.com Signed-off-by: Laurent Dufour <ldufour@linux.vnet.ibm.com> Suggested-by: Jerome Glisse <jglisse@redhat.com> Reviewed-by: Jerome Glisse <jglisse@redhat.com> Acked-by: David Rientjes <rientjes@google.com> Cc: Michal Hocko <mhocko@kernel.org> Cc: "Aneesh Kumar K . V" <aneesh.kumar@linux.vnet.ibm.com> Cc: Michael Ellerman <mpe@ellerman.id.au> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: Paul Mackerras <paulus@samba.org> Cc: Jonathan Corbet <corbet@lwn.net> Cc: Catalin Marinas <catalin.marinas@arm.com> Cc: Will Deacon <will.deacon@arm.com> Cc: Yoshinori Sato <ysato@users.sourceforge.jp> Cc: Rich Felker <dalias@libc.org> Cc: David S. Miller <davem@davemloft.net> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Ingo Molnar <mingo@redhat.com> Cc: Vineet Gupta <vgupta@synopsys.com> Cc: Palmer Dabbelt <palmer@sifive.com> Cc: Albert Ou <albert@sifive.com> Cc: Martin Schwidefsky <schwidefsky@de.ibm.com> Cc: Heiko Carstens <heiko.carstens@de.ibm.com> Cc: David Rientjes <rientjes@google.com> Cc: Robin Murphy <robin.murphy@arm.com> Cc: Christophe LEROY <christophe.leroy@c-s.fr> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2018-06-07 17:06:08 -07:00
comment "GUP_TEST needs to have DEBUG_FS enabled"
depends on !GUP_TEST && !DEBUG_FS
mm: introduce ARCH_HAS_PTE_SPECIAL Currently the PTE special supports is turned on in per architecture header files. Most of the time, it is defined in arch/*/include/asm/pgtable.h depending or not on some other per architecture static definition. This patch introduce a new configuration variable to manage this directly in the Kconfig files. It would later replace __HAVE_ARCH_PTE_SPECIAL. Here notes for some architecture where the definition of __HAVE_ARCH_PTE_SPECIAL is not obvious: arm __HAVE_ARCH_PTE_SPECIAL which is currently defined in arch/arm/include/asm/pgtable-3level.h which is included by arch/arm/include/asm/pgtable.h when CONFIG_ARM_LPAE is set. So select ARCH_HAS_PTE_SPECIAL if ARM_LPAE. powerpc __HAVE_ARCH_PTE_SPECIAL is defined in 2 files: - arch/powerpc/include/asm/book3s/64/pgtable.h - arch/powerpc/include/asm/pte-common.h The first one is included if (PPC_BOOK3S & PPC64) while the second is included in all the other cases. So select ARCH_HAS_PTE_SPECIAL all the time. sparc: __HAVE_ARCH_PTE_SPECIAL is defined if defined(__sparc__) && defined(__arch64__) which are defined through the compiler in sparc/Makefile if !SPARC32 which I assume to be if SPARC64. So select ARCH_HAS_PTE_SPECIAL if SPARC64 There is no functional change introduced by this patch. Link: http://lkml.kernel.org/r/1523433816-14460-2-git-send-email-ldufour@linux.vnet.ibm.com Signed-off-by: Laurent Dufour <ldufour@linux.vnet.ibm.com> Suggested-by: Jerome Glisse <jglisse@redhat.com> Reviewed-by: Jerome Glisse <jglisse@redhat.com> Acked-by: David Rientjes <rientjes@google.com> Cc: Michal Hocko <mhocko@kernel.org> Cc: "Aneesh Kumar K . V" <aneesh.kumar@linux.vnet.ibm.com> Cc: Michael Ellerman <mpe@ellerman.id.au> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: Paul Mackerras <paulus@samba.org> Cc: Jonathan Corbet <corbet@lwn.net> Cc: Catalin Marinas <catalin.marinas@arm.com> Cc: Will Deacon <will.deacon@arm.com> Cc: Yoshinori Sato <ysato@users.sourceforge.jp> Cc: Rich Felker <dalias@libc.org> Cc: David S. Miller <davem@davemloft.net> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Ingo Molnar <mingo@redhat.com> Cc: Vineet Gupta <vgupta@synopsys.com> Cc: Palmer Dabbelt <palmer@sifive.com> Cc: Albert Ou <albert@sifive.com> Cc: Martin Schwidefsky <schwidefsky@de.ibm.com> Cc: Heiko Carstens <heiko.carstens@de.ibm.com> Cc: David Rientjes <rientjes@google.com> Cc: Robin Murphy <robin.murphy@arm.com> Cc: Christophe LEROY <christophe.leroy@c-s.fr> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2018-06-07 17:06:08 -07:00
config GUP_GET_PXX_LOW_HIGH
bool
config DMAPOOL_TEST
tristate "Enable a module to run time tests on dma_pool"
depends on HAS_DMA
help
Provides a test module that will allocate and free many blocks of
various sizes and report how long it takes. This is intended to
provide a consistent way to measure how changes to the
dma_pool_alloc/free routines affect performance.
mm: introduce ARCH_HAS_PTE_SPECIAL Currently the PTE special supports is turned on in per architecture header files. Most of the time, it is defined in arch/*/include/asm/pgtable.h depending or not on some other per architecture static definition. This patch introduce a new configuration variable to manage this directly in the Kconfig files. It would later replace __HAVE_ARCH_PTE_SPECIAL. Here notes for some architecture where the definition of __HAVE_ARCH_PTE_SPECIAL is not obvious: arm __HAVE_ARCH_PTE_SPECIAL which is currently defined in arch/arm/include/asm/pgtable-3level.h which is included by arch/arm/include/asm/pgtable.h when CONFIG_ARM_LPAE is set. So select ARCH_HAS_PTE_SPECIAL if ARM_LPAE. powerpc __HAVE_ARCH_PTE_SPECIAL is defined in 2 files: - arch/powerpc/include/asm/book3s/64/pgtable.h - arch/powerpc/include/asm/pte-common.h The first one is included if (PPC_BOOK3S & PPC64) while the second is included in all the other cases. So select ARCH_HAS_PTE_SPECIAL all the time. sparc: __HAVE_ARCH_PTE_SPECIAL is defined if defined(__sparc__) && defined(__arch64__) which are defined through the compiler in sparc/Makefile if !SPARC32 which I assume to be if SPARC64. So select ARCH_HAS_PTE_SPECIAL if SPARC64 There is no functional change introduced by this patch. Link: http://lkml.kernel.org/r/1523433816-14460-2-git-send-email-ldufour@linux.vnet.ibm.com Signed-off-by: Laurent Dufour <ldufour@linux.vnet.ibm.com> Suggested-by: Jerome Glisse <jglisse@redhat.com> Reviewed-by: Jerome Glisse <jglisse@redhat.com> Acked-by: David Rientjes <rientjes@google.com> Cc: Michal Hocko <mhocko@kernel.org> Cc: "Aneesh Kumar K . V" <aneesh.kumar@linux.vnet.ibm.com> Cc: Michael Ellerman <mpe@ellerman.id.au> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: Paul Mackerras <paulus@samba.org> Cc: Jonathan Corbet <corbet@lwn.net> Cc: Catalin Marinas <catalin.marinas@arm.com> Cc: Will Deacon <will.deacon@arm.com> Cc: Yoshinori Sato <ysato@users.sourceforge.jp> Cc: Rich Felker <dalias@libc.org> Cc: David S. Miller <davem@davemloft.net> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Ingo Molnar <mingo@redhat.com> Cc: Vineet Gupta <vgupta@synopsys.com> Cc: Palmer Dabbelt <palmer@sifive.com> Cc: Albert Ou <albert@sifive.com> Cc: Martin Schwidefsky <schwidefsky@de.ibm.com> Cc: Heiko Carstens <heiko.carstens@de.ibm.com> Cc: David Rientjes <rientjes@google.com> Cc: Robin Murphy <robin.murphy@arm.com> Cc: Christophe LEROY <christophe.leroy@c-s.fr> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2018-06-07 17:06:08 -07:00
config ARCH_HAS_PTE_SPECIAL
bool
config MAPPING_DIRTY_HELPERS
bool
config KMAP_LOCAL
bool
kmap_local: don't assume kmap PTEs are linear arrays in memory The kmap_local conversion broke the ARM architecture, because the new code assumes that all PTEs used for creating kmaps form a linear array in memory, and uses array indexing to look up the kmap PTE belonging to a certain kmap index. On ARM, this cannot work, not only because the PTE pages may be non-adjacent in memory, but also because ARM/!LPAE interleaves hardware entries and extended entries (carrying software-only bits) in a way that is not compatible with array indexing. Fortunately, this only seems to affect configurations with more than 8 CPUs, due to the way the per-CPU kmap slots are organized in memory. Work around this by permitting an architecture to set a Kconfig symbol that signifies that the kmap PTEs do not form a lineary array in memory, and so the only way to locate the appropriate one is to walk the page tables. Link: https://lore.kernel.org/linux-arm-kernel/20211026131249.3731275-1-ardb@kernel.org/ Link: https://lkml.kernel.org/r/20211116094737.7391-1-ardb@kernel.org Fixes: 2a15ba82fa6c ("ARM: highmem: Switch to generic kmap atomic") Signed-off-by: Ard Biesheuvel <ardb@kernel.org> Reported-by: Quanyang Wang <quanyang.wang@windriver.com> Reviewed-by: Linus Walleij <linus.walleij@linaro.org> Acked-by: Russell King (Oracle) <rmk+kernel@armlinux.org.uk> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: <stable@vger.kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2021-11-19 17:43:55 -07:00
config KMAP_LOCAL_NON_LINEAR_PTE_ARRAY
bool
# struct io_mapping based helper. Selected by drivers that need them
config IO_MAPPING
bool
mm: introduce memfd_secret system call to create "secret" memory areas Introduce "memfd_secret" system call with the ability to create memory areas visible only in the context of the owning process and not mapped not only to other processes but in the kernel page tables as well. The secretmem feature is off by default and the user must explicitly enable it at the boot time. Once secretmem is enabled, the user will be able to create a file descriptor using the memfd_secret() system call. The memory areas created by mmap() calls from this file descriptor will be unmapped from the kernel direct map and they will be only mapped in the page table of the processes that have access to the file descriptor. Secretmem is designed to provide the following protections: * Enhanced protection (in conjunction with all the other in-kernel attack prevention systems) against ROP attacks. Seceretmem makes "simple" ROP insufficient to perform exfiltration, which increases the required complexity of the attack. Along with other protections like the kernel stack size limit and address space layout randomization which make finding gadgets is really hard, absence of any in-kernel primitive for accessing secret memory means the one gadget ROP attack can't work. Since the only way to access secret memory is to reconstruct the missing mapping entry, the attacker has to recover the physical page and insert a PTE pointing to it in the kernel and then retrieve the contents. That takes at least three gadgets which is a level of difficulty beyond most standard attacks. * Prevent cross-process secret userspace memory exposures. Once the secret memory is allocated, the user can't accidentally pass it into the kernel to be transmitted somewhere. The secreremem pages cannot be accessed via the direct map and they are disallowed in GUP. * Harden against exploited kernel flaws. In order to access secretmem, a kernel-side attack would need to either walk the page tables and create new ones, or spawn a new privileged uiserspace process to perform secrets exfiltration using ptrace. The file descriptor based memory has several advantages over the "traditional" mm interfaces, such as mlock(), mprotect(), madvise(). File descriptor approach allows explicit and controlled sharing of the memory areas, it allows to seal the operations. Besides, file descriptor based memory paves the way for VMMs to remove the secret memory range from the userspace hipervisor process, for instance QEMU. Andy Lutomirski says: "Getting fd-backed memory into a guest will take some possibly major work in the kernel, but getting vma-backed memory into a guest without mapping it in the host user address space seems much, much worse." memfd_secret() is made a dedicated system call rather than an extension to memfd_create() because it's purpose is to allow the user to create more secure memory mappings rather than to simply allow file based access to the memory. Nowadays a new system call cost is negligible while it is way simpler for userspace to deal with a clear-cut system calls than with a multiplexer or an overloaded syscall. Moreover, the initial implementation of memfd_secret() is completely distinct from memfd_create() so there is no much sense in overloading memfd_create() to begin with. If there will be a need for code sharing between these implementation it can be easily achieved without a need to adjust user visible APIs. The secret memory remains accessible in the process context using uaccess primitives, but it is not exposed to the kernel otherwise; secret memory areas are removed from the direct map and functions in the follow_page()/get_user_page() family will refuse to return a page that belongs to the secret memory area. Once there will be a use case that will require exposing secretmem to the kernel it will be an opt-in request in the system call flags so that user would have to decide what data can be exposed to the kernel. Removing of the pages from the direct map may cause its fragmentation on architectures that use large pages to map the physical memory which affects the system performance. However, the original Kconfig text for CONFIG_DIRECT_GBPAGES said that gigabyte pages in the direct map "... can improve the kernel's performance a tiny bit ..." (commit 00d1c5e05736 ("x86: add gbpages switches")) and the recent report [1] showed that "... although 1G mappings are a good default choice, there is no compelling evidence that it must be the only choice". Hence, it is sufficient to have secretmem disabled by default with the ability of a system administrator to enable it at boot time. Pages in the secretmem regions are unevictable and unmovable to avoid accidental exposure of the sensitive data via swap or during page migration. Since the secretmem mappings are locked in memory they cannot exceed RLIMIT_MEMLOCK. Since these mappings are already locked independently from mlock(), an attempt to mlock()/munlock() secretmem range would fail and mlockall()/munlockall() will ignore secretmem mappings. However, unlike mlock()ed memory, secretmem currently behaves more like long-term GUP: secretmem mappings are unmovable mappings directly consumed by user space. With default limits, there is no excessive use of secretmem and it poses no real problem in combination with ZONE_MOVABLE/CMA, but in the future this should be addressed to allow balanced use of large amounts of secretmem along with ZONE_MOVABLE/CMA. A page that was a part of the secret memory area is cleared when it is freed to ensure the data is not exposed to the next user of that page. The following example demonstrates creation of a secret mapping (error handling is omitted): fd = memfd_secret(0); ftruncate(fd, MAP_SIZE); ptr = mmap(NULL, MAP_SIZE, PROT_READ | PROT_WRITE, MAP_SHARED, fd, 0); [1] https://lore.kernel.org/linux-mm/213b4567-46ce-f116-9cdf-bbd0c884eb3c@linux.intel.com/ [akpm@linux-foundation.org: suppress Kconfig whine] Link: https://lkml.kernel.org/r/20210518072034.31572-5-rppt@kernel.org Signed-off-by: Mike Rapoport <rppt@linux.ibm.com> Acked-by: Hagen Paul Pfeifer <hagen@jauu.net> Acked-by: James Bottomley <James.Bottomley@HansenPartnership.com> Cc: Alexander Viro <viro@zeniv.linux.org.uk> Cc: Andy Lutomirski <luto@kernel.org> Cc: Arnd Bergmann <arnd@arndb.de> Cc: Borislav Petkov <bp@alien8.de> Cc: Catalin Marinas <catalin.marinas@arm.com> Cc: Christopher Lameter <cl@linux.com> Cc: Dan Williams <dan.j.williams@intel.com> Cc: Dave Hansen <dave.hansen@linux.intel.com> Cc: Elena Reshetova <elena.reshetova@intel.com> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Ingo Molnar <mingo@redhat.com> Cc: James Bottomley <jejb@linux.ibm.com> Cc: "Kirill A. Shutemov" <kirill@shutemov.name> Cc: Matthew Wilcox <willy@infradead.org> Cc: Mark Rutland <mark.rutland@arm.com> Cc: Michael Kerrisk <mtk.manpages@gmail.com> Cc: Palmer Dabbelt <palmer@dabbelt.com> Cc: Palmer Dabbelt <palmerdabbelt@google.com> Cc: Paul Walmsley <paul.walmsley@sifive.