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linux/fs/gfs2/incore.h

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/* SPDX-License-Identifier: GPL-2.0-only */
/*
* Copyright (C) Sistina Software, Inc. 1997-2003 All rights reserved.
* Copyright (C) 2004-2008 Red Hat, Inc. All rights reserved.
*/
#ifndef __INCORE_DOT_H__
#define __INCORE_DOT_H__
#include <linux/fs.h>
#include <linux/kobject.h>
#include <linux/workqueue.h>
#include <linux/dlm.h>
#include <linux/buffer_head.h>
#include <linux/rcupdate.h>
#include <linux/rculist_bl.h>
#include <linux/completion.h>
GFS2: Use rbtree for resource groups and clean up bitmap buffer ref count scheme Here is an update of Bob's original rbtree patch which, in addition, also resolves the rather strange ref counting that was being done relating to the bitmap blocks. Originally we had a dual system for journaling resource groups. The metadata blocks were journaled and also the rgrp itself was added to a list. The reason for adding the rgrp to the list in the journal was so that the "repolish clones" code could be run to update the free space, and potentially send any discard requests when the log was flushed. This was done by comparing the "cloned" bitmap with what had been written back on disk during the transaction commit. Due to this, there was a requirement to hang on to the rgrps' bitmap buffers until the journal had been flushed. For that reason, there was a rather complicated set up in the ->go_lock ->go_unlock functions for rgrps involving both a mutex and a spinlock (the ->sd_rindex_spin) to maintain a reference count on the buffers. However, the journal maintains a reference count on the buffers anyway, since they are being journaled as metadata buffers. So by moving the code which deals with the post-journal accounting for bitmap blocks to the metadata journaling code, we can entirely dispense with the rather strange buffer ref counting scheme and also the requirement to journal the rgrps. The net result of all this is that the ->sd_rindex_spin is left to do exactly one job, and that is to look after the rbtree or rgrps. This patch is designed to be a stepping stone towards using RCU for the rbtree of resource groups, however the reduction in the number of uses of the ->sd_rindex_spin is likely to have benefits for multi-threaded workloads, anyway. The patch retains ->go_lock and ->go_unlock for rgrps, however these maybe also be removed in future in favour of calling the functions directly where required in the code. That will allow locking of resource groups without needing to actually read them in - something that could be useful in speeding up statfs. In the mean time though it is valid to dereference ->bi_bh only when the rgrp is locked. This is basically the same rule as before, modulo the references not being valid until the following journal flush. Signed-off-by: Steven Whitehouse <swhiteho@redhat.com> Signed-off-by: Bob Peterson <rpeterso@redhat.com> Cc: Benjamin Marzinski <bmarzins@redhat.com>
2011-08-31 01:53:19 -07:00
#include <linux/rbtree.h>
GFS2: glock statistics gathering The stats are divided into two sets: those relating to the super block and those relating to an individual glock. The super block stats are done on a per cpu basis in order to try and reduce the overhead of gathering them. They are also further divided by glock type. In the case of both the super block and glock statistics, the same information is gathered in each case. The super block statistics are used to provide default values for most of the glock statistics, so that newly created glocks should have, as far as possible, a sensible starting point. The statistics are divided into three pairs of mean and variance, plus two counters. The mean/variance pairs are smoothed exponential estimates and the algorithm used is one which will be very familiar to those used to calculation of round trip times in network code. The three pairs of mean/variance measure the following things: 1. DLM lock time (non-blocking requests) 2. DLM lock time (blocking requests) 3. Inter-request time (again to the DLM) A non-blocking request is one which will complete right away, whatever the state of the DLM lock in question. That currently means any requests when (a) the current state of the lock is exclusive (b) the requested state is either null or unlocked or (c) the "try lock" flag is set. A blocking request covers all the other lock requests. There are two counters. The first is there primarily to show how many lock requests have been made, and thus how much data has gone into the mean/variance calculations. The other counter is counting queueing of holders at the top layer of the glock code. Hopefully that number will be a lot larger than the number of dlm lock requests issued. So why gather these statistics? There are several reasons we'd like to get a better idea of these timings: 1. To be able to better set the glock "min hold time" 2. To spot performance issues more easily 3. To improve the algorithm for selecting resource groups for allocation (to base it on lock wait time, rather than blindly using a "try lock") Due to the smoothing action of the updates, a step change in some input quantity being sampled will only fully be taken into account after 8 samples (or 4 for the variance) and this needs to be carefully considered when interpreting the results. Knowing both the time it takes a lock request to complete and the average time between lock requests for a glock means we can compute the total percentage of the time for which the node is able to use a glock vs. time that the rest of the cluster has its share. That will be very useful when setting the lock min hold time. The other point to remember is that all times are in nanoseconds. Great care has been taken to ensure that we measure exactly the quantities that we want, as accurately as possible. There are always inaccuracies in any measuring system, but I hope this is as accurate as we can reasonably make it. Signed-off-by: Steven Whitehouse <swhiteho@redhat.com>
2012-01-20 03:38:36 -07:00
#include <linux/ktime.h>
#include <linux/percpu.h>
#include <linux/lockref.h>
#include <linux/rhashtable.h>
#include <linux/mutex.h>
#define DIO_WAIT 0x00000010
#define DIO_METADATA 0x00000020
struct gfs2_log_operations;
struct gfs2_bufdata;
struct gfs2_holder;
struct gfs2_glock;
struct gfs2_quota_data;
struct gfs2_trans;
struct gfs2_jdesc;
struct gfs2_sbd;
struct lm_lockops;
typedef void (*gfs2_glop_bh_t) (struct gfs2_glock *gl, unsigned int ret);
struct gfs2_log_header_host {
u64 lh_sequence; /* Sequence number of this transaction */
u32 lh_flags; /* GFS2_LOG_HEAD_... */
u32 lh_tail; /* Block number of log tail */
u32 lh_blkno;
s64 lh_local_total;
s64 lh_local_free;
s64 lh_local_dinodes;
};
/*
* Structure of operations that are associated with each
* type of element in the log.
*/
struct gfs2_log_operations {
void (*lo_before_commit) (struct gfs2_sbd *sdp, struct gfs2_trans *tr);
void (*lo_after_commit) (struct gfs2_sbd *sdp, struct gfs2_trans *tr);
void (*lo_before_scan) (struct gfs2_jdesc *jd,
struct gfs2_log_header_host *head, int pass);
int (*lo_scan_elements) (struct gfs2_jdesc *jd, unsigned int start,
struct gfs2_log_descriptor *ld, __be64 *ptr,
int pass);
void (*lo_after_scan) (struct gfs2_jdesc *jd, int error, int pass);
const char *lo_name;
};
#define GBF_FULL 1
/**
* Clone bitmaps (bi_clone):
*
* - When a block is freed, we remember the previous state of the block in the
* clone bitmap, and only mark the block as free in the real bitmap.
*
* - When looking for a block to allocate, we check for a free block in the
* clone bitmap, and if no clone bitmap exists, in the real bitmap.
*
* - For allocating a block, we mark it as allocated in the real bitmap, and if
* a clone bitmap exists, also in the clone bitmap.
*
* - At the end of a log_flush, we copy the real bitmap into the clone bitmap
* to make the clone bitmap reflect the current allocation state.
* (Alternatively, we could remove the clone bitmap.)
*
* The clone bitmaps are in-core only, and is never written to disk.
*
* These steps ensure that blocks which have been freed in a transaction cannot
* be reallocated in that same transaction.
*/
struct gfs2_bitmap {
struct buffer_head *bi_bh;
char *bi_clone;
unsigned long bi_flags;
u32 bi_offset;
u32 bi_start;
u32 bi_bytes;
u32 bi_blocks;
};
struct gfs2_rgrpd {
GFS2: Use rbtree for resource groups and clean up bitmap buffer ref count scheme Here is an update of Bob's original rbtree patch which, in addition, also resolves the rather strange ref counting that was being done relating to the bitmap blocks. Originally we had a dual system for journaling resource groups. The metadata blocks were journaled and also the rgrp itself was added to a list. The reason for adding the rgrp to the list in the journal was so that the "repolish clones" code could be run to update the free space, and potentially send any discard requests when the log was flushed. This was done by comparing the "cloned" bitmap with what had been written back on disk during the transaction commit. Due to this, there was a requirement to hang on to the rgrps' bitmap buffers until the journal had been flushed. For that reason, there was a rather complicated set up in the ->go_lock ->go_unlock functions for rgrps involving both a mutex and a spinlock (the ->sd_rindex_spin) to maintain a reference count on the buffers. However, the journal maintains a reference count on the buffers anyway, since they are being journaled as metadata buffers. So by moving the code which deals with the post-journal accounting for bitmap blocks to the metadata journaling code, we can entirely dispense with the rather strange buffer ref counting scheme and also the requirement to journal the rgrps. The net result of all this is that the ->sd_rindex_spin is left to do exactly one job, and that is to look after the rbtree or rgrps. This patch is designed to be a stepping stone towards using RCU for the rbtree of resource groups, however the reduction in the number of uses of the ->sd_rindex_spin is likely to have benefits for multi-threaded workloads, anyway. The patch retains ->go_lock and ->go_unlock for rgrps, however these maybe also be removed in future in favour of calling the functions directly where required in the code. That will allow locking of resource groups without needing to actually read them in - something that could be useful in speeding up statfs. In the mean time though it is valid to dereference ->bi_bh only when the rgrp is locked. This is basically the same rule as before, modulo the references not being valid until the following journal flush. Signed-off-by: Steven Whitehouse <swhiteho@redhat.com> Signed-off-by: Bob Peterson <rpeterso@redhat.com> Cc: Benjamin Marzinski <bmarzins@redhat.com>
2011-08-31 01:53:19 -07:00
struct rb_node rd_node; /* Link with superblock */
struct gfs2_glock *rd_gl; /* Glock for this rgrp */
u64 rd_addr; /* grp block disk address */
u64 rd_data0; /* first data location */
u32 rd_length; /* length of rgrp header in fs blocks */
u32 rd_data; /* num of data blocks in rgrp */
u32 rd_bitbytes; /* number of bytes in data bitmaps */
u32 rd_free;
u32 rd_requested; /* number of blocks in rd_rstree */
u32 rd_reserved; /* number of reserved blocks */
u32 rd_free_clone;
u32 rd_dinodes;
u64 rd_igeneration;
struct gfs2_bitmap *rd_bits;
struct gfs2_sbd *rd_sbd;
struct gfs2_rgrp_lvb *rd_rgl;
u32 rd_last_alloc;
u32 rd_flags;
u32 rd_extfail_pt; /* extent failure point */
#define GFS2_RDF_CHECK 0x10000000 /* check for unlinked inodes */
#define GFS2_RDF_ERROR 0x40000000 /* error in rg */
#define GFS2_RDF_PREFERRED 0x80000000 /* This rgrp is preferred */
#define GFS2_RDF_MASK 0xf0000000 /* mask for internal flags */
spinlock_t rd_rsspin; /* protects reservation related vars */
struct mutex rd_mutex;
struct rb_root rd_rstree; /* multi-block reservation tree */
};
enum gfs2_state_bits {
BH_Pinned = BH_PrivateStart,
BH_Escaped = BH_PrivateStart + 1,
};
BUFFER_FNS(Pinned, pinned)
TAS_BUFFER_FNS(Pinned, pinned)
BUFFER_FNS(Escaped, escaped)
TAS_BUFFER_FNS(Escaped, escaped)
struct gfs2_bufdata {
struct buffer_head *bd_bh;
struct gfs2_glock *bd_gl;
u64 bd_blkno;
struct list_head bd_list;
struct gfs2_trans *bd_tr;
struct list_head bd_ail_st_list;
struct list_head bd_ail_gl_list;
};
/*
* Internally, we prefix things with gdlm_ and GDLM_ (for gfs-dlm) since a
* prefix of lock_dlm_ gets awkward.
*/
#define GDLM_STRNAME_BYTES 25
#define GDLM_LVB_SIZE 32
/*
* ls_recover_flags:
*
* DFL_BLOCK_LOCKS: dlm is in recovery and will grant locks that had been
* held by failed nodes whose journals need recovery. Those locks should
* only be used for journal recovery until the journal recovery is done.
* This is set by the dlm recover_prep callback and cleared by the
* gfs2_control thread when journal recovery is complete. To avoid
* races between recover_prep setting and gfs2_control clearing, recover_spin
* is held while changing this bit and reading/writing recover_block
* and recover_start.
*
* DFL_NO_DLM_OPS: dlm lockspace ops/callbacks are not being used.
*
* DFL_FIRST_MOUNT: this node is the first to mount this fs and is doing
* recovery of all journals before allowing other nodes to mount the fs.
* This is cleared when FIRST_MOUNT_DONE is set.
*
* DFL_FIRST_MOUNT_DONE: this node was the first mounter, and has finished
* recovery of all journals, and now allows other nodes to mount the fs.
*
* DFL_MOUNT_DONE: gdlm_mount has completed successfully and cleared
* BLOCK_LOCKS for the first time. The gfs2_control thread should now
* control clearing BLOCK_LOCKS for further recoveries.
*
* DFL_UNMOUNT: gdlm_unmount sets to keep sdp off gfs2_control_wq.
*
* DFL_DLM_RECOVERY: set while dlm is in recovery, between recover_prep()
* and recover_done(), i.e. set while recover_block == recover_start.
*/
enum {
DFL_BLOCK_LOCKS = 0,
DFL_NO_DLM_OPS = 1,
DFL_FIRST_MOUNT = 2,
DFL_FIRST_MOUNT_DONE = 3,
DFL_MOUNT_DONE = 4,
DFL_UNMOUNT = 5,
DFL_DLM_RECOVERY = 6,
};
/*
* We are using struct lm_lockname as an rhashtable key. Avoid holes within
* the struct; padding at the end is fine.
*/
struct lm_lockname {
u64 ln_number;
struct gfs2_sbd *ln_sbd;
unsigned int ln_type;
};
#define lm_name_equal(name1, name2) \
(((name1)->ln_number == (name2)->ln_number) && \
((name1)->ln_type == (name2)->ln_type) && \
((name1)->ln_sbd == (name2)->ln_sbd))
struct gfs2_glock_operations {
int (*go_sync) (struct gfs2_glock *gl);
int (*go_xmote_bh)(struct gfs2_glock *gl);
void (*go_inval) (struct gfs2_glock *gl, int flags);
int (*go_instantiate) (struct gfs2_glock *gl);
int (*go_held)(struct gfs2_holder *gh);
void (*go_dump)(struct seq_file *seq, const struct gfs2_glock *gl,
const char *fs_id_buf);
void (*go_callback)(struct gfs2_glock *gl, bool remote);
void (*go_unlocked)(struct gfs2_glock *gl);
gfs2: set lockdep subclass for iopen glocks This patch introduce a new globs attribute to define the subclass of the glock lockref spinlock. This avoid the following lockdep warning, which occurs when we lock an inode lock while an iopen lock is held: ============================================ WARNING: possible recursive locking detected 5.10.0-rc3+ #4990 Not tainted -------------------------------------------- kworker/0:1/12 is trying to acquire lock: ffff9067d45672d8 (&gl->gl_lockref.lock){+.+.}-{3:3}, at: lockref_get+0x9/0x20 but task is already holding lock: ffff9067da308588 (&gl->gl_lockref.lock){+.+.}-{3:3}, at: delete_work_func+0x164/0x260 other info that might help us debug this: Possible unsafe locking scenario: CPU0 ---- lock(&gl->gl_lockref.lock); lock(&gl->gl_lockref.lock); *** DEADLOCK *** May be due to missing lock nesting notation 3 locks held by kworker/0:1/12: #0: ffff9067c1bfdd38 ((wq_completion)delete_workqueue){+.+.}-{0:0}, at: process_one_work+0x1b7/0x540 #1: ffffac594006be70 ((work_completion)(&(&gl->gl_delete)->work)){+.+.}-{0:0}, at: process_one_work+0x1b7/0x540 #2: ffff9067da308588 (&gl->gl_lockref.lock){+.+.}-{3:3}, at: delete_work_func+0x164/0x260 stack backtrace: CPU: 0 PID: 12 Comm: kworker/0:1 Not tainted 5.10.0-rc3+ #4990 Hardware name: QEMU Standard PC (Q35 + ICH9, 2009), BIOS 1.13.0-2.fc32 04/01/2014 Workqueue: delete_workqueue delete_work_func Call Trace: dump_stack+0x8b/0xb0 __lock_acquire.cold+0x19e/0x2e3 lock_acquire+0x150/0x410 ? lockref_get+0x9/0x20 _raw_spin_lock+0x27/0x40 ? lockref_get+0x9/0x20 lockref_get+0x9/0x20 delete_work_func+0x188/0x260 process_one_work+0x237/0x540 worker_thread+0x4d/0x3b0 ? process_one_work+0x540/0x540 kthread+0x127/0x140 ? __kthread_bind_mask+0x60/0x60 ret_from_fork+0x22/0x30 Suggested-by: Andreas Gruenbacher <agruenba@redhat.com> Signed-off-by: Alexander Aring <aahringo@redhat.com> Signed-off-by: Andreas Gruenbacher <agruenba@redhat.com>
2020-11-23 08:53:35 -07:00
const int go_subclass;
const int go_type;
const unsigned long go_flags;
#define GLOF_ASPACE 1 /* address space attached */
#define GLOF_LVB 2 /* Lock Value Block attached */
#define GLOF_NONDISK 8 /* not I/O related */
};