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Rick Edgecombe <rick.p.edgecombe@intel.com> Cc: Roman Gushchin <guro@fb.com> Cc: Shakeel Butt <shakeelb@google.com> Cc: Shuah Khan <shuah@kernel.org> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Tycho Andersen <tycho@tycho.ws> Cc: Will Deacon <will@kernel.org> Cc: David Hildenbrand <david@redhat.com> Cc: kernel test robot <lkp@intel.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2021-07-07 18:08:03 -07:00
config MEMFD_CREATE
bool "Enable memfd_create() system call" if EXPERT
mm: introduce memfd_secret system call to create "secret" memory areas Introduce "memfd_secret" system call with the ability to create memory areas visible only in the context of the owning process and not mapped not only to other processes but in the kernel page tables as well. The secretmem feature is off by default and the user must explicitly enable it at the boot time. Once secretmem is enabled, the user will be able to create a file descriptor using the memfd_secret() system call. The memory areas created by mmap() calls from this file descriptor will be unmapped from the kernel direct map and they will be only mapped in the page table of the processes that have access to the file descriptor. Secretmem is designed to provide the following protections: * Enhanced protection (in conjunction with all the other in-kernel attack prevention systems) against ROP attacks. Seceretmem makes "simple" ROP insufficient to perform exfiltration, which increases the required complexity of the attack. Along with other protections like the kernel stack size limit and address space layout randomization which make finding gadgets is really hard, absence of any in-kernel primitive for accessing secret memory means the one gadget ROP attack can't work. Since the only way to access secret memory is to reconstruct the missing mapping entry, the attacker has to recover the physical page and insert a PTE pointing to it in the kernel and then retrieve the contents. That takes at least three gadgets which is a level of difficulty beyond most standard attacks. * Prevent cross-process secret userspace memory exposures. Once the secret memory is allocated, the user can't accidentally pass it into the kernel to be transmitted somewhere. The secreremem pages cannot be accessed via the direct map and they are disallowed in GUP. * Harden against exploited kernel flaws. In order to access secretmem, a kernel-side attack would need to either walk the page tables and create new ones, or spawn a new privileged uiserspace process to perform secrets exfiltration using ptrace. The file descriptor based memory has several advantages over the "traditional" mm interfaces, such as mlock(), mprotect(), madvise(). File descriptor approach allows explicit and controlled sharing of the memory areas, it allows to seal the operations. Besides, file descriptor based memory paves the way for VMMs to remove the secret memory range from the userspace hipervisor process, for instance QEMU. Andy Lutomirski says: "Getting fd-backed memory into a guest will take some possibly major work in the kernel, but getting vma-backed memory into a guest without mapping it in the host user address space seems much, much worse." memfd_secret() is made a dedicated system call rather than an extension to memfd_create() because it's purpose is to allow the user to create more secure memory mappings rather than to simply allow file based access to the memory. Nowadays a new system call cost is negligible while it is way simpler for userspace to deal with a clear-cut system calls than with a multiplexer or an overloaded syscall. Moreover, the initial implementation of memfd_secret() is completely distinct from memfd_create() so there is no much sense in overloading memfd_create() to begin with. If there will be a need for code sharing between these implementation it can be easily achieved without a need to adjust user visible APIs. The secret memory remains accessible in the process context using uaccess primitives, but it is not exposed to the kernel otherwise; secret memory areas are removed from the direct map and functions in the follow_page()/get_user_page() family will refuse to return a page that belongs to the secret memory area. Once there will be a use case that will require exposing secretmem to the kernel it will be an opt-in request in the system call flags so that user would have to decide what data can be exposed to the kernel. Removing of the pages from the direct map may cause its fragmentation on architectures that use large pages to map the physical memory which affects the system performance. However, the original Kconfig text for CONFIG_DIRECT_GBPAGES said that gigabyte pages in the direct map "... can improve the kernel's performance a tiny bit ..." (commit 00d1c5e05736 ("x86: add gbpages switches")) and the recent report [1] showed that "... although 1G mappings are a good default choice, there is no compelling evidence that it must be the only choice". Hence, it is sufficient to have secretmem disabled by default with the ability of a system administrator to enable it at boot time. Pages in the secretmem regions are unevictable and unmovable to avoid accidental exposure of the sensitive data via swap or during page migration. Since the secretmem mappings are locked in memory they cannot exceed RLIMIT_MEMLOCK. Since these mappings are already locked independently from mlock(), an attempt to mlock()/munlock() secretmem range would fail and mlockall()/munlockall() will ignore secretmem mappings. However, unlike mlock()ed memory, secretmem currently behaves more like long-term GUP: secretmem mappings are unmovable mappings directly consumed by user space. With default limits, there is no excessive use of secretmem and it poses no real problem in combination with ZONE_MOVABLE/CMA, but in the future this should be addressed to allow balanced use of large amounts of secretmem along with ZONE_MOVABLE/CMA. A page that was a part of the secret memory area is cleared when it is freed to ensure the data is not exposed to the next user of that page. The following example demonstrates creation of a secret mapping (error handling is omitted): fd = memfd_secret(0); ftruncate(fd, MAP_SIZE); ptr = mmap(NULL, MAP_SIZE, PROT_READ | PROT_WRITE, MAP_SHARED, fd, 0); [1] https://lore.kernel.org/linux-mm/213b4567-46ce-f116-9cdf-bbd0c884eb3c@linux.intel.com/ [akpm@linux-foundation.org: suppress Kconfig whine] Link: https://lkml.kernel.org/r/20210518072034.31572-5-rppt@kernel.org Signed-off-by: Mike Rapoport <rppt@linux.ibm.com> Acked-by: Hagen Paul Pfeifer <hagen@jauu.net> Acked-by: James Bottomley <James.Bottomley@HansenPartnership.com> Cc: Alexander Viro <viro@zeniv.linux.org.uk> Cc: Andy Lutomirski <luto@kernel.org> Cc: Arnd Bergmann <arnd@arndb.de> Cc: Borislav Petkov <bp@alien8.de> Cc: Catalin Marinas <catalin.marinas@arm.com> Cc: Christopher Lameter <cl@linux.com> Cc: Dan Williams <dan.j.williams@intel.com> Cc: Dave Hansen <dave.hansen@linux.intel.com> Cc: Elena Reshetova <elena.reshetova@intel.com> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Ingo Molnar <mingo@redhat.com> Cc: James Bottomley <jejb@linux.ibm.com> Cc: "Kirill A. Shutemov" <kirill@shutemov.name> Cc: Matthew Wilcox <willy@infradead.org> Cc: Mark Rutland <mark.rutland@arm.com> Cc: Michael Kerrisk <mtk.manpages@gmail.com> Cc: Palmer Dabbelt <palmer@dabbelt.com> Cc: Palmer Dabbelt <palmerdabbelt@google.com> Cc: Paul Walmsley <paul.walmsley@sifive.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Rick Edgecombe <rick.p.edgecombe@intel.com> Cc: Roman Gushchin <guro@fb.com> Cc: Shakeel Butt <shakeelb@google.com> Cc: Shuah Khan <shuah@kernel.org> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Tycho Andersen <tycho@tycho.ws> Cc: Will Deacon <will@kernel.org> Cc: David Hildenbrand <david@redhat.com> Cc: kernel test robot <lkp@intel.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2021-07-07 18:08:03 -07:00
config SECRETMEM
mm: Kconfig: make config SECRETMEM visible with EXPERT Commit 6a108a14fa35 ("kconfig: rename CONFIG_EMBEDDED to CONFIG_EXPERT") introduces CONFIG_EXPERT to carry the previous intent of CONFIG_EMBEDDED and just gives that intent a much better name. That has been clearly a good and long overdue renaming, and it is clearly an improvement to the kernel build configuration that has shown to help managing the kernel build configuration in the last decade. However, rather than bravely and radically just deleting CONFIG_EMBEDDED, this commit gives CONFIG_EMBEDDED a new intended semantics, but keeps it open for future contributors to implement that intended semantics: A new CONFIG_EMBEDDED option is added that automatically selects CONFIG_EXPERT when enabled and can be used in the future to isolate options that should only be considered for embedded systems (RISC architectures, SLOB, etc). Since then, this CONFIG_EMBEDDED implicitly had two purposes: - It can make even more options visible beyond what CONFIG_EXPERT makes visible. In other words, it may introduce another level of enabling the visibility of configuration options: always visible, visible with CONFIG_EXPERT and visible with CONFIG_EMBEDDED. - Set certain default values of some configurations differently, following the assumption that configuring a kernel build for an embedded system generally starts with a different set of default values compared to kernel builds for all other kind of systems. Considering the second purpose, note that already probably arguing that a kernel build for an embedded system would choose some values differently is already tricky: the set of embedded systems with Linux kernels is already quite diverse. Many embedded system have powerful CPUs and it would not be clear that all embedded systems just optimize towards one specific aspect, e.g., a smaller kernel image size. So, it is unclear if starting with "one set of default configuration" that is induced by CONFIG_EMBEDDED is a good offer for developers configuring their kernels. Also, the differences of needed user-space features in an embedded system compared to a non-embedded system are probably difficult or even impossible to name in some generic way. So it is not surprising that in the last decade hardly anyone has contributed changes to make something default differently in case of CONFIG_EMBEDDED=y. Currently, in v6.0-rc4, SECRETMEM is the only config switched off if CONFIG_EMBEDDED=y. As long as that is actually the only option that currently is selected or deselected, it is better to just make SECRETMEM configurable at build time by experts using menuconfig instead. Make SECRETMEM configurable when EXPERT is set and otherwise default to yes. Further, SECRETMEM needs ARCH_HAS_SET_DIRECT_MAP. This allows us to remove CONFIG_EMBEDDED in the close future. Link: https://lkml.kernel.org/r/20221116131922.25533-1-lukas.bulwahn@gmail.com Signed-off-by: Lukas Bulwahn <lukas.bulwahn@gmail.com> Acked-by: Mike Rapoport <rppt@linux.ibm.com> Acked-by: Arnd Bergmann <arnd@arndb.de> Reviewed-by: Masahiro Yamada <masahiroy@kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2022-11-16 06:19:22 -07:00
default y
bool "Enable memfd_secret() system call" if EXPERT
depends on ARCH_HAS_SET_DIRECT_MAP
help
Enable the memfd_secret() system call with the ability to create
memory areas visible only in the context of the owning process and
not mapped to other processes and other kernel page tables.