GFS2: glock statistics gathering The stats are divided into two sets: those relating to the super block and those relating to an individual glock. The super block stats are done on a per cpu basis in order to try and reduce the overhead of gathering them. They are also further divided by glock type. In the case of both the super block and glock statistics, the same information is gathered in each case. The super block statistics are used to provide default values for most of the glock statistics, so that newly created glocks should have, as far as possible, a sensible starting point. The statistics are divided into three pairs of mean and variance, plus two counters. The mean/variance pairs are smoothed exponential estimates and the algorithm used is one which will be very familiar to those used to calculation of round trip times in network code. The three pairs of mean/variance measure the following things: 1. DLM lock time (non-blocking requests) 2. DLM lock time (blocking requests) 3. Inter-request time (again to the DLM) A non-blocking request is one which will complete right away, whatever the state of the DLM lock in question. That currently means any requests when (a) the current state of the lock is exclusive (b) the requested state is either null or unlocked or (c) the "try lock" flag is set. A blocking request covers all the other lock requests. There are two counters. The first is there primarily to show how many lock requests have been made, and thus how much data has gone into the mean/variance calculations. The other counter is counting queueing of holders at the top layer of the glock code. Hopefully that number will be a lot larger than the number of dlm lock requests issued. So why gather these statistics? There are several reasons we'd like to get a better idea of these timings: 1. To be able to better set the glock "min hold time" 2. To spot performance issues more easily 3. To improve the algorithm for selecting resource groups for allocation (to base it on lock wait time, rather than blindly using a "try lock") Due to the smoothing action of the updates, a step change in some input quantity being sampled will only fully be taken into account after 8 samples (or 4 for the variance) and this needs to be carefully considered when interpreting the results. Knowing both the time it takes a lock request to complete and the average time between lock requests for a glock means we can compute the total percentage of the time for which the node is able to use a glock vs. time that the rest of the cluster has its share. That will be very useful when setting the lock min hold time. The other point to remember is that all times are in nanoseconds. Great care has been taken to ensure that we measure exactly the quantities that we want, as accurately as possible. There are always inaccuracies in any measuring system, but I hope this is as accurate as we can reasonably make it. Signed-off-by: Steven Whitehouse <swhiteho@redhat.com>
2012-01-20 03:38:36 -07:00
enum {
GFS2_LKS_SRTT = 0, /* Non blocking smoothed round trip time */
GFS2_LKS_SRTTVAR = 1, /* Non blocking smoothed variance */
GFS2_LKS_SRTTB = 2, /* Blocking smoothed round trip time */
GFS2_LKS_SRTTVARB = 3, /* Blocking smoothed variance */
GFS2_LKS_SIRT = 4, /* Smoothed Inter-request time */
GFS2_LKS_SIRTVAR = 5, /* Smoothed Inter-request variance */
GFS2_LKS_DCOUNT = 6, /* Count of dlm requests */
GFS2_LKS_QCOUNT = 7, /* Count of gfs2_holder queues */
GFS2_NR_LKSTATS
};
struct gfs2_lkstats {
u64 stats[GFS2_NR_LKSTATS];
GFS2: glock statistics gathering The stats are divided into two sets: those relating to the super block and those relating to an individual glock. The super block stats are done on a per cpu basis in order to try and reduce the overhead of gathering them. They are also further divided by glock type. In the case of both the super block and glock statistics, the same information is gathered in each case. The super block statistics are used to provide default values for most of the glock statistics, so that newly created glocks should have, as far as possible, a sensible starting point. The statistics are divided into three pairs of mean and variance, plus two counters. The mean/variance pairs are smoothed exponential estimates and the algorithm used is one which will be very familiar to those used to calculation of round trip times in network code. The three pairs of mean/variance measure the following things: 1. DLM lock time (non-blocking requests) 2. DLM lock time (blocking requests) 3. Inter-request time (again to the DLM) A non-blocking request is one which will complete right away, whatever the state of the DLM lock in question. That currently means any requests when (a) the current state of the lock is exclusive (b) the requested state is either null or unlocked or (c) the "try lock" flag is set. A blocking request covers all the other lock requests. There are two counters. The first is there primarily to show how many lock requests have been made, and thus how much data has gone into the mean/variance calculations. The other counter is counting queueing of holders at the top layer of the glock code. Hopefully that number will be a lot larger than the number of dlm lock requests issued. So why gather these statistics? There are several reasons we'd like to get a better idea of these timings: 1. To be able to better set the glock "min hold time" 2. To spot performance issues more easily 3. To improve the algorithm for selecting resource groups for allocation (to base it on lock wait time, rather than blindly using a "try lock") Due to the smoothing action of the updates, a step change in some input quantity being sampled will only fully be taken into account after 8 samples (or 4 for the variance) and this needs to be carefully considered when interpreting the results. Knowing both the time it takes a lock request to complete and the average time between lock requests for a glock means we can compute the total percentage of the time for which the node is able to use a glock vs. time that the rest of the cluster has its share. That will be very useful when setting the lock min hold time. The other point to remember is that all times are in nanoseconds. Great care has been taken to ensure that we measure exactly the quantities that we want, as accurately as possible. There are always inaccuracies in any measuring system, but I hope this is as accurate as we can reasonably make it. Signed-off-by: Steven Whitehouse <swhiteho@redhat.com>
2012-01-20 03:38:36 -07:00
};
enum {
/* States */
HIF_HOLDER = 6, /* Set for gh that "holds" the glock */
HIF_WAIT = 10,
};
struct gfs2_holder {
struct list_head gh_list;
struct gfs2_glock *gh_gl;
struct pid *gh_owner_pid;
u16 gh_flags;
u16 gh_state;
int gh_error;
unsigned long gh_iflags; /* HIF_... */
unsigned long gh_ip;
};
/* Number of quota types we support */
#define GFS2_MAXQUOTAS 2
struct gfs2_qadata { /* quota allocation data */
/* Quota stuff */
struct gfs2_quota_data *qa_qd[2 * GFS2_MAXQUOTAS];
struct gfs2_holder qa_qd_ghs[2 * GFS2_MAXQUOTAS];
unsigned int qa_qd_num;
int qa_ref;
};
/* Resource group multi-block reservation, in order of appearance:
Step 1. Function prepares to write, allocates a mb, sets the size hint.
Step 2. User calls inplace_reserve to target an rgrp, sets the rgrp info
Step 3. Function get_local_rgrp locks the rgrp, determines which bits to use
Step 4. Bits are assigned from the rgrp based on either the reservation
or wherever it can.
*/
struct gfs2_blkreserv {
struct rb_node rs_node; /* node within rd_rstree */
struct gfs2_rgrpd *rs_rgd;
u64 rs_start;
u32 rs_requested;
u32 rs_reserved; /* number of reserved blocks */
};
/*
* Allocation parameters
* @target: The number of blocks we'd ideally like to allocate
* @aflags: The flags (e.g. Orlov flag)
*
* The intent is to gradually expand this structure over time in
* order to give more information, e.g. alignment, min extent size
* to the allocation code.
*/
struct gfs2_alloc_parms {
u64 target;
u32 min_target;
u32 aflags;
u64 allowed;
};
enum {
GLF_LOCK = 1,
gfs2: fix GL_SKIP node_scope problems Before this patch, when a glock was locked, the very first holder on the queue would unlock the lockref and call the go_instantiate glops function (if one existed), unless GL_SKIP was specified. When we introduced the new node-scope concept, we allowed multiple holders to lock glocks in EX mode and share the lock. But node-scope introduced a new problem: if the first holder has GL_SKIP and the next one does NOT, since it is not the first holder on the queue, the go_instantiate op was not called. Eventually the GL_SKIP holder may call the instantiate sub-function (e.g. gfs2_rgrp_bh_get) but there was still a window of time in which another non-GL_SKIP holder assumes the instantiate function had been called by the first holder. In the case of rgrp glocks, this led to a NULL pointer dereference on the buffer_heads. This patch tries to fix the problem by introducing two new glock flags: GLF_INSTANTIATE_NEEDED, which keeps track of when the instantiate function needs to be called to "fill in" or "read in" the object before it is referenced. GLF_INSTANTIATE_IN_PROG which is used to determine when a process is in the process of reading in the object. Whenever a function needs to reference the object, it checks the GLF_INSTANTIATE_NEEDED flag, and if set, it sets GLF_INSTANTIATE_IN_PROG and calls the glops "go_instantiate" function. As before, the gl_lockref spin_lock is unlocked during the IO operation, which may take a relatively long amount of time to complete. While unlocked, if another process determines go_instantiate is still needed, it sees GLF_INSTANTIATE_IN_PROG is set, and waits for the go_instantiate glop operation to be completed. Once GLF_INSTANTIATE_IN_PROG is cleared, it needs to check GLF_INSTANTIATE_NEEDED again because the other process's go_instantiate operation may not have been successful. Functions that previously called the instantiate sub-functions now call directly into gfs2_instantiate so the new bits are managed properly. Signed-off-by: Bob Peterson <rpeterso@redhat.com> Signed-off-by: Andreas Gruenbacher <agruenba@redhat.com>
2021-10-06 07:29:18 -07:00
GLF_INSTANTIATE_NEEDED = 2, /* needs instantiate */
GLF_DEMOTE = 3,
GLF_PENDING_DEMOTE = 4,
GLF_DEMOTE_IN_PROGRESS = 5,
GLF_DIRTY = 6,
GLF_LFLUSH = 7,
GLF_INVALIDATE_IN_PROGRESS = 8,
GLF_HAVE_REPLY = 9,
GLF_INITIAL = 10,
GLF_HAVE_FROZEN_REPLY = 11,