mm: introduce memfd_secret system call to create "secret" memory areas Introduce "memfd_secret" system call with the ability to create memory areas visible only in the context of the owning process and not mapped not only to other processes but in the kernel page tables as well. The secretmem feature is off by default and the user must explicitly enable it at the boot time. Once secretmem is enabled, the user will be able to create a file descriptor using the memfd_secret() system call. The memory areas created by mmap() calls from this file descriptor will be unmapped from the kernel direct map and they will be only mapped in the page table of the processes that have access to the file descriptor. Secretmem is designed to provide the following protections: * Enhanced protection (in conjunction with all the other in-kernel attack prevention systems) against ROP attacks. Seceretmem makes "simple" ROP insufficient to perform exfiltration, which increases the required complexity of the attack. Along with other protections like the kernel stack size limit and address space layout randomization which make finding gadgets is really hard, absence of any in-kernel primitive for accessing secret memory means the one gadget ROP attack can't work. Since the only way to access secret memory is to reconstruct the missing mapping entry, the attacker has to recover the physical page and insert a PTE pointing to it in the kernel and then retrieve the contents. That takes at least three gadgets which is a level of difficulty beyond most standard attacks. * Prevent cross-process secret userspace memory exposures. Once the secret memory is allocated, the user can't accidentally pass it into the kernel to be transmitted somewhere. The secreremem pages cannot be accessed via the direct map and they are disallowed in GUP. * Harden against exploited kernel flaws. In order to access secretmem, a kernel-side attack would need to either walk the page tables and create new ones, or spawn a new privileged uiserspace process to perform secrets exfiltration using ptrace. The file descriptor based memory has several advantages over the "traditional" mm interfaces, such as mlock(), mprotect(), madvise(). File descriptor approach allows explicit and controlled sharing of the memory areas, it allows to seal the operations. Besides, file descriptor based memory paves the way for VMMs to remove the secret memory range from the userspace hipervisor process, for instance QEMU. Andy Lutomirski says: "Getting fd-backed memory into a guest will take some possibly major work in the kernel, but getting vma-backed memory into a guest without mapping it in the host user address space seems much, much worse." memfd_secret() is made a dedicated system call rather than an extension to memfd_create() because it's purpose is to allow the user to create more secure memory mappings rather than to simply allow file based access to the memory. Nowadays a new system call cost is negligible while it is way simpler for userspace to deal with a clear-cut system calls than with a multiplexer or an overloaded syscall. Moreover, the initial implementation of memfd_secret() is completely distinct from memfd_create() so there is no much sense in overloading memfd_create() to begin with. If there will be a need for code sharing between these implementation it can be easily achieved without a need to adjust user visible APIs. The secret memory remains accessible in the process context using uaccess primitives, but it is not exposed to the kernel otherwise; secret memory areas are removed from the direct map and functions in the follow_page()/get_user_page() family will refuse to return a page that belongs to the secret memory area. Once there will be a use case that will require exposing secretmem to the kernel it will be an opt-in request in the system call flags so that user would have to decide what data can be exposed to the kernel. Removing of the pages from the direct map may cause its fragmentation on architectures that use large pages to map the physical memory which affects the system performance. However, the original Kconfig text for CONFIG_DIRECT_GBPAGES said that gigabyte pages in the direct map "... can improve the kernel's performance a tiny bit ..." (commit 00d1c5e05736 ("x86: add gbpages switches")) and the recent report [1] showed that "... although 1G mappings are a good default choice, there is no compelling evidence that it must be the only choice". Hence, it is sufficient to have secretmem disabled by default with the ability of a system administrator to enable it at boot time. Pages in the secretmem regions are unevictable and unmovable to avoid accidental exposure of the sensitive data via swap or during page migration. Since the secretmem mappings are locked in memory they cannot exceed RLIMIT_MEMLOCK. Since these mappings are already locked independently from mlock(), an attempt to mlock()/munlock() secretmem range would fail and mlockall()/munlockall() will ignore secretmem mappings. However, unlike mlock()ed memory, secretmem currently behaves more like long-term GUP: secretmem mappings are unmovable mappings directly consumed by user space. With default limits, there is no excessive use of secretmem and it poses no real problem in combination with ZONE_MOVABLE/CMA, but in the future this should be addressed to allow balanced use of large amounts of secretmem along with ZONE_MOVABLE/CMA. A page that was a part of the secret memory area is cleared when it is freed to ensure the data is not exposed to the next user of that page. The following example demonstrates creation of a secret mapping (error handling is omitted): fd = memfd_secret(0); ftruncate(fd, MAP_SIZE); ptr = mmap(NULL, MAP_SIZE, PROT_READ | PROT_WRITE, MAP_SHARED, fd, 0); [1] https://lore.kernel.org/linux-mm/213b4567-46ce-f116-9cdf-bbd0c884eb3c@linux.intel.com/ [akpm@linux-foundation.org: suppress Kconfig whine] Link: https://lkml.kernel.org/r/20210518072034.31572-5-rppt@kernel.org Signed-off-by: Mike Rapoport <rppt@linux.ibm.com> Acked-by: Hagen Paul Pfeifer <hagen@jauu.net> Acked-by: James Bottomley <James.Bottomley@HansenPartnership.com> Cc: Alexander Viro <viro@zeniv.linux.org.uk> Cc: Andy Lutomirski <luto@kernel.org> Cc: Arnd Bergmann <arnd@arndb.de> Cc: Borislav Petkov <bp@alien8.de> Cc: Catalin Marinas <catalin.marinas@arm.com> Cc: Christopher Lameter <cl@linux.com> Cc: Dan Williams <dan.j.williams@intel.com> Cc: Dave Hansen <dave.hansen@linux.intel.com> Cc: Elena Reshetova <elena.reshetova@intel.com> Cc: "H. Peter Anvin" <hpa@zytor.com> Cc: Ingo Molnar <mingo@redhat.com> Cc: James Bottomley <jejb@linux.ibm.com> Cc: "Kirill A. Shutemov" <kirill@shutemov.name> Cc: Matthew Wilcox <willy@infradead.org> Cc: Mark Rutland <mark.rutland@arm.com> Cc: Michael Kerrisk <mtk.manpages@gmail.com> Cc: Palmer Dabbelt <palmer@dabbelt.com> Cc: Palmer Dabbelt <palmerdabbelt@google.com> Cc: Paul Walmsley <paul.walmsley@sifive.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Rick Edgecombe <rick.p.edgecombe@intel.com> Cc: Roman Gushchin <guro@fb.com> Cc: Shakeel Butt <shakeelb@google.com> Cc: Shuah Khan <shuah@kernel.org> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Tycho Andersen <tycho@tycho.ws> Cc: Will Deacon <will@kernel.org> Cc: David Hildenbrand <david@redhat.com> Cc: kernel test robot <lkp@intel.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2021-07-07 18:08:03 -07:00
mm: add a field to store names for private anonymous memory In many userspace applications, and especially in VM based applications like Android uses heavily, there are multiple different allocators in use. At a minimum there is libc malloc and the stack, and in many cases there are libc malloc, the stack, direct syscalls to mmap anonymous memory, and multiple VM heaps (one for small objects, one for big objects, etc.). Each of these layers usually has its own tools to inspect its usage; malloc by compiling a debug version, the VM through heap inspection tools, and for direct syscalls there is usually no way to track them. On Android we heavily use a set of tools that use an extended version of the logic covered in Documentation/vm/pagemap.txt to walk all pages mapped in userspace and slice their usage by process, shared (COW) vs. unique mappings, backing, etc. This can account for real physical memory usage even in cases like fork without exec (which Android uses heavily to share as many private COW pages as possible between processes), Kernel SamePage Merging, and clean zero pages. It produces a measurement of the pages that only exist in that process (USS, for unique), and a measurement of the physical memory usage of that process with the cost of shared pages being evenly split between processes that share them (PSS). If all anonymous memory is indistinguishable then figuring out the real physical memory usage (PSS) of each heap requires either a pagemap walking tool that can understand the heap debugging of every layer, or for every layer's heap debugging tools to implement the pagemap walking logic, in which case it is hard to get a consistent view of memory across the whole system. Tracking the information in userspace leads to all sorts of problems. It either needs to be stored inside the process, which means every process has to have an API to export its current heap information upon request, or it has to be stored externally in a filesystem that somebody needs to clean up on crashes. It needs to be readable while the process is still running, so it has to have some sort of synchronization with every layer of userspace. Efficiently tracking the ranges requires reimplementing something like the kernel vma trees, and linking to it from every layer of userspace. It requires more memory, more syscalls, more runtime cost, and more complexity to separately track regions that the kernel is already tracking. This patch adds a field to /proc/pid/maps and /proc/pid/smaps to show a userspace-provided name for anonymous vmas. The names of named anonymous vmas are shown in /proc/pid/maps and /proc/pid/smaps as [anon:<name>]. Userspace can set the name for a region of memory by calling prctl(PR_SET_VMA, PR_SET_VMA_ANON_NAME, start, len, (unsigned long)name) Setting the name to NULL clears it. The name length limit is 80 bytes including NUL-terminator and is checked to contain only printable ascii characters (including space), except '[',']','\','$' and '`'. Ascii strings are being used to have a descriptive identifiers for vmas, which can be understood by the users reading /proc/pid/maps or /proc/pid/smaps. Names can be standardized for a given system and they can include some variable parts such as the name of the allocator or a library, tid of the thread using it, etc. The name is stored in a pointer in the shared union in vm_area_struct that points to a null terminated string. Anonymous vmas with the same name (equivalent strings) and are otherwise mergeable will be merged. The name pointers are not shared between vmas even if they contain the same name. The name pointer is stored in a union with fields that are only used on file-backed mappings, so it does not increase memory usage. CONFIG_ANON_VMA_NAME kernel configuration is introduced to enable this feature. It keeps the feature disabled by default to prevent any additional memory overhead and to avoid confusing procfs parsers on systems which are not ready to support named anonymous vmas. The patch is based on the original patch developed by Colin Cross, more specifically on its latest version [1] posted upstream by Sumit Semwal. It used a userspace pointer to store vma names. In that design, name pointers could be shared between vmas. However during the last upstreaming attempt, Kees Cook raised concerns [2] about this approach and suggested to copy the name into kernel memory space, perform validity checks [3] and store as a string referenced from vm_area_struct. One big concern is about fork() performance which would need to strdup anonymous vma names. Dave Hansen suggested experimenting with worst-case scenario of forking a process with 64k vmas having longest possible names [4]. I ran this experiment on an ARM64 Android device and recorded a worst-case regression of almost 40% when forking such a process. This regression is addressed in the followup patch which replaces the pointer to a name with a refcounted structure that allows sharing the name pointer between vmas of the same name. Instead of duplicating the string during fork() or when splitting a vma it increments the refcount. [1] https://lore.kernel.org/linux-mm/20200901161459.11772-4-sumit.semwal@linaro.org/ [2] https://lore.kernel.org/linux-mm/202009031031.D32EF57ED@keescook/ [3] https://lore.kernel.org/linux-mm/202009031022.3834F692@keescook/ [4] https://lore.kernel.org/linux-mm/5d0358ab-8c47-2f5f-8e43-23b89d6a8e95@intel.com/ Changes for prctl(2) manual page (in the options section): PR_SET_VMA Sets an attribute specified in arg2 for virtual memory areas starting from the address specified in arg3 and spanning the size specified in arg4. arg5 specifies the value of the attribute to be set. Note that assigning an attribute to a virtual memory area might prevent it from being merged with adjacent virtual memory areas due to the difference in that attribute's value. Currently, arg2 must be one of: PR_SET_VMA_ANON_NAME Set a name for anonymous virtual memory areas. arg5 should be a pointer to a null-terminated string containing the name. The name length including null byte cannot exceed 80 bytes. If arg5 is NULL, the name of the appropriate anonymous virtual memory areas will be reset. The name can contain only printable ascii characters (including space), except '[',']','\','$' and '`'. This feature is available only if the kernel is built with the CONFIG_ANON_VMA_NAME option enabled. [surenb@google.com: docs: proc.rst: /proc/PID/maps: fix malformed table] Link: https://lkml.kernel.org/r/20211123185928.2513763-1-surenb@google.com [surenb: rebased over v5.15-rc6, replaced userpointer with a kernel copy, added input sanitization and CONFIG_ANON_VMA_NAME config. The bulk of the work here was done by Colin Cross, therefore, with his permission, keeping him as the author] Link: https://lkml.kernel.org/r/20211019215511.3771969-2-surenb@google.com Signed-off-by: Colin Cross <ccross@google.com> Signed-off-by: Suren Baghdasaryan <surenb@google.com> Reviewed-by: Kees Cook <keescook@chromium.org> Cc: Stephen Rothwell <sfr@canb.auug.org.au> Cc: Al Viro <viro@zeniv.linux.org.uk> Cc: Cyrill Gorcunov <gorcunov@openvz.org> Cc: Dave Hansen <dave.hansen@intel.com> Cc: David Rientjes <rientjes@google.com> Cc: "Eric W. Biederman" <ebiederm@xmission.com> Cc: Hugh Dickins <hughd@google.com> Cc: Ingo Molnar <mingo@kernel.org> Cc: Jan Glauber <jan.glauber@gmail.com> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: John Stultz <john.stultz@linaro.org> Cc: Mel Gorman <mgorman@suse.de> Cc: Minchan Kim <minchan@kernel.org> Cc: Oleg Nesterov <oleg@redhat.com> Cc: Pekka Enberg <penberg@kernel.org> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Rob Landley <rob@landley.net> Cc: "Serge E. Hallyn" <serge.hallyn@ubuntu.com> Cc: Shaohua Li <shli@fusionio.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2022-01-14 15:05:59 -07:00
config ANON_VMA_NAME
bool "Anonymous VMA name support"
depends on PROC_FS && ADVISE_SYSCALLS && MMU
help
Allow naming anonymous virtual memory areas.
This feature allows assigning names to virtual memory areas. Assigned
names can be later retrieved from /proc/pid/maps and /proc/pid/smaps
and help identifying individual anonymous memory areas.
Assigning a name to anonymous virtual memory area might prevent that
area from being merged with adjacent virtual memory areas due to the
difference in their name.
config HAVE_ARCH_USERFAULTFD_WP
bool
help
Arch has userfaultfd write protection support
config HAVE_ARCH_USERFAULTFD_MINOR
bool
help
Arch has userfaultfd minor fault support
menuconfig USERFAULTFD
bool "Enable userfaultfd() system call"
depends on MMU
help
Enable the userfaultfd() system call that allows to intercept and
handle page faults in userland.
if USERFAULTFD
mm/uffd: PTE_MARKER_UFFD_WP This patch introduces the 1st user of pte marker: the uffd-wp marker. When the pte marker is installed with the uffd-wp bit set, it means this pte was wr-protected by uffd. We will use this special pte to arm the ptes that got either unmapped or swapped out for a file-backed region that was previously wr-protected. This special pte could trigger a page fault just like swap entries. This idea is greatly inspired by Hugh and Andrea in the discussion, which is referenced in the links below. Some helpers are introduced to detect whether a swap pte is uffd wr-protected. After the pte marker introduced, one swap pte can be wr-protected in two forms: either it is a normal swap pte and it has _PAGE_SWP_UFFD_WP set, or it's a pte marker that has PTE_MARKER_UFFD_WP set. [peterx@redhat.com: fixup] Link: https://lkml.kernel.org/r/YkzKiM8tI4+qOfXF@xz-m1.local Link: https://lore.kernel.org/lkml/20201126222359.8120-1-peterx@redhat.com/ Link: https://lore.kernel.org/lkml/20201130230603.46187-1-peterx@redhat.com/ Link: https://lkml.kernel.org/r/20220405014838.14131-1-peterx@redhat.com Signed-off-by: Peter Xu <peterx@redhat.com> Suggested-by: Andrea Arcangeli <aarcange@redhat.com> Suggested-by: Hugh Dickins <hughd@google.com> Cc: Alistair Popple <apopple@nvidia.com> Cc: Axel Rasmussen <axelrasmussen@google.com> Cc: David Hildenbrand <david@redhat.com> Cc: Jerome Glisse <jglisse@redhat.com> Cc: "Kirill A . Shutemov" <kirill@shutemov.name> Cc: Matthew Wilcox <willy@infradead.org> Cc: Mike Kravetz <mike.kravetz@oracle.com> Cc: Mike Rapoport <rppt@linux.vnet.ibm.com> Cc: Nadav Amit <nadav.amit@gmail.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2022-05-12 20:22:52 -07:00
config PTE_MARKER_UFFD_WP
bool "Userfaultfd write protection support for shmem/hugetlbfs"
default y
depends on HAVE_ARCH_USERFAULTFD_WP
mm/uffd: PTE_MARKER_UFFD_WP This patch introduces the 1st user of pte marker: the uffd-wp marker. When the pte marker is installed with the uffd-wp bit set, it means this pte was wr-protected by uffd. We will use this special pte to arm the ptes that got either unmapped or swapped out for a file-backed region that was previously wr-protected. This special pte could trigger a page fault just like swap entries. This idea is greatly inspired by Hugh and Andrea in the discussion, which is referenced in the links below. Some helpers are introduced to detect whether a swap pte is uffd wr-protected. After the pte marker introduced, one swap pte can be wr-protected in two forms: either it is a normal swap pte and it has _PAGE_SWP_UFFD_WP set, or it's a pte marker that has PTE_MARKER_UFFD_WP set. [peterx@redhat.com: fixup] Link: https://lkml.kernel.org/r/YkzKiM8tI4+qOfXF@xz-m1.local Link: https://lore.kernel.org/lkml/20201126222359.8120-1-peterx@redhat.com/ Link: https://lore.kernel.org/lkml/20201130230603.46187-1-peterx@redhat.com/ Link: https://lkml.kernel.org/r/20220405014838.14131-1-peterx@redhat.com Signed-off-by: Peter Xu <peterx@redhat.com> Suggested-by: Andrea Arcangeli <aarcange@redhat.com> Suggested-by: Hugh Dickins <hughd@google.com> Cc: Alistair Popple <apopple@nvidia.com> Cc: Axel Rasmussen <axelrasmussen@google.com> Cc: David Hildenbrand <david@redhat.com> Cc: Jerome Glisse <jglisse@redhat.com> Cc: "Kirill A . Shutemov" <kirill@shutemov.name> Cc: Matthew Wilcox <willy@infradead.org> Cc: Mike Kravetz <mike.kravetz@oracle.com> Cc: Mike Rapoport <rppt@linux.vnet.ibm.com> Cc: Nadav Amit <nadav.amit@gmail.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2022-05-12 20:22:52 -07:00
help
Allows to create marker PTEs for userfaultfd write protection
purposes. It is required to enable userfaultfd write protection on
file-backed memory types like shmem and hugetlbfs.