gfs2: fix GL_SKIP node_scope problems Before this patch, when a glock was locked, the very first holder on the queue would unlock the lockref and call the go_instantiate glops function (if one existed), unless GL_SKIP was specified. When we introduced the new node-scope concept, we allowed multiple holders to lock glocks in EX mode and share the lock. But node-scope introduced a new problem: if the first holder has GL_SKIP and the next one does NOT, since it is not the first holder on the queue, the go_instantiate op was not called. Eventually the GL_SKIP holder may call the instantiate sub-function (e.g. gfs2_rgrp_bh_get) but there was still a window of time in which another non-GL_SKIP holder assumes the instantiate function had been called by the first holder. In the case of rgrp glocks, this led to a NULL pointer dereference on the buffer_heads. This patch tries to fix the problem by introducing two new glock flags: GLF_INSTANTIATE_NEEDED, which keeps track of when the instantiate function needs to be called to "fill in" or "read in" the object before it is referenced. GLF_INSTANTIATE_IN_PROG which is used to determine when a process is in the process of reading in the object. Whenever a function needs to reference the object, it checks the GLF_INSTANTIATE_NEEDED flag, and if set, it sets GLF_INSTANTIATE_IN_PROG and calls the glops "go_instantiate" function. As before, the gl_lockref spin_lock is unlocked during the IO operation, which may take a relatively long amount of time to complete. While unlocked, if another process determines go_instantiate is still needed, it sees GLF_INSTANTIATE_IN_PROG is set, and waits for the go_instantiate glop operation to be completed. Once GLF_INSTANTIATE_IN_PROG is cleared, it needs to check GLF_INSTANTIATE_NEEDED again because the other process's go_instantiate operation may not have been successful. Functions that previously called the instantiate sub-functions now call directly into gfs2_instantiate so the new bits are managed properly. Signed-off-by: Bob Peterson <rpeterso@redhat.com> Signed-off-by: Andreas Gruenbacher <agruenba@redhat.com>
2021-10-06 07:29:18 -07:00
GLF_INSTANTIATE_IN_PROG = 12, /* instantiate happening now */
GLF_LRU = 13,
GLF_OBJECT = 14, /* Used only for tracing */
GFS2: glock statistics gathering The stats are divided into two sets: those relating to the super block and those relating to an individual glock. The super block stats are done on a per cpu basis in order to try and reduce the overhead of gathering them. They are also further divided by glock type. In the case of both the super block and glock statistics, the same information is gathered in each case. The super block statistics are used to provide default values for most of the glock statistics, so that newly created glocks should have, as far as possible, a sensible starting point. The statistics are divided into three pairs of mean and variance, plus two counters. The mean/variance pairs are smoothed exponential estimates and the algorithm used is one which will be very familiar to those used to calculation of round trip times in network code. The three pairs of mean/variance measure the following things: 1. DLM lock time (non-blocking requests) 2. DLM lock time (blocking requests) 3. Inter-request time (again to the DLM) A non-blocking request is one which will complete right away, whatever the state of the DLM lock in question. That currently means any requests when (a) the current state of the lock is exclusive (b) the requested state is either null or unlocked or (c) the "try lock" flag is set. A blocking request covers all the other lock requests. There are two counters. The first is there primarily to show how many lock requests have been made, and thus how much data has gone into the mean/variance calculations. The other counter is counting queueing of holders at the top layer of the glock code. Hopefully that number will be a lot larger than the number of dlm lock requests issued. So why gather these statistics? There are several reasons we'd like to get a better idea of these timings: 1. To be able to better set the glock "min hold time" 2. To spot performance issues more easily 3. To improve the algorithm for selecting resource groups for allocation (to base it on lock wait time, rather than blindly using a "try lock") Due to the smoothing action of the updates, a step change in some input quantity being sampled will only fully be taken into account after 8 samples (or 4 for the variance) and this needs to be carefully considered when interpreting the results. Knowing both the time it takes a lock request to complete and the average time between lock requests for a glock means we can compute the total percentage of the time for which the node is able to use a glock vs. time that the rest of the cluster has its share. That will be very useful when setting the lock min hold time. The other point to remember is that all times are in nanoseconds. Great care has been taken to ensure that we measure exactly the quantities that we want, as accurately as possible. There are always inaccuracies in any measuring system, but I hope this is as accurate as we can reasonably make it. Signed-off-by: Steven Whitehouse <swhiteho@redhat.com>
2012-01-20 03:38:36 -07:00
GLF_BLOCKING = 15,
GLF_UNLOCKED = 16, /* Wait for glock to be unlocked */
GLF_TRY_TO_EVICT = 17, /* iopen glocks only */
GLF_VERIFY_EVICT = 18, /* iopen glocks only */
};
struct gfs2_glock {
unsigned long gl_flags; /* GLF_... */
struct lm_lockname gl_name;
struct lockref gl_lockref;
/* State fields protected by gl_lockref.lock */
unsigned int gl_state:2, /* Current state */
gl_target:2, /* Target state */
gl_demote_state:2, /* State requested by remote node */
gl_req:2, /* State in last dlm request */
gl_reply:8; /* Last reply from the dlm */
unsigned long gl_demote_time; /* time of first demote request */
long gl_hold_time;
struct list_head gl_holders;
const struct gfs2_glock_operations *gl_ops;
GFS2: glock statistics gathering The stats are divided into two sets: those relating to the super block and those relating to an individual glock. The super block stats are done on a per cpu basis in order to try and reduce the overhead of gathering them. They are also further divided by glock type. In the case of both the super block and glock statistics, the same information is gathered in each case. The super block statistics are used to provide default values for most of the glock statistics, so that newly created glocks should have, as far as possible, a sensible starting point. The statistics are divided into three pairs of mean and variance, plus two counters. The mean/variance pairs are smoothed exponential estimates and the algorithm used is one which will be very familiar to those used to calculation of round trip times in network code. The three pairs of mean/variance measure the following things: 1. DLM lock time (non-blocking requests) 2. DLM lock time (blocking requests) 3. Inter-request time (again to the DLM) A non-blocking request is one which will complete right away, whatever the state of the DLM lock in question. That currently means any requests when (a) the current state of the lock is exclusive (b) the requested state is either null or unlocked or (c) the "try lock" flag is set. A blocking request covers all the other lock requests. There are two counters. The first is there primarily to show how many lock requests have been made, and thus how much data has gone into the mean/variance calculations. The other counter is counting queueing of holders at the top layer of the glock code. Hopefully that number will be a lot larger than the number of dlm lock requests issued. So why gather these statistics? There are several reasons we'd like to get a better idea of these timings: 1. To be able to better set the glock "min hold time" 2. To spot performance issues more easily 3. To improve the algorithm for selecting resource groups for allocation (to base it on lock wait time, rather than blindly using a "try lock") Due to the smoothing action of the updates, a step change in some input quantity being sampled will only fully be taken into account after 8 samples (or 4 for the variance) and this needs to be carefully considered when interpreting the results. Knowing both the time it takes a lock request to complete and the average time between lock requests for a glock means we can compute the total percentage of the time for which the node is able to use a glock vs. time that the rest of the cluster has its share. That will be very useful when setting the lock min hold time. The other point to remember is that all times are in nanoseconds. Great care has been taken to ensure that we measure exactly the quantities that we want, as accurately as possible. There are always inaccuracies in any measuring system, but I hope this is as accurate as we can reasonably make it. Signed-off-by: Steven Whitehouse <swhiteho@redhat.com>
2012-01-20 03:38:36 -07:00
ktime_t gl_dstamp;
struct gfs2_lkstats gl_stats;
struct dlm_lksb gl_lksb;
unsigned long gl_tchange;
void *gl_object;
struct list_head gl_lru;
struct list_head gl_ail_list;
atomic_t gl_ail_count;
atomic_t gl_revokes;
struct delayed_work gl_work;
/* For iopen glocks only */
struct {
struct delayed_work gl_delete;
u64 gl_no_formal_ino;
};
struct rcu_head gl_rcu;
struct rhash_head gl_node;
};
enum {
GIF_QD_LOCKED = 1,
GIF_ALLOC_FAILED = 2,
GIF_SW_PAGED = 3,
GIF_FREE_VFS_INODE = 5,
GIF_GLOP_PENDING = 6,
GIF_DEFERRED_DELETE = 7,
};
struct gfs2_inode {
struct inode i_inode;
u64 i_no_addr;
u64 i_no_formal_ino;
u64 i_generation;
u64 i_eattr;
unsigned long i_flags; /* GIF_... */
struct gfs2_glock *i_gl;
struct gfs2_holder i_iopen_gh;
struct gfs2_qadata *i_qadata; /* quota allocation data */
struct gfs2_holder i_rgd_gh;
struct gfs2_blkreserv i_res; /* rgrp multi-block reservation */
u64 i_goal; /* goal block for allocations */
atomic_t i_sizehint; /* hint of the write size */
struct rw_semaphore i_rw_mutex;
struct list_head i_ordered;
__be64 *i_hash_cache;
u32 i_entries;
u32 i_diskflags;
u8 i_height;
u8 i_depth;
u16 i_rahead;
};
/*
* Since i_inode is the first element of struct gfs2_inode,
* this is effectively a cast.
*/
static inline struct gfs2_inode *GFS2_I(struct inode *inode)
{
return container_of(inode, struct gfs2_inode, i_inode);
}
static inline struct gfs2_sbd *GFS2_SB(const struct inode *inode)
{
return inode->i_sb->s_fs_info;
}
struct gfs2_file {
struct mutex f_fl_mutex;
struct gfs2_holder f_fl_gh;
};
struct gfs2_revoke_replay {
struct list_head rr_list;
u64 rr_blkno;
unsigned int rr_where;
};
enum {
QDF_CHANGE = 1,
QDF_LOCKED = 2,
QDF_REFRESH = 3,
QDF_QMSG_QUIET = 4,
};
struct gfs2_quota_data {
struct hlist_bl_node qd_hlist;
struct list_head qd_list;
struct kqid qd_id;
struct gfs2_sbd *qd_sbd;
struct lockref qd_lockref;
struct list_head qd_lru;
unsigned qd_hash;
unsigned long qd_flags; /* QDF_... */
s64 qd_change;
s64 qd_change_sync;
unsigned int qd_slot;
unsigned int qd_slot_ref;
struct buffer_head *qd_bh;
struct gfs2_quota_change *qd_bh_qc;
unsigned int qd_bh_count;
struct gfs2_glock *qd_gl;
struct gfs2_quota_lvb qd_qb;
u64 qd_sync_gen;
unsigned long qd_last_warn;
struct rcu_head qd_rcu;
};
enum {
TR_TOUCHED = 1,
TR_ATTACHED = 2,
TR_ONSTACK = 3,
};
struct gfs2_trans {
unsigned long tr_ip;
unsigned int tr_blocks;
unsigned int tr_revokes;
unsigned int tr_reserved;
unsigned long tr_flags;
unsigned int tr_num_buf_new;
[GFS2] assertion failure after writing to journaled file, umount This patch passes all my nasty tests that were causing the code to fail under one circumstance or another. Here is a complete summary of all changes from today's git tree, in order of appearance: 1. There are now separate variables for metadata buffer accounting. 2. Variable sd_log_num_hdrs is no longer needed, since the header accounting is taken care of by the reserve/refund sequence. 3. Fixed a tiny grammatical problem in a comment. 4. Added a new function "calc_reserved" to calculate the reserved log space. This isn't entirely necessary, but it has two benefits: First, it simplifies the gfs2_log_refund function greatly. Second, it allows for easier debugging because I could sprinkle the code with calls to this function to make sure the accounting is proper (by adding asserts and printks) at strategic point of the code. 5. In log_pull_tail there apparently was a kludge to fix up the accounting based on a "pull" parameter. The buffer accounting is now done properly, so the kludge was removed. 6. File sync operations were making a call to gfs2_log_flush that writes another journal header. Since that header was unplanned for (reserved) by the reserve/refund sequence, the free space had to be decremented so that when log_pull_tail gets called, the free space is be adjusted properly. (Did I hear you call that a kludge? well, maybe, but a lot more justifiable than the one I removed). 7. In the gfs2_log_shutdown code, it optionally syncs the log by specifying the PULL parameter to log_write_header. I'm not sure this is necessary anymore. It just seems to me there could be cases where shutdown is called while there are outstanding log buffers. 8. In the (data)buf_lo_before_commit functions, I changed some offset values from being calculated on the fly to being constants. That simplified some code and we might as well let the compiler do the calculation once rather than redoing those cycles at run time. 9. This version has my rewritten databuf_lo_add function. This version is much more like its predecessor, buf_lo_add, which makes it easier to understand. Again, this might not be necessary, but it seems as if this one works as well as the previous one, maybe even better, so I decided to leave it in. 10. In databuf_lo_before_commit, a previous data corruption problem was caused by going off the end of the buffer. The proper solution is to have the proper limit in place, rather than stopping earlier. (Thus my previous attempt to fix it is wrong). If you don't wrap the buffer, you're stopping too early and that causes more log buffer accounting problems. 11. In lops.h there are two new (previously mentioned) constants for figuring out the data offset for the journal buffers. 12. There are also two new functions, buf_limit and databuf_limit to calculate how many entries will fit in the buffer. 13. In function gfs2_meta_wipe, it needs to distinguish between pinned metadata buffers and journaled data buffers for proper journal buffer accounting. It can't use the JDATA gfs2_inode flag because it's sometimes passed the "real" inode and sometimes the "metadata inode" and the inode flags will be random bits in a metadata gfs2_inode. It needs to base its decision on which was passed in. Signed-off-by: Bob Peterson <rpeterso@redhat.com> Signed-off-by: Steven Whitehouse <swhiteho@redhat.com>
2007-06-18 12:50:20 -07:00
unsigned int tr_num_databuf_new;
unsigned int tr_num_buf_rm;
[GFS2] assertion failure after writing to journaled file, umount This patch passes all my nasty tests that were causing the code to fail under one circumstance or another. Here is a complete summary of all changes from today's git tree, in order of appearance: 1. There are now separate variables for metadata buffer accounting. 2. Variable sd_log_num_hdrs is no longer needed, since the header accounting is taken care of by the reserve/refund sequence. 3. Fixed a tiny grammatical problem in a comment. 4. Added a new function "calc_reserved" to calculate the reserved log space. This isn't entirely necessary, but it has two benefits: First, it simplifies the gfs2_log_refund function greatly. Second, it allows for easier debugging because I could sprinkle the code with calls to this function to make sure the accounting is proper (by adding asserts and printks) at strategic point of the code. 5. In log_pull_tail there apparently was a kludge to fix up the accounting based on a "pull" parameter. The buffer accounting is now done properly, so the kludge was removed. 6. File sync operations were making a call to gfs2_log_flush that writes another journal header. Since that header was unplanned for (reserved) by the reserve/refund sequence, the free space had to be decremented so that when log_pull_tail gets called, the free space is be adjusted properly. (Did I hear you call that a kludge? well, maybe, but a lot more justifiable than the one I removed). 7. In the gfs2_log_shutdown code, it optionally syncs the log by specifying the PULL parameter to log_write_header. I'm not sure this is necessary anymore. It just seems to me there could be cases where shutdown is called while there are outstanding log buffers. 8. In the (data)buf_lo_before_commit functions, I changed some offset values from being calculated on the fly to being constants. That simplified some code and we might as well let the compiler do the calculation once rather than redoing those cycles at run time. 9. This version has my rewritten databuf_lo_add function. This version is much more like its predecessor, buf_lo_add, which makes it easier to understand. Again, this might not be necessary, but it seems as if this one works as well as the previous one, maybe even better, so I decided to leave it in. 10. In databuf_lo_before_commit, a previous data corruption problem was caused by going off the end of the buffer. The proper solution is to have the proper limit in place, rather than stopping earlier. (Thus my previous attempt to fix it is wrong). If you don't wrap the buffer, you're stopping too early and that causes more log buffer accounting problems. 11. In lops.h there are two new (previously mentioned) constants for figuring out the data offset for the journal buffers. 12. There are also two new functions, buf_limit and databuf_limit to calculate how many entries will fit in the buffer. 13. In function gfs2_meta_wipe, it needs to distinguish between pinned metadata buffers and journaled data buffers for proper journal buffer accounting. It can't use the JDATA gfs2_inode flag because it's sometimes passed the "real" inode and sometimes the "metadata inode" and the inode flags will be random bits in a metadata gfs2_inode. It needs to base its decision on which was passed in. Signed-off-by: Bob Peterson <rpeterso@redhat.com> Signed-off-by: Steven Whitehouse <swhiteho@redhat.com>
2007-06-18 12:50:20 -07:00
unsigned int tr_num_databuf_rm;
unsigned int tr_num_revoke;
struct list_head tr_list;
struct list_head tr_databuf;
struct list_head tr_buf;
unsigned int tr_first;
struct list_head tr_ail1_list;
struct list_head tr_ail2_list;
};
struct gfs2_journal_extent {
struct list_head list;
unsigned int lblock; /* First logical block */
u64 dblock; /* First disk block */
u64 blocks;
};
struct gfs2_jdesc {
struct list_head jd_list;
struct list_head extent_list;
unsigned int nr_extents;
struct work_struct jd_work;
struct inode *jd_inode;
struct bio *jd_log_bio;
unsigned long jd_flags;
#define JDF_RECOVERY 1
unsigned int jd_jid;
u32 jd_blocks;
int jd_recover_error;
/* Replay stuff */
unsigned int jd_found_blocks;
unsigned int jd_found_revokes;
unsigned int jd_replayed_blocks;
struct list_head jd_revoke_list;
unsigned int jd_replay_tail;
u64 jd_no_addr;
};
struct gfs2_statfs_change_host {
s64 sc_total;
s64 sc_free;
s64 sc_dinodes;
};
#define GFS2_QUOTA_DEFAULT GFS2_QUOTA_OFF
#define GFS2_QUOTA_OFF 0
#define GFS2_QUOTA_ACCOUNT 1
#define GFS2_QUOTA_ON 2
#define GFS2_QUOTA_QUIET 3 /* on but not complaining */
#define GFS2_DATA_DEFAULT GFS2_DATA_ORDERED
#define GFS2_DATA_WRITEBACK 1
#define GFS2_DATA_ORDERED 2
#define GFS2_ERRORS_DEFAULT GFS2_ERRORS_WITHDRAW
#define GFS2_ERRORS_WITHDRAW 0
#define GFS2_ERRORS_CONTINUE 1 /* place holder for future feature */
#define GFS2_ERRORS_RO 2 /* place holder for future feature */
#define GFS2_ERRORS_PANIC 3
struct gfs2_args {
char ar_lockproto[GFS2_LOCKNAME_LEN]; /* Name of the Lock Protocol */
char ar_locktable[GFS2_LOCKNAME_LEN]; /* Name of the Lock Table */
char ar_hostdata[GFS2_LOCKNAME_LEN]; /* Host specific data */
unsigned int ar_spectator:1; /* Don't get a journal */
unsigned int ar_localflocks:1; /* Let the VFS do flock|fcntl */
unsigned int ar_debug:1; /* Oops on errors */
unsigned int ar_posix_acl:1; /* Enable posix acls */
unsigned int ar_quota:2; /* off/account/on */
unsigned int ar_suiddir:1; /* suiddir support */
unsigned int ar_data:2; /* ordered/writeback */
unsigned int ar_meta:1; /* mount metafs */
unsigned int ar_discard:1; /* discard requests */
unsigned int ar_errors:2; /* errors=withdraw | panic */
unsigned int ar_nobarrier:1; /* do not send barriers */
unsigned int ar_rgrplvb:1; /* use lvbs for rgrp info */
unsigned int ar_got_rgrplvb:1; /* Was the rgrplvb opt given? */