endif # USERFAULTFD
mm/uffd: PTE_MARKER_UFFD_WP This patch introduces the 1st user of pte marker: the uffd-wp marker. When the pte marker is installed with the uffd-wp bit set, it means this pte was wr-protected by uffd. We will use this special pte to arm the ptes that got either unmapped or swapped out for a file-backed region that was previously wr-protected. This special pte could trigger a page fault just like swap entries. This idea is greatly inspired by Hugh and Andrea in the discussion, which is referenced in the links below. Some helpers are introduced to detect whether a swap pte is uffd wr-protected. After the pte marker introduced, one swap pte can be wr-protected in two forms: either it is a normal swap pte and it has _PAGE_SWP_UFFD_WP set, or it's a pte marker that has PTE_MARKER_UFFD_WP set. [peterx@redhat.com: fixup] Link: https://lkml.kernel.org/r/YkzKiM8tI4+qOfXF@xz-m1.local Link: https://lore.kernel.org/lkml/20201126222359.8120-1-peterx@redhat.com/ Link: https://lore.kernel.org/lkml/20201130230603.46187-1-peterx@redhat.com/ Link: https://lkml.kernel.org/r/20220405014838.14131-1-peterx@redhat.com Signed-off-by: Peter Xu <peterx@redhat.com> Suggested-by: Andrea Arcangeli <aarcange@redhat.com> Suggested-by: Hugh Dickins <hughd@google.com> Cc: Alistair Popple <apopple@nvidia.com> Cc: Axel Rasmussen <axelrasmussen@google.com> Cc: David Hildenbrand <david@redhat.com> Cc: Jerome Glisse <jglisse@redhat.com> Cc: "Kirill A . Shutemov" <kirill@shutemov.name> Cc: Matthew Wilcox <willy@infradead.org> Cc: Mike Kravetz <mike.kravetz@oracle.com> Cc: Mike Rapoport <rppt@linux.vnet.ibm.com> Cc: Nadav Amit <nadav.amit@gmail.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2022-05-12 20:22:52 -07:00
mm: multi-gen LRU: minimal implementation To avoid confusion, the terms "promotion" and "demotion" will be applied to the multi-gen LRU, as a new convention; the terms "activation" and "deactivation" will be applied to the active/inactive LRU, as usual. The aging produces young generations. Given an lruvec, it increments max_seq when max_seq-min_seq+1 approaches MIN_NR_GENS. The aging promotes hot pages to the youngest generation when it finds them accessed through page tables; the demotion of cold pages happens consequently when it increments max_seq. Promotion in the aging path does not involve any LRU list operations, only the updates of the gen counter and lrugen->nr_pages[]; demotion, unless as the result of the increment of max_seq, requires LRU list operations, e.g., lru_deactivate_fn(). The aging has the complexity O(nr_hot_pages), since it is only interested in hot pages. The eviction consumes old generations. Given an lruvec, it increments min_seq when lrugen->lists[] indexed by min_seq%MAX_NR_GENS becomes empty. A feedback loop modeled after the PID controller monitors refaults over anon and file types and decides which type to evict when both types are available from the same generation. The protection of pages accessed multiple times through file descriptors takes place in the eviction path. Each generation is divided into multiple tiers. A page accessed N times through file descriptors is in tier order_base_2(N). Tiers do not have dedicated lrugen->lists[], only bits in folio->flags. The aforementioned feedback loop also monitors refaults over all tiers and decides when to protect pages in which tiers (N>1), using the first tier (N=0,1) as a baseline. The first tier contains single-use unmapped clean pages, which are most likely the best choices. In contrast to promotion in the aging path, the protection of a page in the eviction path is achieved by moving this page to the next generation, i.e., min_seq+1, if the feedback loop decides so. This approach has the following advantages: 1. It removes the cost of activation in the buffered access path by inferring whether pages accessed multiple times through file descriptors are statistically hot and thus worth protecting in the eviction path. 2. It takes pages accessed through page tables into account and avoids overprotecting pages accessed multiple times through file descriptors. (Pages accessed through page tables are in the first tier, since N=0.) 3. More tiers provide better protection for pages accessed more than twice through file descriptors, when under heavy buffered I/O workloads. Server benchmark results: Single workload: fio (buffered I/O): +[30, 32]% IOPS BW 5.19-rc1: 2673k 10.2GiB/s patch1-6: 3491k 13.3GiB/s Single workload: memcached (anon): -[4, 6]% Ops/sec KB/sec 5.19-rc1: 1161501.04 45177.25 patch1-6: 1106168.46 43025.04 Configurations: CPU: two Xeon 6154 Mem: total 256G Node 1 was only used as a ram disk to reduce the variance in the results. patch drivers/block/brd.c <<EOF 99,100c99,100 < gfp_flags = GFP_NOIO | __GFP_ZERO | __GFP_HIGHMEM; < page = alloc_page(gfp_flags); --- > gfp_flags = GFP_NOIO | __GFP_ZERO | __GFP_HIGHMEM | __GFP_THISNODE; > page = alloc_pages_node(1, gfp_flags, 0); EOF cat >>/etc/systemd/system.conf <<EOF CPUAffinity=numa NUMAPolicy=bind NUMAMask=0 EOF cat >>/etc/memcached.conf <<EOF -m 184320 -s /var/run/memcached/memcached.sock -a 0766 -t 36 -B binary EOF cat fio.sh modprobe brd rd_nr=1 rd_size=113246208 swapoff -a mkfs.ext4 /dev/ram0 mount -t ext4 /dev/ram0 /mnt mkdir /sys/fs/cgroup/user.slice/test echo 38654705664 >/sys/fs/cgroup/user.slice/test/memory.max echo $$ >/sys/fs/cgroup/user.slice/test/cgroup.procs fio -name=mglru --numjobs=72 --directory=/mnt --size=1408m \ --buffered=1 --ioengine=io_uring --iodepth=128 \ --iodepth_batch_submit=32 --iodepth_batch_complete=32 \ --rw=randread --random_distribution=random --norandommap \ --time_based --ramp_time=10m --runtime=5m --group_reporting cat memcached.sh modprobe brd rd_nr=1 rd_size=113246208 swapoff -a mkswap /dev/ram0 swapon /dev/ram0 memtier_benchmark -S /var/run/memcached/memcached.sock \ -P memcache_binary -n allkeys --key-minimum=1 \ --key-maximum=65000000 --key-pattern=P:P -c 1 -t 36 \ --ratio 1:0 --pipeline 8 -d 2000 memtier_benchmark -S /var/run/memcached/memcached.sock \ -P memcache_binary -n allkeys --key-minimum=1 \ --key-maximum=65000000 --key-pattern=R:R -c 1 -t 36 \ --ratio 0:1 --pipeline 8 --randomize --distinct-client-seed Client benchmark results: kswapd profiles: 5.19-rc1 40.33% page_vma_mapped_walk (overhead) 21.80% lzo1x_1_do_compress (real work) 7.53% do_raw_spin_lock 3.95% _raw_spin_unlock_irq 2.52% vma_interval_tree_iter_next 2.37% folio_referenced_one 2.28% vma_interval_tree_subtree_search 1.97% anon_vma_interval_tree_iter_first 1.60% ptep_clear_flush 1.06% __zram_bvec_write patch1-6 39.03% lzo1x_1_do_compress (real work) 18.47% page_vma_mapped_walk (overhead) 6.74% _raw_spin_unlock_irq 3.97% do_raw_spin_lock 2.49% ptep_clear_flush 2.48% anon_vma_interval_tree_iter_first 1.92% folio_referenced_one 1.88% __zram_bvec_write 1.48% memmove 1.31% vma_interval_tree_iter_next Configurations: CPU: single Snapdragon 7c Mem: total 4G ChromeOS MemoryPressure [1] [1] https://chromium.googlesource.com/chromiumos/platform/tast-tests/ Link: https://lkml.kernel.org/r/20220918080010.2920238-7-yuzhao@google.com Signed-off-by: Yu Zhao <yuzhao@google.com> Acked-by: Brian Geffon <bgeffon@google.com> Acked-by: Jan Alexander Steffens (heftig) <heftig@archlinux.org> Acked-by: Oleksandr Natalenko <oleksandr@natalenko.name> Acked-by: Steven Barrett <steven@liquorix.net> Acked-by: Suleiman Souhlal <suleiman@google.com> Tested-by: Daniel Byrne <djbyrne@mtu.edu> Tested-by: Donald Carr <d@chaos-reins.com> Tested-by: Holger Hoffstätte <holger@applied-asynchrony.com> Tested-by: Konstantin Kharlamov <Hi-Angel@yandex.ru> Tested-by: Shuang Zhai <szhai2@cs.rochester.edu> Tested-by: Sofia Trinh <sofia.trinh@edi.works> Tested-by: Vaibhav Jain <vaibhav@linux.ibm.com> Cc: Andi Kleen <ak@linux.intel.com> Cc: Aneesh Kumar K.V <aneesh.kumar@linux.ibm.com> Cc: Barry Song <baohua@kernel.org> Cc: Catalin Marinas <catalin.marinas@arm.com> Cc: Dave Hansen <dave.hansen@linux.intel.com> Cc: Hillf Danton <hdanton@sina.com> Cc: Jens Axboe <axboe@kernel.dk> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Jonathan Corbet <corbet@lwn.net> Cc: Linus Torvalds <torvalds@linux-foundation.org> Cc: Matthew Wilcox <willy@infradead.org> Cc: Mel Gorman <mgorman@suse.de> Cc: Miaohe Lin <linmiaohe@huawei.com> Cc: Michael Larabel <Michael@MichaelLarabel.com> Cc: Michal Hocko <mhocko@kernel.org> Cc: Mike Rapoport <rppt@kernel.org> Cc: Mike Rapoport <rppt@linux.ibm.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Qi Zheng <zhengqi.arch@bytedance.com> Cc: Tejun Heo <tj@kernel.org> Cc: Vlastimil Babka <vbabka@suse.cz> Cc: Will Deacon <will@kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2022-09-18 01:00:03 -07:00
# multi-gen LRU {
mm: multi-gen LRU: groundwork Evictable pages are divided into multiple generations for each lruvec. The youngest generation number is stored in lrugen->max_seq for both anon and file types as they are aged on an equal footing. The oldest generation numbers are stored in lrugen->min_seq[] separately for anon and file types as clean file pages can be evicted regardless of swap constraints. These three variables are monotonically increasing. Generation numbers are truncated into order_base_2(MAX_NR_GENS+1) bits in order to fit into the gen counter in folio->flags. Each truncated generation number is an index to lrugen->lists[]. The sliding window technique is used to track at least MIN_NR_GENS and at most MAX_NR_GENS generations. The gen counter stores a value within [1, MAX_NR_GENS] while a page is on one of lrugen->lists[]. Otherwise it stores 0. There are two conceptually independent procedures: "the aging", which produces young generations, and "the eviction", which consumes old generations. They form a closed-loop system, i.e., "the page reclaim". Both procedures can be invoked from userspace for the purposes of working set estimation and proactive reclaim. These techniques are commonly used to optimize job scheduling (bin packing) in data centers [1][2]. To avoid confusion, the terms "hot" and "cold" will be applied to the multi-gen LRU, as a new convention; the terms "active" and "inactive" will be applied to the active/inactive LRU, as usual. The protection of hot pages and the selection of cold pages are based on page access channels and patterns. There are two access channels: one through page tables and the other through file descriptors. The protection of the former channel is by design stronger because: 1. The uncertainty in determining the access patterns of the former channel is higher due to the approximation of the accessed bit. 2. The cost of evicting the former channel is higher due to the TLB flushes required and the likelihood of encountering the dirty bit. 3. The penalty of underprotecting the former channel is higher because applications usually do not prepare themselves for major page faults like they do for blocked I/O. E.g., GUI applications commonly use dedicated I/O threads to avoid blocking rendering threads. There are also two access patterns: one with temporal locality and the other without. For the reasons listed above, the former channel is assumed to follow the former pattern unless VM_SEQ_READ or VM_RAND_READ is present; the latter channel is assumed to follow the latter pattern unless outlying refaults have been observed [3][4]. The next patch will address the "outlying refaults". Three macros, i.e., LRU_REFS_WIDTH, LRU_REFS_PGOFF and LRU_REFS_MASK, used later are added in this patch to make the entire patchset less diffy. A page is added to the youngest generation on faulting. The aging needs to check the accessed bit at least twice before handing this page over to the eviction. The first check takes care of the accessed bit set on the initial fault; the second check makes sure this page has not been used since then. This protocol, AKA second chance, requires a minimum of two generations, hence MIN_NR_GENS. [1] https://dl.acm.org/doi/10.1145/3297858.3304053 [2] https://dl.acm.org/doi/10.1145/3503222.3507731 [3] https://lwn.net/Articles/495543/ [4] https://lwn.net/Articles/815342/ Link: https://lkml.kernel.org/r/20220918080010.2920238-6-yuzhao@google.com Signed-off-by: Yu Zhao <yuzhao@google.com> Acked-by: Brian Geffon <bgeffon@google.com> Acked-by: Jan Alexander Steffens (heftig) <heftig@archlinux.org> Acked-by: Oleksandr Natalenko <oleksandr@natalenko.name> Acked-by: Steven Barrett <steven@liquorix.net> Acked-by: Suleiman Souhlal <suleiman@google.com> Tested-by: Daniel Byrne <djbyrne@mtu.edu> Tested-by: Donald Carr <d@chaos-reins.com> Tested-by: Holger Hoffstätte <holger@applied-asynchrony.com> Tested-by: Konstantin Kharlamov <Hi-Angel@yandex.ru> Tested-by: Shuang Zhai <szhai2@cs.rochester.edu> Tested-by: Sofia Trinh <sofia.trinh@edi.works> Tested-by: Vaibhav Jain <vaibhav@linux.ibm.com> Cc: Andi Kleen <ak@linux.intel.com> Cc: Aneesh Kumar K.V <aneesh.kumar@linux.ibm.com> Cc: Barry Song <baohua@kernel.org> Cc: Catalin Marinas <catalin.marinas@arm.com> Cc: Dave Hansen <dave.hansen@linux.intel.com> Cc: Hillf Danton <hdanton@sina.com> Cc: Jens Axboe <axboe@kernel.dk> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Jonathan Corbet <corbet@lwn.net> Cc: Linus Torvalds <torvalds@linux-foundation.org> Cc: Matthew Wilcox <willy@infradead.org> Cc: Mel Gorman <mgorman@suse.de> Cc: Miaohe Lin <linmiaohe@huawei.com> Cc: Michael Larabel <Michael@MichaelLarabel.com> Cc: Michal Hocko <mhocko@kernel.org> Cc: Mike Rapoport <rppt@kernel.org> Cc: Mike Rapoport <rppt@linux.ibm.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Qi Zheng <zhengqi.arch@bytedance.com> Cc: Tejun Heo <tj@kernel.org> Cc: Vlastimil Babka <vbabka@suse.cz> Cc: Will Deacon <will@kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2022-09-18 01:00:02 -07:00
config LRU_GEN
bool "Multi-Gen LRU"
depends on MMU
# make sure folio->flags has enough spare bits
depends on 64BIT || !SPARSEMEM || SPARSEMEM_VMEMMAP
help
mm: multi-gen LRU: admin guide Add an admin guide. Link: https://lkml.kernel.org/r/20220918080010.2920238-14-yuzhao@google.com Signed-off-by: Yu Zhao <yuzhao@google.com> Acked-by: Brian Geffon <bgeffon@google.com> Acked-by: Jan Alexander Steffens (heftig) <heftig@archlinux.org> Acked-by: Oleksandr Natalenko <oleksandr@natalenko.name> Acked-by: Steven Barrett <steven@liquorix.net> Acked-by: Suleiman Souhlal <suleiman@google.com> Acked-by: Mike Rapoport <rppt@linux.ibm.com> Tested-by: Daniel Byrne <djbyrne@mtu.edu> Tested-by: Donald Carr <d@chaos-reins.com> Tested-by: Holger Hoffstätte <holger@applied-asynchrony.com> Tested-by: Konstantin Kharlamov <Hi-Angel@yandex.ru> Tested-by: Shuang Zhai <szhai2@cs.rochester.edu> Tested-by: Sofia Trinh <sofia.trinh@edi.works> Tested-by: Vaibhav Jain <vaibhav@linux.ibm.com> Cc: Andi Kleen <ak@linux.intel.com> Cc: Aneesh Kumar K.V <aneesh.kumar@linux.ibm.com> Cc: Barry Song <baohua@kernel.org> Cc: Catalin Marinas <catalin.marinas@arm.com> Cc: Dave Hansen <dave.hansen@linux.intel.com> Cc: Hillf Danton <hdanton@sina.com> Cc: Jens Axboe <axboe@kernel.dk> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Jonathan Corbet <corbet@lwn.net> Cc: Linus Torvalds <torvalds@linux-foundation.org> Cc: Matthew Wilcox <willy@infradead.org> Cc: Mel Gorman <mgorman@suse.de> Cc: Miaohe Lin <linmiaohe@huawei.com> Cc: Michael Larabel <Michael@MichaelLarabel.com> Cc: Michal Hocko <mhocko@kernel.org> Cc: Mike Rapoport <rppt@kernel.org> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Qi Zheng <zhengqi.arch@bytedance.com> Cc: Tejun Heo <tj@kernel.org> Cc: Vlastimil Babka <vbabka@suse.cz> Cc: Will Deacon <will@kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2022-09-18 01:00:10 -07:00
A high performance LRU implementation to overcommit memory. See
Documentation/admin-guide/mm/multigen_lru.rst for details.