gfs2: change gfs2 readdir cookie gfs2 currently returns 31 bits of filename hash as a cookie that readdir uses for an offset into the directory. When there are a large number of directory entries, the likelihood of a collision goes up way too quickly. GFS2 will now return cookies that are guaranteed unique for a while, and then fail back to using 30 bits of filename hash. Specifically, the directory leaf blocks are divided up into chunks based on the minimum size of a gfs2 directory entry (48 bytes). Each entry's cookie is based off the chunk where it starts, in the linked list of leaf blocks that it hashes to (there are 131072 hash buckets). Directory entries will have unique names until they take reach chunk 8192. Assuming the largest filenames possible, and the least efficient spacing possible, this new method will still be able to return unique names when the previous method has statistically more than a 99% chance of a collision. The non-unique names it fails back to are guaranteed to not collide with the unique names. unique cookies will be in this format: - 1 bit "0" to make sure the the returned cookie is positive - 17 bits for the hash table index - 1 bit for the mode "0" - 13 bits for the offset non-unique cookies will be in this format: - 1 bit "0" to make sure the the returned cookie is positive - 17 bits for the hash table index - 1 bit for the mode "1" - 13 more bits of the name hash Another benefit of location based cookies, is that once a directory's exhash table is fully extended (so that multiple hash table indexs do not use the same leaf blocks), gfs2 can skip sorting the directory entries until it reaches the non-unique ones, and then it only needs to sort these. This provides a significant speed up for directory reads of very large directories. The only issue is that for these cookies to continue to point to the correct entry as files are added and removed from the directory, gfs2 must keep the entries at the same offset in the leaf block when they are split (see my previous patch). This means that until all the nodes in a cluster are running with code that will split the directory leaf blocks this way, none of the nodes can use the new cookie code. To deal with this, gfs2 now has the mount option loccookie, which, if set, will make it return these new location based cookies. This option must not be set until all nodes in the cluster are at least running this version of the kernel code, and you have guaranteed that there are no outstanding cookies required by other software, such as NFS. gfs2 uses some of the extra space at the end of the gfs2_dirent structure to store the calculated readdir cookies. This keeps us from needing to allocate a seperate array to hold these values. gfs2 recomputes the cookie stored in de_cookie for every readdir call. The time it takes to do so is small, and if gfs2 expected this value to be saved on disk, the new code wouldn't work correctly on filesystems created with an earlier version of gfs2. One issue with adding de_cookie to the union in the gfs2_dirent structure is that it caused the union to align itself to a 4 byte boundary, instead of its previous 2 byte boundary. This changed the offset of de_rahead. To solve that, I pulled de_rahead out of the union, since it does not need to be there. Signed-off-by: Benjamin Marzinski <bmarzins@redhat.com> Signed-off-by: Bob Peterson <rpeterso@redhat.com>
2015-12-01 07:46:55 -07:00
unsigned int ar_loccookie:1; /* use location based readdir
cookies */
s32 ar_commit; /* Commit interval */
s32 ar_statfs_quantum; /* The fast statfs interval */
s32 ar_quota_quantum; /* The quota interval */
s32 ar_statfs_percent; /* The % change to force sync */
};
struct gfs2_tune {
spinlock_t gt_spin;
unsigned int gt_logd_secs;
unsigned int gt_quota_warn_period; /* Secs between quota warn msgs */
unsigned int gt_quota_scale_num; /* Numerator */
unsigned int gt_quota_scale_den; /* Denominator */
unsigned int gt_quota_quantum; /* Secs between syncs to quota file */
unsigned int gt_new_files_jdata;
unsigned int gt_max_readahead; /* Max bytes to read-ahead from disk */
unsigned int gt_complain_secs;
unsigned int gt_statfs_quantum;
unsigned int gt_statfs_slow;
};
enum {
SDF_JOURNAL_CHECKED = 0,
SDF_JOURNAL_LIVE = 1,
SDF_WITHDRAWN = 2,
SDF_NOBARRIERS = 3,
SDF_NORECOVERY = 4,
SDF_DEMOTE = 5,
SDF_NOJOURNALID = 6,
SDF_RORECOVERY = 7, /* read only recovery */
SDF_SKIP_DLM_UNLOCK = 8,
SDF_FORCE_AIL_FLUSH = 9,
SDF_FREEZE_INITIATOR = 10,
SDF_WITHDRAWING = 11, /* Will withdraw eventually */
gfs2: Force withdraw to replay journals and wait for it to finish When a node withdraws from a file system, it often leaves its journal in an incomplete state. This is especially true when the withdraw is caused by io errors writing to the journal. Before this patch, a withdraw would try to write a "shutdown" record to the journal, tell dlm it's done with the file system, and none of the other nodes know about the problem. Later, when the problem is fixed and the withdrawn node is rebooted, it would then discover that its own journal was incomplete, and replay it. However, replaying it at this point is almost guaranteed to introduce corruption because the other nodes are likely to have used affected resource groups that appeared in the journal since the time of the withdraw. Replaying the journal later will overwrite any changes made, and not through any fault of dlm, which was instructed during the withdraw to release those resources. This patch makes file system withdraws seen by the entire cluster. Withdrawing nodes dequeue their journal glock to allow recovery. The remaining nodes check all the journals to see if they are clean or in need of replay. They try to replay dirty journals, but only the journals of withdrawn nodes will be "not busy" and therefore available for replay. Until the journal replay is complete, no i/o related glocks may be given out, to ensure that the replay does not cause the aforementioned corruption: We cannot allow any journal replay to overwrite blocks associated with a glock once it is held. The "live" glock which is now used to signal when a withdraw occurs. When a withdraw occurs, the node signals its withdraw by dequeueing the "live" glock and trying to enqueue it in EX mode, thus forcing the other nodes to all see a demote request, by way of a "1CB" (one callback) try lock. The "live" glock is not granted in EX; the callback is only just used to indicate a withdraw has occurred. Note that all nodes in the cluster must wait for the recovering node to finish replaying the withdrawing node's journal before continuing. To this end, it checks that the journals are clean multiple times in a retry loop. Also note that the withdraw function may be called from a wide variety of situations, and therefore, we need to take extra precautions to make sure pointers are valid before using them in many circumstances. We also need to take care when glocks decide to withdraw, since the withdraw code now uses glocks. Also, before this patch, if a process encountered an error and decided to withdraw, if another process was already withdrawing, the second withdraw would be silently ignored, which set it free to unlock its glocks. That's correct behavior if the original withdrawer encounters further errors down the road. But if secondary waiters don't wait for the journal replay, unlocking glocks will allow other nodes to use them, despite the fact that the journal containing those blocks is being replayed. The replay needs to finish before our glocks are released to other nodes. IOW, secondary withdraws need to wait for the first withdraw to finish. For example, if an rgrp glock is unlocked by a process that didn't wait for the first withdraw, a journal replay could introduce file system corruption by replaying a rgrp block that has already been granted to a different cluster node. Signed-off-by: Bob Peterson <rpeterso@redhat.com>
2020-01-28 12:23:45 -07:00
SDF_WITHDRAW_IN_PROG = 12, /* Withdraw is in progress */
SDF_REMOTE_WITHDRAW = 13, /* Performing remote recovery */
SDF_WITHDRAW_RECOVERY = 14, /* Wait for journal recovery when we are
withdrawing */
SDF_KILL = 15,
SDF_EVICTING = 16,
SDF_FROZEN = 17,
};
#define GFS2_FSNAME_LEN 256
struct gfs2_inum_host {
u64 no_formal_ino;
u64 no_addr;
};
struct gfs2_sb_host {
u32 sb_magic;
u32 sb_type;
u32 sb_fs_format;
u32 sb_multihost_format;
u32 sb_bsize;
u32 sb_bsize_shift;
struct gfs2_inum_host sb_master_dir;
struct gfs2_inum_host sb_root_dir;
char sb_lockproto[GFS2_LOCKNAME_LEN];
char sb_locktable[GFS2_LOCKNAME_LEN];
};
/*
* lm_mount() return values
*
* ls_jid - the journal ID this node should use
* ls_first - this node is the first to mount the file system
* ls_lockspace - lock module's context for this file system
* ls_ops - lock module's functions
*/
struct lm_lockstruct {
int ls_jid;
unsigned int ls_first;
const struct lm_lockops *ls_ops;
dlm_lockspace_t *ls_dlm;
int ls_recover_jid_done; /* These two are deprecated, */
int ls_recover_jid_status; /* used previously by gfs_controld */
struct dlm_lksb ls_mounted_lksb; /* mounted_lock */
struct dlm_lksb ls_control_lksb; /* control_lock */
char ls_control_lvb[GDLM_LVB_SIZE]; /* control_lock lvb */
struct completion ls_sync_wait; /* {control,mounted}_{lock,unlock} */
char *ls_lvb_bits;
spinlock_t ls_recover_spin; /* protects following fields */
unsigned long ls_recover_flags; /* DFL_ */
uint32_t ls_recover_mount; /* gen in first recover_done cb */
uint32_t ls_recover_start; /* gen in last recover_done cb */
uint32_t ls_recover_block; /* copy recover_start in last recover_prep */
uint32_t ls_recover_size; /* size of recover_submit, recover_result */
uint32_t *ls_recover_submit; /* gen in last recover_slot cb per jid */
uint32_t *ls_recover_result; /* result of last jid recovery */
};
GFS2: glock statistics gathering The stats are divided into two sets: those relating to the super block and those relating to an individual glock. The super block stats are done on a per cpu basis in order to try and reduce the overhead of gathering them. They are also further divided by glock type. In the case of both the super block and glock statistics, the same information is gathered in each case. The super block statistics are used to provide default values for most of the glock statistics, so that newly created glocks should have, as far as possible, a sensible starting point. The statistics are divided into three pairs of mean and variance, plus two counters. The mean/variance pairs are smoothed exponential estimates and the algorithm used is one which will be very familiar to those used to calculation of round trip times in network code. The three pairs of mean/variance measure the following things: 1. DLM lock time (non-blocking requests) 2. DLM lock time (blocking requests) 3. Inter-request time (again to the DLM) A non-blocking request is one which will complete right away, whatever the state of the DLM lock in question. That currently means any requests when (a) the current state of the lock is exclusive (b) the requested state is either null or unlocked or (c) the "try lock" flag is set. A blocking request covers all the other lock requests. There are two counters. The first is there primarily to show how many lock requests have been made, and thus how much data has gone into the mean/variance calculations. The other counter is counting queueing of holders at the top layer of the glock code. Hopefully that number will be a lot larger than the number of dlm lock requests issued. So why gather these statistics? There are several reasons we'd like to get a better idea of these timings: 1. To be able to better set the glock "min hold time" 2. To spot performance issues more easily 3. To improve the algorithm for selecting resource groups for allocation (to base it on lock wait time, rather than blindly using a "try lock") Due to the smoothing action of the updates, a step change in some input quantity being sampled will only fully be taken into account after 8 samples (or 4 for the variance) and this needs to be carefully considered when interpreting the results. Knowing both the time it takes a lock request to complete and the average time between lock requests for a glock means we can compute the total percentage of the time for which the node is able to use a glock vs. time that the rest of the cluster has its share. That will be very useful when setting the lock min hold time. The other point to remember is that all times are in nanoseconds. Great care has been taken to ensure that we measure exactly the quantities that we want, as accurately as possible. There are always inaccuracies in any measuring system, but I hope this is as accurate as we can reasonably make it. Signed-off-by: Steven Whitehouse <swhiteho@redhat.com>
2012-01-20 03:38:36 -07:00
struct gfs2_pcpu_lkstats {
/* One struct for each glock type */
struct gfs2_lkstats lkstats[10];
};
/* List of local (per node) statfs inodes */
struct local_statfs_inode {
struct list_head si_list;
struct inode *si_sc_inode;
unsigned int si_jid; /* journal id this statfs inode corresponds to */
};
struct gfs2_sbd {
struct super_block *sd_vfs;
GFS2: glock statistics gathering The stats are divided into two sets: those relating to the super block and those relating to an individual glock. The super block stats are done on a per cpu basis in order to try and reduce the overhead of gathering them. They are also further divided by glock type. In the case of both the super block and glock statistics, the same information is gathered in each case. The super block statistics are used to provide default values for most of the glock statistics, so that newly created glocks should have, as far as possible, a sensible starting point. The statistics are divided into three pairs of mean and variance, plus two counters. The mean/variance pairs are smoothed exponential estimates and the algorithm used is one which will be very familiar to those used to calculation of round trip times in network code. The three pairs of mean/variance measure the following things: 1. DLM lock time (non-blocking requests) 2. DLM lock time (blocking requests) 3. Inter-request time (again to the DLM) A non-blocking request is one which will complete right away, whatever the state of the DLM lock in question. That currently means any requests when (a) the current state of the lock is exclusive (b) the requested state is either null or unlocked or (c) the "try lock" flag is set. A blocking request covers all the other lock requests. There are two counters. The first is there primarily to show how many lock requests have been made, and thus how much data has gone into the mean/variance calculations. The other counter is counting queueing of holders at the top layer of the glock code. Hopefully that number will be a lot larger than the number of dlm lock requests issued. So why gather these statistics? There are several reasons we'd like to get a better idea of these timings: 1. To be able to better set the glock "min hold time" 2. To spot performance issues more easily 3. To improve the algorithm for selecting resource groups for allocation (to base it on lock wait time, rather than blindly using a "try lock") Due to the smoothing action of the updates, a step change in some input quantity being sampled will only fully be taken into account after 8 samples (or 4 for the variance) and this needs to be carefully considered when interpreting the results. Knowing both the time it takes a lock request to complete and the average time between lock requests for a glock means we can compute the total percentage of the time for which the node is able to use a glock vs. time that the rest of the cluster has its share. That will be very useful when setting the lock min hold time. The other point to remember is that all times are in nanoseconds. Great care has been taken to ensure that we measure exactly the quantities that we want, as accurately as possible. There are always inaccuracies in any measuring system, but I hope this is as accurate as we can reasonably make it. Signed-off-by: Steven Whitehouse <swhiteho@redhat.com>
2012-01-20 03:38:36 -07:00
struct gfs2_pcpu_lkstats __percpu *sd_lkstats;
struct kobject sd_kobj;
gfs2: use-after-free in sysfs deregistration syzkaller found the following splat with CONFIG_DEBUG_KOBJECT_RELEASE=y: Read of size 1 at addr ffff000028e896b8 by task kworker/1:2/228 CPU: 1 PID: 228 Comm: kworker/1:2 Tainted: G S 5.9.0-rc8+ #101 Hardware name: linux,dummy-virt (DT) Workqueue: events kobject_delayed_cleanup Call trace: dump_backtrace+0x0/0x4d8 show_stack+0x34/0x48 dump_stack+0x174/0x1f8 print_address_description.constprop.0+0x5c/0x550 kasan_report+0x13c/0x1c0 __asan_report_load1_noabort+0x34/0x60 memcmp+0xd0/0xd8 gfs2_uevent+0xc4/0x188 kobject_uevent_env+0x54c/0x1240 kobject_uevent+0x2c/0x40 __kobject_del+0x190/0x1d8 kobject_delayed_cleanup+0x2bc/0x3b8 process_one_work+0x96c/0x18c0 worker_thread+0x3f0/0xc30 kthread+0x390/0x498 ret_from_fork+0x10/0x18 Allocated by task 1110: kasan_save_stack+0x28/0x58 __kasan_kmalloc.isra.0+0xc8/0xe8 kasan_kmalloc+0x10/0x20 kmem_cache_alloc_trace+0x1d8/0x2f0 alloc_super+0x64/0x8c0 sget_fc+0x110/0x620 get_tree_bdev+0x190/0x648 gfs2_get_tree+0x50/0x228 vfs_get_tree+0x84/0x2e8 path_mount+0x1134/0x1da8 do_mount+0x124/0x138 __arm64_sys_mount+0x164/0x238 el0_svc_common.constprop.0+0x15c/0x598 do_el0_svc+0x60/0x150 el0_svc+0x34/0xb0 el0_sync_handler+0xc8/0x5b4 el0_sync+0x15c/0x180 Freed by task 228: kasan_save_stack+0x28/0x58 kasan_set_track+0x28/0x40 kasan_set_free_info+0x24/0x48 __kasan_slab_free+0x118/0x190 kasan_slab_free+0x14/0x20 slab_free_freelist_hook+0x6c/0x210 kfree+0x13c/0x460 Use the same pattern as f2fs + ext4 where the kobject destruction must complete before allowing the FS itself to be freed. This means that we need an explicit free_sbd in the callers. Cc: Bob Peterson <rpeterso@redhat.com> Cc: Andreas Gruenbacher <agruenba@redhat.com> Signed-off-by: Jamie Iles <jamie@nuviainc.com> [Also go to fail_free when init_names fails.] Signed-off-by: Andreas Gruenbacher <agruenba@redhat.com>
2020-10-12 06:13:09 -07:00
struct completion sd_kobj_unregister;
unsigned long sd_flags; /* SDF_... */
struct gfs2_sb_host sd_sb;
/* Constants computed on mount */
u32 sd_fsb2bb;
u32 sd_fsb2bb_shift;
u32 sd_diptrs; /* Number of pointers in a dinode */
u32 sd_inptrs; /* Number of pointers in a indirect block */
u32 sd_ldptrs; /* Number of pointers in a log descriptor block */
u32 sd_jbsize; /* Size of a journaled data block */
u32 sd_hash_bsize; /* sizeof(exhash block) */
u32 sd_hash_bsize_shift;
u32 sd_hash_ptrs; /* Number of pointers in a hash block */
u32 sd_qc_per_block;
GFS2: Speed up gfs2_rbm_from_block This patch is a rewrite of function gfs2_rbm_from_block. Rather than looping to find the right bitmap, the code now does a few simple math calculations. I compared the performance of both algorithms side by side and the new algorithm is noticeably faster. Sample instrumentation output from a "fast" machine: 5 million calls: millisec spent: Orig: 166 New: 113 5 million calls: millisec spent: Orig: 189 New: 114 In addition, I ran postmark (on a somewhat slowr CPU) before the after the new algorithm was put in place and postmark showed a decent improvement: Before the new algorithm: ------------------------- Time: 645 seconds total 584 seconds of transactions (171 per second) Files: 150087 created (232 per second) Creation alone: 100000 files (2083 per second) Mixed with transactions: 50087 files (85 per second) 49995 read (85 per second) 49991 appended (85 per second) 150087 deleted (232 per second) Deletion alone: 100174 files (7705 per second) Mixed with transactions: 49913 files (85 per second) Data: 273.42 megabytes read (434.08 kilobytes per second) 852.13 megabytes written (1.32 megabytes per second) With the new algorithm: ----------------------- Time: 599 seconds total 530 seconds of transactions (188 per second) Files: 150087 created (250 per second) Creation alone: 100000 files (1886 per second) Mixed with transactions: 50087 files (94 per second) 49995 read (94 per second) 49991 appended (94 per second) 150087 deleted (250 per second) Deletion alone: 100174 files (6260 per second) Mixed with transactions: 49913 files (94 per second) Data: 273.42 megabytes read (467.42 kilobytes per second) 852.13 megabytes written (1.42 megabytes per second) Signed-off-by: Bob Peterson <rpeterso@redhat.com> Signed-off-by: Steven Whitehouse <swhiteho@redhat.com>
2012-10-19 05:32:51 -07:00
u32 sd_blocks_per_bitmap;
u32 sd_max_dirres; /* Max blocks needed to add a directory entry */
u32 sd_max_height; /* Max height of a file's metadata tree */
u64 sd_heightsize[GFS2_MAX_META_HEIGHT + 1];