mm: multi-gen LRU: groundwork Evictable pages are divided into multiple generations for each lruvec. The youngest generation number is stored in lrugen->max_seq for both anon and file types as they are aged on an equal footing. The oldest generation numbers are stored in lrugen->min_seq[] separately for anon and file types as clean file pages can be evicted regardless of swap constraints. These three variables are monotonically increasing. Generation numbers are truncated into order_base_2(MAX_NR_GENS+1) bits in order to fit into the gen counter in folio->flags. Each truncated generation number is an index to lrugen->lists[]. The sliding window technique is used to track at least MIN_NR_GENS and at most MAX_NR_GENS generations. The gen counter stores a value within [1, MAX_NR_GENS] while a page is on one of lrugen->lists[]. Otherwise it stores 0. There are two conceptually independent procedures: "the aging", which produces young generations, and "the eviction", which consumes old generations. They form a closed-loop system, i.e., "the page reclaim". Both procedures can be invoked from userspace for the purposes of working set estimation and proactive reclaim. These techniques are commonly used to optimize job scheduling (bin packing) in data centers [1][2]. To avoid confusion, the terms "hot" and "cold" will be applied to the multi-gen LRU, as a new convention; the terms "active" and "inactive" will be applied to the active/inactive LRU, as usual. The protection of hot pages and the selection of cold pages are based on page access channels and patterns. There are two access channels: one through page tables and the other through file descriptors. The protection of the former channel is by design stronger because: 1. The uncertainty in determining the access patterns of the former channel is higher due to the approximation of the accessed bit. 2. The cost of evicting the former channel is higher due to the TLB flushes required and the likelihood of encountering the dirty bit. 3. The penalty of underprotecting the former channel is higher because applications usually do not prepare themselves for major page faults like they do for blocked I/O. E.g., GUI applications commonly use dedicated I/O threads to avoid blocking rendering threads. There are also two access patterns: one with temporal locality and the other without. For the reasons listed above, the former channel is assumed to follow the former pattern unless VM_SEQ_READ or VM_RAND_READ is present; the latter channel is assumed to follow the latter pattern unless outlying refaults have been observed [3][4]. The next patch will address the "outlying refaults". Three macros, i.e., LRU_REFS_WIDTH, LRU_REFS_PGOFF and LRU_REFS_MASK, used later are added in this patch to make the entire patchset less diffy. A page is added to the youngest generation on faulting. The aging needs to check the accessed bit at least twice before handing this page over to the eviction. The first check takes care of the accessed bit set on the initial fault; the second check makes sure this page has not been used since then. This protocol, AKA second chance, requires a minimum of two generations, hence MIN_NR_GENS. [1] https://dl.acm.org/doi/10.1145/3297858.3304053 [2] https://dl.acm.org/doi/10.1145/3503222.3507731 [3] https://lwn.net/Articles/495543/ [4] https://lwn.net/Articles/815342/ Link: https://lkml.kernel.org/r/20220918080010.2920238-6-yuzhao@google.com Signed-off-by: Yu Zhao <yuzhao@google.com> Acked-by: Brian Geffon <bgeffon@google.com> Acked-by: Jan Alexander Steffens (heftig) <heftig@archlinux.org> Acked-by: Oleksandr Natalenko <oleksandr@natalenko.name> Acked-by: Steven Barrett <steven@liquorix.net> Acked-by: Suleiman Souhlal <suleiman@google.com> Tested-by: Daniel Byrne <djbyrne@mtu.edu> Tested-by: Donald Carr <d@chaos-reins.com> Tested-by: Holger Hoffstätte <holger@applied-asynchrony.com> Tested-by: Konstantin Kharlamov <Hi-Angel@yandex.ru> Tested-by: Shuang Zhai <szhai2@cs.rochester.edu> Tested-by: Sofia Trinh <sofia.trinh@edi.works> Tested-by: Vaibhav Jain <vaibhav@linux.ibm.com> Cc: Andi Kleen <ak@linux.intel.com> Cc: Aneesh Kumar K.V <aneesh.kumar@linux.ibm.com> Cc: Barry Song <baohua@kernel.org> Cc: Catalin Marinas <catalin.marinas@arm.com> Cc: Dave Hansen <dave.hansen@linux.intel.com> Cc: Hillf Danton <hdanton@sina.com> Cc: Jens Axboe <axboe@kernel.dk> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Jonathan Corbet <corbet@lwn.net> Cc: Linus Torvalds <torvalds@linux-foundation.org> Cc: Matthew Wilcox <willy@infradead.org> Cc: Mel Gorman <mgorman@suse.de> Cc: Miaohe Lin <linmiaohe@huawei.com> Cc: Michael Larabel <Michael@MichaelLarabel.com> Cc: Michal Hocko <mhocko@kernel.org> Cc: Mike Rapoport <rppt@kernel.org> Cc: Mike Rapoport <rppt@linux.ibm.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Qi Zheng <zhengqi.arch@bytedance.com> Cc: Tejun Heo <tj@kernel.org> Cc: Vlastimil Babka <vbabka@suse.cz> Cc: Will Deacon <will@kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2022-09-18 01:00:02 -07:00
mm: multi-gen LRU: kill switch Add /sys/kernel/mm/lru_gen/enabled as a kill switch. Components that can be disabled include: 0x0001: the multi-gen LRU core 0x0002: walking page table, when arch_has_hw_pte_young() returns true 0x0004: clearing the accessed bit in non-leaf PMD entries, when CONFIG_ARCH_HAS_NONLEAF_PMD_YOUNG=y [yYnN]: apply to all the components above E.g., echo y >/sys/kernel/mm/lru_gen/enabled cat /sys/kernel/mm/lru_gen/enabled 0x0007 echo 5 >/sys/kernel/mm/lru_gen/enabled cat /sys/kernel/mm/lru_gen/enabled 0x0005 NB: the page table walks happen on the scale of seconds under heavy memory pressure, in which case the mmap_lock contention is a lesser concern, compared with the LRU lock contention and the I/O congestion. So far the only well-known case of the mmap_lock contention happens on Android, due to Scudo [1] which allocates several thousand VMAs for merely a few hundred MBs. The SPF and the Maple Tree also have provided their own assessments [2][3]. However, if walking page tables does worsen the mmap_lock contention, the kill switch can be used to disable it. In this case the multi-gen LRU will suffer a minor performance degradation, as shown previously. Clearing the accessed bit in non-leaf PMD entries can also be disabled, since this behavior was not tested on x86 varieties other than Intel and AMD. [1] https://source.android.com/devices/tech/debug/scudo [2] https://lore.kernel.org/r/20220128131006.67712-1-michel@lespinasse.org/ [3] https://lore.kernel.org/r/20220426150616.3937571-1-Liam.Howlett@oracle.com/ Link: https://lkml.kernel.org/r/20220918080010.2920238-11-yuzhao@google.com Signed-off-by: Yu Zhao <yuzhao@google.com> Acked-by: Brian Geffon <bgeffon@google.com> Acked-by: Jan Alexander Steffens (heftig) <heftig@archlinux.org> Acked-by: Oleksandr Natalenko <oleksandr@natalenko.name> Acked-by: Steven Barrett <steven@liquorix.net> Acked-by: Suleiman Souhlal <suleiman@google.com> Tested-by: Daniel Byrne <djbyrne@mtu.edu> Tested-by: Donald Carr <d@chaos-reins.com> Tested-by: Holger Hoffstätte <holger@applied-asynchrony.com> Tested-by: Konstantin Kharlamov <Hi-Angel@yandex.ru> Tested-by: Shuang Zhai <szhai2@cs.rochester.edu> Tested-by: Sofia Trinh <sofia.trinh@edi.works> Tested-by: Vaibhav Jain <vaibhav@linux.ibm.com> Cc: Andi Kleen <ak@linux.intel.com> Cc: Aneesh Kumar K.V <aneesh.kumar@linux.ibm.com> Cc: Barry Song <baohua@kernel.org> Cc: Catalin Marinas <catalin.marinas@arm.com> Cc: Dave Hansen <dave.hansen@linux.intel.com> Cc: Hillf Danton <hdanton@sina.com> Cc: Jens Axboe <axboe@kernel.dk> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Jonathan Corbet <corbet@lwn.net> Cc: Linus Torvalds <torvalds@linux-foundation.org> Cc: Matthew Wilcox <willy@infradead.org> Cc: Mel Gorman <mgorman@suse.de> Cc: Miaohe Lin <linmiaohe@huawei.com> Cc: Michael Larabel <Michael@MichaelLarabel.com> Cc: Michal Hocko <mhocko@kernel.org> Cc: Mike Rapoport <rppt@kernel.org> Cc: Mike Rapoport <rppt@linux.ibm.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Qi Zheng <zhengqi.arch@bytedance.com> Cc: Tejun Heo <tj@kernel.org> Cc: Vlastimil Babka <vbabka@suse.cz> Cc: Will Deacon <will@kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2022-09-18 01:00:07 -07:00
config LRU_GEN_ENABLED
bool "Enable by default"
depends on LRU_GEN
help
This option enables the multi-gen LRU by default.
mm: multi-gen LRU: minimal implementation To avoid confusion, the terms "promotion" and "demotion" will be applied to the multi-gen LRU, as a new convention; the terms "activation" and "deactivation" will be applied to the active/inactive LRU, as usual. The aging produces young generations. Given an lruvec, it increments max_seq when max_seq-min_seq+1 approaches MIN_NR_GENS. The aging promotes hot pages to the youngest generation when it finds them accessed through page tables; the demotion of cold pages happens consequently when it increments max_seq. Promotion in the aging path does not involve any LRU list operations, only the updates of the gen counter and lrugen->nr_pages[]; demotion, unless as the result of the increment of max_seq, requires LRU list operations, e.g., lru_deactivate_fn(). The aging has the complexity O(nr_hot_pages), since it is only interested in hot pages. The eviction consumes old generations. Given an lruvec, it increments min_seq when lrugen->lists[] indexed by min_seq%MAX_NR_GENS becomes empty. A feedback loop modeled after the PID controller monitors refaults over anon and file types and decides which type to evict when both types are available from the same generation. The protection of pages accessed multiple times through file descriptors takes place in the eviction path. Each generation is divided into multiple tiers. A page accessed N times through file descriptors is in tier order_base_2(N). Tiers do not have dedicated lrugen->lists[], only bits in folio->flags. The aforementioned feedback loop also monitors refaults over all tiers and decides when to protect pages in which tiers (N>1), using the first tier (N=0,1) as a baseline. The first tier contains single-use unmapped clean pages, which are most likely the best choices. In contrast to promotion in the aging path, the protection of a page in the eviction path is achieved by moving this page to the next generation, i.e., min_seq+1, if the feedback loop decides so. This approach has the following advantages: 1. It removes the cost of activation in the buffered access path by inferring whether pages accessed multiple times through file descriptors are statistically hot and thus worth protecting in the eviction path. 2. It takes pages accessed through page tables into account and avoids overprotecting pages accessed multiple times through file descriptors. (Pages accessed through page tables are in the first tier, since N=0.) 3. More tiers provide better protection for pages accessed more than twice through file descriptors, when under heavy buffered I/O workloads. Server benchmark results: Single workload: fio (buffered I/O): +[30, 32]% IOPS BW 5.19-rc1: 2673k 10.2GiB/s patch1-6: 3491k 13.3GiB/s Single workload: memcached (anon): -[4, 6]% Ops/sec KB/sec 5.19-rc1: 1161501.04 45177.25 patch1-6: 1106168.46 43025.04 Configurations: CPU: two Xeon 6154 Mem: total 256G Node 1 was only used as a ram disk to reduce the variance in the results. patch drivers/block/brd.c <<EOF 99,100c99,100 < gfp_flags = GFP_NOIO | __GFP_ZERO | __GFP_HIGHMEM; < page = alloc_page(gfp_flags); --- > gfp_flags = GFP_NOIO | __GFP_ZERO | __GFP_HIGHMEM | __GFP_THISNODE; > page = alloc_pages_node(1, gfp_flags, 0); EOF cat >>/etc/systemd/system.conf <<EOF CPUAffinity=numa NUMAPolicy=bind NUMAMask=0 EOF cat >>/etc/memcached.conf <<EOF -m 184320 -s /var/run/memcached/memcached.sock -a 0766 -t 36 -B binary EOF cat fio.sh modprobe brd rd_nr=1 rd_size=113246208 swapoff -a mkfs.ext4 /dev/ram0 mount -t ext4 /dev/ram0 /mnt mkdir /sys/fs/cgroup/user.slice/test echo 38654705664 >/sys/fs/cgroup/user.slice/test/memory.max echo $$ >/sys/fs/cgroup/user.slice/test/cgroup.procs fio -name=mglru --numjobs=72 --directory=/mnt --size=1408m \ --buffered=1 --ioengine=io_uring --iodepth=128 \ --iodepth_batch_submit=32 --iodepth_batch_complete=32 \ --rw=randread --random_distribution=random --norandommap \ --time_based --ramp_time=10m --runtime=5m --group_reporting cat memcached.sh modprobe brd rd_nr=1 rd_size=113246208 swapoff -a mkswap /dev/ram0 swapon /dev/ram0 memtier_benchmark -S /var/run/memcached/memcached.sock \ -P memcache_binary -n allkeys --key-minimum=1 \ --key-maximum=65000000 --key-pattern=P:P -c 1 -t 36 \ --ratio 1:0 --pipeline 8 -d 2000 memtier_benchmark -S /var/run/memcached/memcached.sock \ -P memcache_binary -n allkeys --key-minimum=1 \ --key-maximum=65000000 --key-pattern=R:R -c 1 -t 36 \ --ratio 0:1 --pipeline 8 --randomize --distinct-client-seed Client benchmark results: kswapd profiles: 5.19-rc1 40.33% page_vma_mapped_walk (overhead) 21.80% lzo1x_1_do_compress (real work) 7.53% do_raw_spin_lock 3.95% _raw_spin_unlock_irq 2.52% vma_interval_tree_iter_next 2.37% folio_referenced_one 2.28% vma_interval_tree_subtree_search 1.97% anon_vma_interval_tree_iter_first 1.60% ptep_clear_flush 1.06% __zram_bvec_write patch1-6 39.03% lzo1x_1_do_compress (real work) 18.47% page_vma_mapped_walk (overhead) 6.74% _raw_spin_unlock_irq 3.97% do_raw_spin_lock 2.49% ptep_clear_flush 2.48% anon_vma_interval_tree_iter_first 1.92% folio_referenced_one 1.88% __zram_bvec_write 1.48% memmove 1.31% vma_interval_tree_iter_next Configurations: CPU: single Snapdragon 7c Mem: total 4G ChromeOS MemoryPressure [1] [1] https://chromium.googlesource.com/chromiumos/platform/tast-tests/ Link: https://lkml.kernel.org/r/20220918080010.2920238-7-yuzhao@google.com Signed-off-by: Yu Zhao <yuzhao@google.com> Acked-by: Brian Geffon <bgeffon@google.com> Acked-by: Jan Alexander Steffens (heftig) <heftig@archlinux.org> Acked-by: Oleksandr Natalenko <oleksandr@natalenko.name> Acked-by: Steven Barrett <steven@liquorix.net> Acked-by: Suleiman Souhlal <suleiman@google.com> Tested-by: Daniel Byrne <djbyrne@mtu.edu> Tested-by: Donald Carr <d@chaos-reins.com> Tested-by: Holger Hoffstätte <holger@applied-asynchrony.com> Tested-by: Konstantin Kharlamov <Hi-Angel@yandex.ru> Tested-by: Shuang Zhai <szhai2@cs.rochester.edu> Tested-by: Sofia Trinh <sofia.trinh@edi.works> Tested-by: Vaibhav Jain <vaibhav@linux.ibm.com> Cc: Andi Kleen <ak@linux.intel.com> Cc: Aneesh Kumar K.V <aneesh.kumar@linux.ibm.com> Cc: Barry Song <baohua@kernel.org> Cc: Catalin Marinas <catalin.marinas@arm.com> Cc: Dave Hansen <dave.hansen@linux.intel.com> Cc: Hillf Danton <hdanton@sina.com> Cc: Jens Axboe <axboe@kernel.dk> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Jonathan Corbet <corbet@lwn.net> Cc: Linus Torvalds <torvalds@linux-foundation.org> Cc: Matthew Wilcox <willy@infradead.org> Cc: Mel Gorman <mgorman@suse.de> Cc: Miaohe Lin <linmiaohe@huawei.com> Cc: Michael Larabel <Michael@MichaelLarabel.com> Cc: Michal Hocko <mhocko@kernel.org> Cc: Mike Rapoport <rppt@kernel.org> Cc: Mike Rapoport <rppt@linux.ibm.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Qi Zheng <zhengqi.arch@bytedance.com> Cc: Tejun Heo <tj@kernel.org> Cc: Vlastimil Babka <vbabka@suse.cz> Cc: Will Deacon <will@kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2022-09-18 01:00:03 -07:00
config LRU_GEN_STATS
bool "Full stats for debugging"
depends on LRU_GEN
help
Do not enable this option unless you plan to look at historical stats
from evicted generations for debugging purpose.
This option has a per-memcg and per-node memory overhead.
config LRU_GEN_WALKS_MMU
def_bool y
depends on LRU_GEN && ARCH_HAS_HW_PTE_YOUNG
mm: multi-gen LRU: minimal implementation To avoid confusion, the terms "promotion" and "demotion" will be applied to the multi-gen LRU, as a new convention; the terms "activation" and "deactivation" will be applied to the active/inactive LRU, as usual. The aging produces young generations. Given an lruvec, it increments max_seq when max_seq-min_seq+1 approaches MIN_NR_GENS. The aging promotes hot pages to the youngest generation when it finds them accessed through page tables; the demotion of cold pages happens consequently when it increments max_seq. Promotion in the aging path does not involve any LRU list operations, only the updates of the gen counter and lrugen->nr_pages[]; demotion, unless as the result of the increment of max_seq, requires LRU list operations, e.g., lru_deactivate_fn(). The aging has the complexity O(nr_hot_pages), since it is only interested in hot pages. The eviction consumes old generations. Given an lruvec, it increments min_seq when lrugen->lists[] indexed by min_seq%MAX_NR_GENS becomes empty. A feedback loop modeled after the PID controller monitors refaults over anon and file types and decides which type to evict when both types are available from the same generation. The protection of pages accessed multiple times through file descriptors takes place in the eviction path. Each generation is divided into multiple tiers. A page accessed N times through file descriptors is in tier order_base_2(N). Tiers do not have dedicated lrugen->lists[], only bits in folio->flags. The aforementioned feedback loop also monitors refaults over all tiers and decides when to protect pages in which tiers (N>1), using the first tier (N=0,1) as a baseline. The first tier contains single-use unmapped clean pages, which are most likely the best choices. In contrast to promotion in the aging path, the protection of a page in the eviction path is achieved by moving this page to the next generation, i.e., min_seq+1, if the feedback loop decides so. This approach has the following advantages: 1. It removes the cost of activation in the buffered access path by inferring whether pages accessed multiple times through file descriptors are statistically hot and thus worth protecting in the eviction path. 2. It takes pages accessed through page tables into account and avoids overprotecting pages accessed multiple times through file descriptors. (Pages accessed through page tables are in the first tier, since N=0.) 3. More tiers provide better protection for pages accessed more than twice through file descriptors, when under heavy buffered I/O workloads. Server benchmark results: Single workload: fio (buffered I/O): +[30, 32]% IOPS BW 5.19-rc1: 2673k 10.2GiB/s patch1-6: 3491k 13.3GiB/s Single workload: memcached (anon): -[4, 6]% Ops/sec KB/sec 5.19-rc1: 1161501.04 45177.25 patch1-6: 1106168.46 43025.04 Configurations: CPU: two Xeon 6154 Mem: total 256G Node 1 was only used as a ram disk to reduce the variance in the results. patch drivers/block/brd.c <<EOF 99,100c99,100 < gfp_flags = GFP_NOIO | __GFP_ZERO | __GFP_HIGHMEM; < page = alloc_page(gfp_flags); --- > gfp_flags = GFP_NOIO | __GFP_ZERO | __GFP_HIGHMEM | __GFP_THISNODE; > page = alloc_pages_node(1, gfp_flags, 0); EOF cat >>/etc/systemd/system.conf <<EOF CPUAffinity=numa NUMAPolicy=bind NUMAMask=0 EOF cat >>/etc/memcached.conf <<EOF -m 184320 -s /var/run/memcached/memcached.sock -a 0766 -t 36 -B binary EOF cat fio.sh modprobe brd rd_nr=1 rd_size=113246208 swapoff -a mkfs.ext4 /dev/ram0 mount -t ext4 /dev/ram0 /mnt mkdir /sys/fs/cgroup/user.slice/test echo 38654705664 >/sys/fs/cgroup/user.slice/test/memory.max echo $$ >/sys/fs/cgroup/user.slice/test/cgroup.procs fio -name=mglru --numjobs=72 --directory=/mnt --size=1408m \ --buffered=1 --ioengine=io_uring --iodepth=128 \ --iodepth_batch_submit=32 --iodepth_batch_complete=32 \ --rw=randread --random_distribution=random --norandommap \ --time_based --ramp_time=10m --runtime=5m --group_reporting cat memcached.sh modprobe brd rd_nr=1 rd_size=113246208 swapoff -a mkswap /dev/ram0 swapon /dev/ram0 memtier_benchmark -S /var/run/memcached/memcached.sock \ -P memcache_binary -n allkeys --key-minimum=1 \ --key-maximum=65000000 --key-pattern=P:P -c 1 -t 36 \ --ratio 1:0 --pipeline 8 -d 2000 memtier_benchmark -S /var/run/memcached/memcached.sock \ -P memcache_binary -n allkeys --key-minimum=1 \ --key-maximum=65000000 --key-pattern=R:R -c 1 -t 36 \ --ratio 0:1 --pipeline 8 --randomize --distinct-client-seed Client benchmark results: kswapd profiles: 5.19-rc1 40.33% page_vma_mapped_walk (overhead) 21.80% lzo1x_1_do_compress (real work) 7.53% do_raw_spin_lock 3.95% _raw_spin_unlock_irq 2.52% vma_interval_tree_iter_next 2.37% folio_referenced_one 2.28% vma_interval_tree_subtree_search 1.97% anon_vma_interval_tree_iter_first 1.60% ptep_clear_flush 1.06% __zram_bvec_write patch1-6 39.03% lzo1x_1_do_compress (real work) 18.47% page_vma_mapped_walk (overhead) 6.74% _raw_spin_unlock_irq 3.97% do_raw_spin_lock 2.49% ptep_clear_flush 2.48% anon_vma_interval_tree_iter_first 1.92% folio_referenced_one 1.88% __zram_bvec_write 1.48% memmove 1.31% vma_interval_tree_iter_next Configurations: CPU: single Snapdragon 7c Mem: total 4G ChromeOS MemoryPressure [1] [1] https://chromium.googlesource.com/chromiumos/platform/tast-tests/ Link: https://lkml.kernel.org/r/20220918080010.2920238-7-yuzhao@google.com Signed-off-by: Yu Zhao <yuzhao@google.com> Acked-by: Brian Geffon <bgeffon@google.com> Acked-by: Jan Alexander Steffens (heftig) <heftig@archlinux.org> Acked-by: Oleksandr Natalenko <oleksandr@natalenko.name> Acked-by: Steven Barrett <steven@liquorix.net> Acked-by: Suleiman Souhlal <suleiman@google.com> Tested-by: Daniel Byrne <djbyrne@mtu.edu> Tested-by: Donald Carr <d@chaos-reins.com> Tested-by: Holger Hoffstätte <holger@applied-asynchrony.com> Tested-by: Konstantin Kharlamov <Hi-Angel@yandex.ru> Tested-by: Shuang Zhai <szhai2@cs.rochester.edu> Tested-by: Sofia Trinh <sofia.trinh@edi.works> Tested-by: Vaibhav Jain <vaibhav@linux.ibm.com> Cc: Andi Kleen <ak@linux.intel.com> Cc: Aneesh Kumar K.V <aneesh.kumar@linux.ibm.com> Cc: Barry Song <baohua@kernel.org> Cc: Catalin Marinas <catalin.marinas@arm.com> Cc: Dave Hansen <dave.hansen@linux.intel.com> Cc: Hillf Danton <hdanton@sina.com> Cc: Jens Axboe <axboe@kernel.dk> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Jonathan Corbet <corbet@lwn.net> Cc: Linus Torvalds <torvalds@linux-foundation.org> Cc: Matthew Wilcox <willy@infradead.org> Cc: Mel Gorman <mgorman@suse.de> Cc: Miaohe Lin <linmiaohe@huawei.com> Cc: Michael Larabel <Michael@MichaelLarabel.com> Cc: Michal Hocko <mhocko@kernel.org> Cc: Mike Rapoport <rppt@kernel.org> Cc: Mike Rapoport <rppt@linux.ibm.com> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Qi Zheng <zhengqi.arch@bytedance.com> Cc: Tejun Heo <tj@kernel.org> Cc: Vlastimil Babka <vbabka@suse.cz> Cc: Will Deacon <will@kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2022-09-18 01:00:03 -07:00
# }
mm: introduce CONFIG_PER_VMA_LOCK Patch series "Per-VMA locks", v4. LWN article describing the feature: https://lwn.net/Articles/906852/ Per-vma locks idea that was discussed during SPF [1] discussion at LSF/MM last year [2], which concluded with suggestion that “a reader/writer semaphore could be put into the VMA itself; that would have the effect of using the VMA as a sort of range lock. There would still be contention at the VMA level, but it would be an improvement.” This patchset implements this suggested approach. When handling page faults we lookup the VMA that contains the faulting page under RCU protection and try to acquire its lock. If that fails we fall back to using mmap_lock, similar to how SPF handled this situation. One notable way the implementation deviates from the proposal is the way VMAs are read-locked. During some of mm updates, multiple VMAs need to be locked until the end of the update (e.g. vma_merge, split_vma, etc). Tracking all the locked VMAs, avoiding recursive locks, figuring out when it's safe to unlock previously locked VMAs would make the code more complex. So, instead of the usual lock/unlock pattern, the proposed solution marks a VMA as locked and provides an efficient way to: 1. Identify locked VMAs. 2. Unlock all locked VMAs in bulk. We also postpone unlocking the locked VMAs until the end of the update, when we do mmap_write_unlock. Potentially this keeps a VMA locked for longer than is absolutely necessary but it results in a big reduction of code complexity. Read-locking a VMA is done using two sequence numbers - one in the vm_area_struct and one in the mm_struct. VMA is considered read-locked when these sequence numbers are equal. To read-lock a VMA we set the sequence number in vm_area_struct to be equal to the sequence number in mm_struct. To unlock all VMAs we increment mm_struct's seq number. This allows for an efficient way to track locked VMAs and to drop the locks on all VMAs at the end of the update. The patchset implements per-VMA locking only for anonymous pages which are not in swap and avoids userfaultfs as their implementation is more complex. Additional support for file-back page faults, swapped and user pages can be added incrementally. Performance benchmarks show similar although slightly smaller benefits as with SPF patchset (~75% of SPF benefits). Still, with lower complexity this approach might be more desirable. Since RFC was posted in September 2022, two separate Google teams outside of Android evaluated the patchset and confirmed positive results. Here are the known usecases when per-VMA locks show benefits: Android: Apps with high number of threads (~100) launch times improve by up to 20%. Each thread mmaps several areas upon startup (Stack and Thread-local storage (TLS), thread signal stack, indirect ref table), which requires taking mmap_lock in write mode. Page faults take mmap_lock in read mode. During app launch, both thread creation and page faults establishing the active workinget are happening in parallel and that causes lock contention between mm writers and readers even if updates and page faults are happening in different VMAs. Per-vma locks prevent this contention by providing more granular lock. Google Fibers: We have several dynamically sized thread pools that spawn new threads under increased load and reduce their number when idling. For example, Google's in-process scheduling/threading framework, UMCG/Fibers, is backed by such a thread pool. When idling, only a small number of idle worker threads are available; when a spike of incoming requests arrive, each request is handled in its own "fiber", which is a work item posted onto a UMCG worker thread; quite often these spikes lead to a number of new threads spawning. Each new thread needs to allocate and register an RSEQ section on its TLS, then register itself with the kernel as a UMCG worker thread, and only after that it can be considered by the in-process UMCG/Fiber scheduler as available to do useful work. In short, during an incoming workload spike new threads have to be spawned, and they perform several syscalls (RSEQ registration, UMCG worker registration, memory allocations) before they can actually start doing useful work. Removing any bottlenecks on this thread startup path will greatly improve our services' latencies when faced with request/workload spikes. At high scale, mmap_lock contention during thread creation and stack page faults leads to user-visible multi-second serving latencies in a similar pattern to Android app startup. Per-VMA locking patchset has been run successfully in limited experiments with user-facing production workloads. In these experiments, we observed that the peak thread creation rate was high enough that thread creation is no longer a bottleneck. TCP zerocopy receive: From the point of view of TCP zerocopy receive, the per-vma lock patch is massively beneficial. In today's implementation, a process with N threads where N - 1 are performing zerocopy receive and 1 thread is performing madvise() with the write lock taken (e.g. needs to change vm_flags) will result in all N -1 receive threads blocking until the madvise is done. Conversely, on a busy process receiving a lot of data, an madvise operation that does need to take the mmap lock in write mode will need to wait for all of the receives to be done - a lose:lose proposition. Per-VMA locking _removes_ by definition this source of contention entirely. There are other benefits for receive as well, chiefly a reduction in cacheline bouncing across receiving threads for locking/unlocking the single mmap lock. On an RPC style synthetic workload with 4KB RPCs: 1a) The find+lock+unlock VMA path in the base case, without the per-vma lock patchset, is about 0.7% of cycles as measured by perf. 1b) mmap_read_lock + mmap_read_unlock in the base case is about 0.5% cycles overall - most of this is within the TCP read hotpath (a small fraction is 'other' usage in the system). 2a) The find+lock+unlock VMA path, with the per-vma patchset and a trivial patch written to take advantage of it in TCP, is about 0.4% of cycles (down from 0.7% above) 2b) mmap_read_lock + mmap_read_unlock in the per-vma patchset is < 0.1% cycles and is out of the TCP read hotpath entirely (down from 0.5% before, the remaining usage is the 'other' usage in the system). So, in addition to entirely removing an onerous source of contention, it also reduces the CPU cycles of TCP receive zerocopy by about 0.5%+ (compared to overall cycles in perf) for the 'small' RPC scenario. In https://lkml.kernel.org/r/87fsaqouyd.fsf_-_@stealth, Punit demonstrated throughput improvements of as much as 188% from this patchset. This patch (of 25): This configuration variable will be used to build the support for VMA locking during page fault handling. This is enabled on supported architectures with SMP and MMU set. The architecture support is needed since the page fault handler is called from the architecture's page faulting code which needs modifications to handle faults under VMA lock. Link: https://lkml.kernel.org/r/20230227173632.3292573-1-surenb@google.com Link: https://lkml.kernel.org/r/20230227173632.3292573-10-surenb@google.com Signed-off-by: Suren Baghdasaryan <surenb@google.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2023-02-27 10:36:08 -07:00
config ARCH_SUPPORTS_PER_VMA_LOCK
def_bool n
config PER_VMA_LOCK
def_bool y
depends on ARCH_SUPPORTS_PER_VMA_LOCK && MMU && SMP
help
Allow per-vma locking during page fault handling.
This feature allows locking each virtual memory area separately when
handling page faults instead of taking mmap_lock.
mm: introduce new 'lock_mm_and_find_vma()' page fault helper .. and make x86 use it. This basically extracts the existing x86 "find and expand faulting vma" code, but extends it to also take the mmap lock for writing in case we actually do need to expand the vma. We've historically short-circuited that case, and have some rather ugly special logic to serialize the stack segment expansion (since we only hold the mmap lock for reading) that doesn't match the normal VM locking. That slight violation of locking worked well, right up until it didn't: the maple tree code really does want proper locking even for simple extension of an existing vma. So extract the code for "look up the vma of the fault" from x86, fix it up to do the necessary write locking, and make it available as a helper function for other architectures that can use the common helper. Note: I say "common helper", but it really only handles the normal stack-grows-down case. Which is all architectures except for PA-RISC and IA64. So some rare architectures can't use the helper, but if they care they'll just need to open-code this logic. It's also worth pointing out that this code really would like to have an optimistic "mmap_upgrade_trylock()" to make it quicker to go from a read-lock (for the common case) to taking the write lock (for having to extend the vma) in the normal single-threaded situation where there is no other locking activity. But that _is_ all the very uncommon special case, so while it would be nice to have such an operation, it probably doesn't matter in reality. I did put in the skeleton code for such a possible future expansion, even if it only acts as pseudo-documentation for what we're doing. Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2023-06-15 15:17:36 -07:00
config LOCK_MM_AND_FIND_VMA
bool
depends on !STACK_GROWSUP
config IOMMU_MM_DATA
bool
config EXECMEM
bool
mm: introduce numa_memblks Move code dealing with numa_memblks from arch/x86 to mm/ and add Kconfig options to let x86 select it in its Kconfig. This code will be later reused by arch_numa. No functional changes. Link: https://lkml.kernel.org/r/20240807064110.1003856-18-rppt@kernel.org Signed-off-by: Mike Rapoport (Microsoft) <rppt@kernel.org> Tested-by: Zi Yan <ziy@nvidia.com> # for x86_64 and arm64 Reviewed-by: Jonathan Cameron <Jonathan.Cameron@huawei.com> Tested-by: Jonathan Cameron <Jonathan.Cameron@huawei.com> [arm64 + CXL via QEMU] Acked-by: Dan Williams <dan.j.williams@intel.com> Acked-by: David Hildenbrand <david@redhat.com> Cc: Alexander Gordeev <agordeev@linux.ibm.com> Cc: Andreas Larsson <andreas@gaisler.com> Cc: Arnd Bergmann <arnd@arndb.de> Cc: Borislav Petkov <bp@alien8.de> Cc: Catalin Marinas <catalin.marinas@arm.com> Cc: Christophe Leroy <christophe.leroy@csgroup.eu> Cc: Dave Hansen <dave.hansen@linux.intel.com> Cc: Davidlohr Bueso <dave@stgolabs.net> Cc: David S. Miller <davem@davemloft.net> Cc: Greg Kroah-Hartman <gregkh@linuxfoundation.org> Cc: Heiko Carstens <hca@linux.ibm.com> Cc: Huacai Chen <chenhuacai@kernel.org> Cc: Ingo Molnar <mingo@redhat.com> Cc: Jiaxun Yang <jiaxun.yang@flygoat.com> Cc: John Paul Adrian Glaubitz <glaubitz@physik.fu-berlin.de> Cc: Jonathan Corbet <corbet@lwn.net> Cc: Michael Ellerman <mpe@ellerman.id.au> Cc: Palmer Dabbelt <palmer@dabbelt.com> Cc: Rafael J. Wysocki <rafael@kernel.org> Cc: Rob Herring (Arm) <robh@kernel.org> Cc: Samuel Holland <samuel.holland@sifive.com> Cc: Thomas Bogendoerfer <tsbogend@alpha.franken.de> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Vasily Gorbik <gor@linux.ibm.com> Cc: Will Deacon <will@kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2024-08-06 23:41:01 -07:00
config NUMA_MEMBLKS
bool
mm: introduce numa_emulation Move numa_emulation code from arch/x86 to mm/numa_emulation.c This code will be later reused by arch_numa. No functional changes. Link: https://lkml.kernel.org/r/20240807064110.1003856-20-rppt@kernel.org Signed-off-by: Mike Rapoport (Microsoft) <rppt@kernel.org> Tested-by: Zi Yan <ziy@nvidia.com> # for x86_64 and arm64 Reviewed-by: Jonathan Cameron <Jonathan.Cameron@huawei.com> Tested-by: Jonathan Cameron <Jonathan.Cameron@huawei.com> [arm64 + CXL via QEMU] Acked-by: Dan Williams <dan.j.williams@intel.com> Cc: Alexander Gordeev <agordeev@linux.ibm.com> Cc: Andreas Larsson <andreas@gaisler.com> Cc: Arnd Bergmann <arnd@arndb.de> Cc: Borislav Petkov <bp@alien8.de> Cc: Catalin Marinas <catalin.marinas@arm.com> Cc: Christophe Leroy <christophe.leroy@csgroup.eu> Cc: Dave Hansen <dave.hansen@linux.intel.com> Cc: David Hildenbrand <david@redhat.com> Cc: Davidlohr Bueso <dave@stgolabs.net> Cc: David S. Miller <davem@davemloft.net> Cc: Greg Kroah-Hartman <gregkh@linuxfoundation.org> Cc: Heiko Carstens <hca@linux.ibm.com> Cc: Huacai Chen <chenhuacai@kernel.org> Cc: Ingo Molnar <mingo@redhat.com> Cc: Jiaxun Yang <jiaxun.yang@flygoat.com> Cc: John Paul Adrian Glaubitz <glaubitz@physik.fu-berlin.de> Cc: Jonathan Corbet <corbet@lwn.net> Cc: Michael Ellerman <mpe@ellerman.id.au> Cc: Palmer Dabbelt <palmer@dabbelt.com> Cc: Rafael J. Wysocki <rafael@kernel.org> Cc: Rob Herring (Arm) <robh@kernel.org> Cc: Samuel Holland <samuel.holland@sifive.com> Cc: Thomas Bogendoerfer <tsbogend@alpha.franken.de> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Vasily Gorbik <gor@linux.ibm.com> Cc: Will Deacon <will@kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
2024-08-06 23:41:03 -07:00
config NUMA_EMU
bool "NUMA emulation"
depends on NUMA_MEMBLKS
help
Enable NUMA emulation. A flat machine will be split
into virtual nodes when booted with "numa=fake=N", where N is the
number of nodes. This is only useful for debugging.
mm: introduce Data Access MONitor (DAMON) Patch series "Introduce Data Access MONitor (DAMON)", v34. Introduction ============ DAMON is a data access monitoring framework for the Linux kernel. The core mechanisms of DAMON called 'region based sampling' and 'adaptive regions adjustment' (refer to 'mechanisms.rst' in the 11th patch of this patchset for the detail) make it - accurate (The monitored information is useful for DRAM level memory management. It might not appropriate for Cache-level accuracy, though.), - light-weight (The monitoring overhead is low enough to be applied online while making no impact on the performance of the target workloads.), and - scalable (the upper-bound of the instrumentation overhead is controllable regardless of the size of target workloads.). Using this framework, therefore, several memory management mechanisms such as reclamation and THP can be optimized to aware real data access patterns. Experimental access pattern aware memory management optimization works that incurring high instrumentation overhead will be able to have another try. Though DAMON is for kernel subsystems, it can be easily exposed to the user space by writing a DAMON-wrapper kernel subsystem. Then, user space users who have some special workloads will be able to write personalized tools or applications for deeper understanding and specialized optimizations of their systems. DAMON is also merged in two public Amazon Linux kernel trees that based on v5.4.y[1] and v5.10.y[2]. [1] https://github.com/amazonlinux/linux/tree/amazon-5.4.y/master/mm/damon [2] https://github.com/amazonlinux/linux/tree/amazon-5.10.y/master/mm/damon The userspace tool[1] is available, released under GPLv2, and actively being maintained. I am also planning to implement another basic user interface in perf[2]. Also, the basic test suite for DAMON is available under GPLv2[3]. [1] https://github.com/awslabs/damo [2] https://lore.kernel.org/linux-mm/20210107120729.22328-1-sjpark@amazon.com/ [3] https://github.com/awslabs/damon-tests Long-term Plan -------------- DAMON is a part of a project called Data Access-aware Operating System (DAOS). As the name implies, I want to improve the performance and efficiency of systems using fine-grained data access patterns. The optimizations are for both kernel and user spaces. I will therefore modify or create kernel subsystems, export some of those to user space and implement user space library / tools. Below shows the layers and components for the project. --------------------------------------------------------------------------- Primitives: PTE Accessed bit, PG_idle, rmap, (Intel CMT), ... Framework: DAMON Features: DAMOS, virtual addr, physical addr, ... Applications: DAMON-debugfs, (DARC), ... ^^^^^^^^^^^^^^^^^^^^^^^ KERNEL SPACE ^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^ Raw Interface: debugfs, (sysfs), (damonfs), tracepoints, (sys_damon), ... vvvvvvvvvvvvvvvvvvvvvvv USER SPACE vvvvvvvvvvvvvvvvvvvvvvvvvvvvvvvv Library: (libdamon), ... Tools: DAMO, (perf), ... --------------------------------------------------------------------------- The components in parentheses or marked as '...' are not implemented yet but in the future plan. IOW, those are the TODO tasks of DAOS project. For more detail, please refer to the plans: https://lore.kernel.org/linux-mm/20201202082731.24828-1-sjpark@amazon.com/ Evaluations =========== We evaluated DAMON's overhead, monitoring quality and usefulness using 24 realistic workloads on my QEMU/KVM based virtual machine running a kernel that v24 DAMON patchset is applied. DAMON is lightweight. It increases system memory usage by 0.39% and slows target workloads down by 1.16%. DAMON is accurate and useful for memory management optimizations. An experimental DAMON-based operation scheme for THP, namely 'ethp', removes 76.15% of THP memory overheads while preserving 51.25% of THP speedup. Another experimental DAMON-based 'proactive reclamation' implementation, 'prcl', reduces 93.38% of residential sets and 23.63% of system memory footprint while incurring only 1.22% runtime overhead in the best case (parsec3/freqmine). NOTE that the experimental THP optimization and proactive reclamation are not for production but only for proof of concepts. Please refer to the official document[1] or "Documentation/admin-guide/mm: Add a document for DAMON" patch in this patchset for detailed evaluation setup and results. [1] https://damonitor.github.io/doc/html/latest-damon/admin-guide/mm/damon/eval.html Real-world User Story ===================== In summary, DAMON has used on production systems and proved its usefulness. DAMON as a profiler ------------------- We analyzed characteristics of a large scale production systems of our customers using DAMON. The systems utilize 70GB DRAM and 36 CPUs. From this, we were able to find interesting things below. There were obviously different access pattern under idle workload and active workload. Under the idle workload, it accessed large memory regions with low frequency, while the active workload accessed small memory regions with high freuqnecy. DAMON found a 7GB memory region that showing obviously high access frequency under the active workload. We believe this is the performance-effective working set and need to be protected. There was a 4KB memory region that showing highest access frequency under not only active but also idle workloads. We think this must be a hottest code section like thing that should never be paged out. For this analysis, DAMON used only 0.3-1% of single CPU time. Because we used recording-based analysis, it consumed about 3-12 MB of disk space per 20 minutes. This is only small amount of disk space, but we can further reduce the disk usage by using non-recording-based DAMON features. I'd like to argue that only DAMON can do such detailed analysis (finding 4KB highest region in 70GB memory) with the light overhead. DAMON as a system optimization tool ----------------------------------- We also found below potential performance problems on the systems and made DAMON-based solutions. The system doesn't want to make the workload suffer from the page reclamation and thus it utilizes enough DRAM but no swap device. However, we found the system is actively reclaiming file-backed pages, because the system has intensive file IO. The file IO turned out to be not performance critical for the workload, but the customer wanted to ensure performance critical file-backed pages like code section to not mistakenly be evicted. Using direct IO should or `mlock()` would be a straightforward solution, but modifying the user space code is not easy for the customer. Alternatively, we could use DAMON-based operation scheme[1]. By using it, we can ask DAMON to track access frequency of each region and make 'process_madvise(MADV_WILLNEED)[2]' call for regions having specific size and access frequency for a time interval. We also found the system is having high number of TLB misses. We tried 'always' THP enabled policy and it greatly reduced TLB misses, but the page reclamation also been more frequent due to the THP internal fragmentation caused memory bloat. We could try another DAMON-based operation scheme that applies 'MADV_HUGEPAGE' to memory regions having >=2MB size and high access frequency, while applying 'MADV_NOHUGEPAGE' to regions having <2MB size and low access frequency. We do not own the systems so we only reported the analysis results and possible optimization solutions to the customers. The customers satisfied about the analysis results and promised to try the optimization guides. [1] https://lore.kernel.org/linux-mm/20201006123931.5847-1-sjpark@amazon.com/ [2] https://lore.kernel.org/linux-api/20200622192900.22757-4-minchan@kernel.org/ Comparison with Idle Page Tracking ================================== Idle Page Tracking allows users to set and read idleness of pages using a bitmap file which represents each page with each bit of the file. One recommended usage of it is working set size detection. Users can do that by 1. find PFN of each page for workloads in interest, 2. set all the pages as idle by doing writes to the bitmap file, 3. wait until the workload accesses its working set, and 4. read the idleness of the pages again and count pages became not idle. NOTE: While Idle Page Tracking is for user space users, DAMON is primarily designed for kernel subsystems though it can easily exposed to the user space. Hence, this section only assumes such user space use of DAMON. For what use cases Idle Page Tracking would be better? ------------------------------------------------------ 1. Flexible usecases other than hotness monitoring. Because Idle Page Tracking allows users to control the primitive (Page idleness) by themselves, Idle Page Tracking users can do anything they want. Meanwhile, DAMON is primarily designed to monitor the hotness of each memory region. For this, DAMON asks users to provide sampling interval and aggregation interval. For the reason, there could be some use case that using Idle Page Tracking is simpler. 2. Physical memory monitoring. Idle Page Tracking receives PFN range as input, so natively supports physical memory monitoring. DAMON is designed to be extensible for multiple address spaces and use cases by implementing and using primitives for the given use case. Therefore, by theory, DAMON has no limitation in the type of target address space as long as primitives for the given address space exists. However, the default primitives introduced by this patchset supports only virtual address spaces. Therefore, for physical memory monitoring, you should implement your own primitives and use it, or simply use Idle Page Tracking. Nonetheless, RFC patchsets[1] for the physical memory address space primitives is already available. It also supports user memory same to Idle Page Tracking. [1] https://lore.kernel.org/linux-mm/20200831104730.28970-1-sjpark@amazon.com/ For what use cases DAMON is better? ----------------------------------- 1. Hotness Monitoring. Idle Page Tracking let users know only if a page frame is accessed or not. For hotness check, the user should write more code and use more memory. DAMON do that by itself. 2. Low Monitoring Overhead DAMON receives user's monitoring request with one step and then provide the results. So, roughly speaking, DAMON require only O(1) user/kernel context switches. In case of Idle Page Tracking, however, because the interface receives contiguous page frames, the number of user/kernel context switches increases as the monitoring target becomes complex and huge. As a result, the context switch overhead could be not negligible. Moreover, DAMON is born to handle with the monitoring overhead. Because the core mechanism is pure logical, Idle Page Tracking users might be able to implement the mechanism on their own, but it would be time consuming and the user/kernel context switching will still more frequent than that of DAMON. Also, the kernel subsystems cannot use the logic in this case. 3. Page granularity working set size detection. Until v22 of this patchset, this was categorized as the thing Idle Page Tracking could do better, because DAMON basically maintains additional metadata for each of the monitoring target regions. So, in the page granularity working set size detection use case, DAMON would incur (number of monitoring target pages * size of metadata) memory overhead. Size of the single metadata item is about 54 bytes, so assuming 4KB pages, about 1.3% of monitoring target pages will be additionally used. All essential metadata for Idle Page Tracking are embedded in 'struct page' and page table entries. Therefore, in this use case, only one counter variable for working set size accounting is required if Idle Page Tracking is used. There are more details to consider, but roughly speaking, this is true in most cases. However, the situation changed from v23. Now DAMON supports arbitrary types of monitoring targets, which don't use the metadata. Using that, DAMON can do the working set size detection with no additional space overhead but less user-kernel context switch. A first draft for the implementation of monitoring primitives for this usage is available in a DAMON development tree[1]. An RFC patchset for it based on this patchset will also be available soon. Since v24, the arbitrary type support is dropped from this patchset because this patchset doesn't introduce real use of the type. You can still get it from the DAMON development tree[2], though. [1] https://github.com/sjp38/linux/tree/damon/pgidle_hack [2] https://github.com/sjp38/linux/tree/damon/master 4. More future usecases While Idle Page Tracking has tight coupling with base primitives (PG_Idle and page table Accessed bits), DAMON is designed to be extensible for many use cases and address spaces. If you need some special address type or want to use special h/w access check primitives, you can write your own primitives for that and configure DAMON to use those. Therefore, if your use case could be changed a lot in future, using DAMON could be better. Can I use both Idle Page Tracking and DAMON? -------------------------------------------- Yes, though using them concurrently for overlapping memory regions could result in interference to each other. Nevertheless, such use case would be rare or makes no sense at all. Even in the case, the noise would bot be really significant. So, you can choose whatever you want depending on the characteristics of your use cases. More Information ================ We prepared a showcase web site[1] that you can get more information. There are - the official documentations[2], - the heatmap format dynamic access pattern of various realistic workloads for heap area[3], mmap()-ed area[4], and stack[5] area, - the dynamic working set size distribution[6] and chronological working set size changes[7], and - the latest performance test results[8]. [1] https://damonitor.github.io/_index [2] https://damonitor.github.io/doc/html/latest-damon [3] https://damonitor.github.io/test/result/visual/latest/rec.heatmap.0.png.html [4] https://damonitor.github.io/test/result/visual/latest/rec.heatmap.1.png.html [5] https://damonitor.github.io/test/result/visual/latest/rec.heatmap.2.png.html [6] https://damonitor.github.io/test/result/visual/latest/rec.wss_sz.png.html [7] https://damonitor.github.io/test/result/visual/latest/rec.wss_time.png.html [8] https://damonitor.github.io/test/result/perf/latest/html/index.html Baseline and Complete Git Trees =============================== The patches are based on the latest -mm tree, specifically v5.14-rc1-mmots-2021-07-15-18-47 of https://github.com/hnaz/linux-mm. You can also clone the complete git tree: $ git clone git://github.com/sjp38/linux -b damon/patches/v34 The web is also available: https://github.com/sjp38/linux/releases/tag/damon/patches/v34 Development Trees ----------------- There are a couple of trees for entire DAMON patchset series and features for future release. - For latest release: https://github.com/sjp38/linux/tree/damon/master - For next release: https://github.com/sjp38/linux/tree/damon/next Long-term Support Trees ----------------------- For people who want to test DAMON but using LTS kernels, there are another couple of trees based on two latest LTS kernels respectively and containing the 'damon/master' backports. - For v5.4.y: https://github.com/sjp38/linux/tree/damon/for-v5.4.y - For v5.10.y: https://github.com/sjp38/linux/tree/damon/for-v5.10.y Amazon Linux Kernel Trees ------------------------- DAMON is also merged in two public Amazon Linux kernel trees that based on v5.4.y[1] and v5.10.y[2]. [1] https://github.com/amazonlinux/linux/tree/amazon-5.4.y/master/mm/damon [2] https://github.com/amazonlinux/linux/tree/amazon-5.10.y/master/mm/damon Git Tree for Diff of Patches ============================ For easy review of diff between different versions of each patch, I prepared a git tree containing all versions of the DAMON patchset series: https://github.com/sjp38/damon-patches You can clone it and use 'diff' for easy review of changes between different versions of the patchset. For example: $ git clone https://github.com/sjp38/damon-patches && cd damon-patches $ diff -u damon/v33 damon/v34 Sequence Of Patches =================== First three patches implement the core logics of DAMON. The 1st patch introduces basic sampling based hotness monitoring for arbitrary types of targets. Following two patches implement the core mechanisms for control of overhead and accuracy, namely regions based sampling (patch 2) and adaptive regions adjustment (patch 3). Now the essential parts of DAMON is complete, but it cannot work unless someone provides monitoring primitives for a specific use case. The following two patches make it just work for virtual address spaces monitoring. The 4th patch makes 'PG_idle' can be used by DAMON and the 5th patch implements the virtual memory address space specific monitoring primitives using page table Accessed bits and the 'PG_idle' page flag. Now DAMON just works for virtual address space monitoring via the kernel space api. To let the user space users can use DAMON, following four patches add interfaces for them. The 6th patch adds a tracepoint for monitoring results. The 7th patch implements a DAMON application kernel module, namely damon-dbgfs, that simply wraps DAMON and exposes DAMON interface to the user space via the debugfs interface. The 8th patch further exports pid of monitoring thread (kdamond) to user space for easier cpu usage accounting, and the 9th patch makes the debugfs interface to support multiple contexts. Three patches for maintainability follows. The 10th patch adds documentations for both the user space and the kernel space. The 11th patch provides unit tests (based on the kunit) while the 12th patch adds user space tests (based on the kselftest). Finally, the last patch (13th) updates the MAINTAINERS file. This patch (of 13): DAMON is a data access monitoring framework for the Linux kernel. The core mechanisms of DAMON make it - accurate (the monitoring output is useful enough for DRAM level performance-centric memory management; It might be inappropriate for CPU cache levels, though), - light-weight (the monitoring overhead is normally low enough to be applied online), and - scalable (the upper-bound of the overhead is in constant range regardless of the size of target workloads). Using this framework, hence, we can easily write efficient kernel space data access monitoring applications. For example, the kernel's memory management mechanisms can make advanced decisions using this. Experimental data access aware optimization works that incurring high access monitoring overhead could again be implemented on top of this. Due to its simple and flexible interface, providing user space interface would be also easy. Then, user space users who have some special workloads can write personalized applications for better understanding and optimizations of their workloads and systems. === Nevertheless, this commit is defining and implementing only basic access check part without the overhead-accuracy handling core logic. The basic access check is as below. The output of DAMON says what memory regions are how frequently accessed for a given duration. The resolution of the access frequency is controlled by setting ``sampling interval`` and ``aggregation interval``. In detail, DAMON checks access to each page per ``sampling interval`` and aggregates the results. In other words, counts the number of the accesses to each region. After each ``aggregation interval`` passes, DAMON calls callback functions that previously registered by users so that users can read the aggregated results and then clears the results. This can be described in below simple pseudo-code:: init() while monitoring_on: for page in monitoring_target: if accessed(page): nr_accesses[page] += 1 if time() % aggregation_interval == 0: for callback in user_registered_callbacks: callback(monitoring_target, nr_accesses) for page in monitoring_target: nr_accesses[page] = 0 if time() % update_interval == 0: update() sleep(sampling interval) The target regions constructed at the beginning of the monitoring and updated after each ``regions_update_interval``, because the target regions could be dynamically changed (e.g., mmap() or memory hotplug). The monitoring overhead of this mechanism will arbitrarily increase as the size of the target workload grows. The basic monitoring primitives for actual access check and dynamic target regions construction aren't in the core part of DAMON. Instead, it allows users to implement their own primitives that are optimized for their use case and configure DAMON to use those. In other words, users cannot use current version of DAMON without some additional works. Following commits will implement the core mechanisms for the overhead-accuracy control and default primitives implementations. Link: https://lkml.kernel.org/r/20210716081449.22187-1-sj38.park@gmail.com Link: https://lkml.kernel.org/r/20210716081449.22187-2-sj38.park@gmail.com Signed-off-by: SeongJae Park <sjpark@amazon.de> Reviewed-by: Leonard Foerster <foersleo@amazon.de> Reviewed-by: Fernand Sieber <sieberf@amazon.com> Acked-by: Shakeel Butt <shakeelb@google.com> Cc: Jonathan Cameron <Jonathan.Cameron@huawei.com> Cc: Alexander Shishkin <alexander.shishkin@linux.intel.com> Cc: Amit Shah <amit@kernel.org> Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org> Cc: Jonathan Corbet <corbet@lwn.net> Cc: David Hildenbrand <david@redhat.com> Cc: David Woodhouse <dwmw@amazon.com> Cc: Marco Elver <elver@google.com> Cc: Fan Du <fan.du@intel.com> Cc: Greg Kroah-Hartman <greg@kroah.com> Cc: Greg Thelen <gthelen@google.com> Cc: Joe Perches <joe@perches.com> Cc: Mel Gorman <mgorman@suse.de> Cc: Maximilian Heyne <mheyne@amazon.de> Cc: Minchan Kim <minchan@kernel.org> Cc: Ingo Molnar <mingo@redhat.com> Cc: Namhyung Kim <namhyung@kernel.org> Cc: Peter Zijlstra <peterz@infradead.org> Cc: Rik van Riel <riel@surriel.com> Cc: David Rientjes <rientjes@google.com> Cc: Steven Rostedt (VMware) <rostedt@goodmis.org> Cc: Shuah Khan <shuah@kernel.org> Cc: Vlastimil Babka <vbabka@suse.cz> Cc: Vladimir Davydov <vdavydov.dev@gmail.com> Cc: Brendan Higgins <brendanhiggins@google.com> Cc: Markus Boehme <markubo@amazon.de> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2021-09-07 19:56:28 -07:00
source "mm/damon/Kconfig"
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