gfs2: change gfs2 readdir cookie gfs2 currently returns 31 bits of filename hash as a cookie that readdir uses for an offset into the directory. When there are a large number of directory entries, the likelihood of a collision goes up way too quickly. GFS2 will now return cookies that are guaranteed unique for a while, and then fail back to using 30 bits of filename hash. Specifically, the directory leaf blocks are divided up into chunks based on the minimum size of a gfs2 directory entry (48 bytes). Each entry's cookie is based off the chunk where it starts, in the linked list of leaf blocks that it hashes to (there are 131072 hash buckets). Directory entries will have unique names until they take reach chunk 8192. Assuming the largest filenames possible, and the least efficient spacing possible, this new method will still be able to return unique names when the previous method has statistically more than a 99% chance of a collision. The non-unique names it fails back to are guaranteed to not collide with the unique names. unique cookies will be in this format: - 1 bit "0" to make sure the the returned cookie is positive - 17 bits for the hash table index - 1 bit for the mode "0" - 13 bits for the offset non-unique cookies will be in this format: - 1 bit "0" to make sure the the returned cookie is positive - 17 bits for the hash table index - 1 bit for the mode "1" - 13 more bits of the name hash Another benefit of location based cookies, is that once a directory's exhash table is fully extended (so that multiple hash table indexs do not use the same leaf blocks), gfs2 can skip sorting the directory entries until it reaches the non-unique ones, and then it only needs to sort these. This provides a significant speed up for directory reads of very large directories. The only issue is that for these cookies to continue to point to the correct entry as files are added and removed from the directory, gfs2 must keep the entries at the same offset in the leaf block when they are split (see my previous patch). This means that until all the nodes in a cluster are running with code that will split the directory leaf blocks this way, none of the nodes can use the new cookie code. To deal with this, gfs2 now has the mount option loccookie, which, if set, will make it return these new location based cookies. This option must not be set until all nodes in the cluster are at least running this version of the kernel code, and you have guaranteed that there are no outstanding cookies required by other software, such as NFS. gfs2 uses some of the extra space at the end of the gfs2_dirent structure to store the calculated readdir cookies. This keeps us from needing to allocate a seperate array to hold these values. gfs2 recomputes the cookie stored in de_cookie for every readdir call. The time it takes to do so is small, and if gfs2 expected this value to be saved on disk, the new code wouldn't work correctly on filesystems created with an earlier version of gfs2. One issue with adding de_cookie to the union in the gfs2_dirent structure is that it caused the union to align itself to a 4 byte boundary, instead of its previous 2 byte boundary. This changed the offset of de_rahead. To solve that, I pulled de_rahead out of the union, since it does not need to be there. Signed-off-by: Benjamin Marzinski <bmarzins@redhat.com> Signed-off-by: Bob Peterson <rpeterso@redhat.com>
2015-12-01 07:46:55 -07:00
u32 sd_max_dents_per_leaf; /* Max number of dirents in a leaf block */
struct gfs2_args sd_args; /* Mount arguments */
struct gfs2_tune sd_tune; /* Filesystem tuning structure */
/* Lock Stuff */
struct lm_lockstruct sd_lockstruct;
struct gfs2_holder sd_live_gh;
struct gfs2_glock *sd_rename_gl;
GFS2: remove transaction glock GFS2 has a transaction glock, which must be grabbed for every transaction, whose purpose is to deal with freezing the filesystem. Aside from this involving a large amount of locking, it is very easy to make the current fsfreeze code hang on unfreezing. This patch rewrites how gfs2 handles freezing the filesystem. The transaction glock is removed. In it's place is a freeze glock, which is cached (but not held) in a shared state by every node in the cluster when the filesystem is mounted. This lock only needs to be grabbed on freezing, and actions which need to be safe from freezing, like recovery. When a node wants to freeze the filesystem, it grabs this glock exclusively. When the freeze glock state changes on the nodes (either from shared to unlocked, or shared to exclusive), the filesystem does a special log flush. gfs2_log_flush() does all the work for flushing out the and shutting down the incore log, and then it tries to grab the freeze glock in a shared state again. Since the filesystem is stuck in gfs2_log_flush, no new transaction can start, and nothing can be written to disk. Unfreezing the filesytem simply involes dropping the freeze glock, allowing gfs2_log_flush() to grab and then release the shared lock, so it is cached for next time. However, in order for the unfreezing ioctl to occur, gfs2 needs to get a shared lock on the filesystem root directory inode to check permissions. If that glock has already been grabbed exclusively, fsfreeze will be unable to get the shared lock and unfreeze the filesystem. In order to allow the unfreeze, this patch makes gfs2 grab a shared lock on the filesystem root directory during the freeze, and hold it until it unfreezes the filesystem. The functions which need to grab a shared lock in order to allow the unfreeze ioctl to be issued now use the lock grabbed by the freeze code instead. The freeze and unfreeze code take care to make sure that this shared lock will not be dropped while another process is using it. Signed-off-by: Benjamin Marzinski <bmarzins@redhat.com> Signed-off-by: Steven Whitehouse <swhiteho@redhat.com>
2014-05-01 20:26:55 -07:00
struct gfs2_glock *sd_freeze_gl;
struct work_struct sd_freeze_work;
wait_queue_head_t sd_kill_wait;
gfs2: Use async glocks for rename Because s_vfs_rename_mutex is not cluster-wide, multiple nodes can reverse the roles of which directories are "old" and which are "new" for the purposes of rename. This can cause deadlocks where two nodes end up waiting for each other. There can be several layers of directory dependencies across many nodes. This patch fixes the problem by acquiring all gfs2_rename's inode glocks asychronously and waiting for all glocks to be acquired. That way all inodes are locked regardless of the order. The timeout value for multiple asynchronous glocks is calculated to be the total of the individual wait times for each glock times two. Since gfs2_exchange is very similar to gfs2_rename, both functions are patched in the same way. A new async glock wait queue, sd_async_glock_wait, keeps a list of waiters for these events. If gfs2's holder_wake function detects an async holder, it wakes up any waiters for the event. The waiter only tests whether any of its requests are still pending. Since the glocks are sent to dlm asychronously, the wait function needs to check to see which glocks, if any, were granted. If a glock is granted by dlm (and therefore held), its minimum hold time is checked and adjusted as necessary, as other glock grants do. If the event times out, all glocks held thus far must be dequeued to resolve any existing deadlocks. Then, if there are any outstanding locking requests, we need to loop around and wait for dlm to respond to those requests too. After we release all requests, we return -ESTALE to the caller (vfs rename) which loops around and retries the request. Node1 Node2 --------- --------- 1. Enqueue A Enqueue B 2. Enqueue B Enqueue A 3. A granted 6. B granted 7. Wait for B 8. Wait for A 9. A times out (since Node 1 holds A) 10. Dequeue B (since it was granted) 11. Wait for all requests from DLM 12. B Granted (since Node2 released it in step 10) 13. Rename 14. Dequeue A 15. DLM Grants A 16. Dequeue A (due to the timeout and since we no longer have B held for our task). 17. Dequeue B 18. Return -ESTALE to vfs 19. VFS retries the operation, goto step 1. This release-all-locks / acquire-all-locks may slow rename / exchange down as both nodes struggle in the same way and do the same thing. However, this will only happen when there is contention for the same inodes, which ought to be rare. Signed-off-by: Bob Peterson <rpeterso@redhat.com> Signed-off-by: Andreas Gruenbacher <agruenba@redhat.com>
2019-08-30 10:31:02 -07:00
wait_queue_head_t sd_async_glock_wait;
atomic_t sd_glock_disposal;
struct completion sd_locking_init;
struct completion sd_wdack;
struct delayed_work sd_control_work;
/* Inode Stuff */
struct dentry *sd_master_dir;
struct dentry *sd_root_dir;
struct inode *sd_jindex;
struct inode *sd_statfs_inode;
struct inode *sd_sc_inode;
struct list_head sd_sc_inodes_list;
struct inode *sd_qc_inode;
struct inode *sd_rindex;
struct inode *sd_quota_inode;
/* StatFS stuff */
spinlock_t sd_statfs_spin;
struct gfs2_statfs_change_host sd_statfs_master;
struct gfs2_statfs_change_host sd_statfs_local;
int sd_statfs_force_sync;
/* Resource group stuff */
int sd_rindex_uptodate;
spinlock_t sd_rindex_spin;
GFS2: Use rbtree for resource groups and clean up bitmap buffer ref count scheme Here is an update of Bob's original rbtree patch which, in addition, also resolves the rather strange ref counting that was being done relating to the bitmap blocks. Originally we had a dual system for journaling resource groups. The metadata blocks were journaled and also the rgrp itself was added to a list. The reason for adding the rgrp to the list in the journal was so that the "repolish clones" code could be run to update the free space, and potentially send any discard requests when the log was flushed. This was done by comparing the "cloned" bitmap with what had been written back on disk during the transaction commit. Due to this, there was a requirement to hang on to the rgrps' bitmap buffers until the journal had been flushed. For that reason, there was a rather complicated set up in the ->go_lock ->go_unlock functions for rgrps involving both a mutex and a spinlock (the ->sd_rindex_spin) to maintain a reference count on the buffers. However, the journal maintains a reference count on the buffers anyway, since they are being journaled as metadata buffers. So by moving the code which deals with the post-journal accounting for bitmap blocks to the metadata journaling code, we can entirely dispense with the rather strange buffer ref counting scheme and also the requirement to journal the rgrps. The net result of all this is that the ->sd_rindex_spin is left to do exactly one job, and that is to look after the rbtree or rgrps. This patch is designed to be a stepping stone towards using RCU for the rbtree of resource groups, however the reduction in the number of uses of the ->sd_rindex_spin is likely to have benefits for multi-threaded workloads, anyway. The patch retains ->go_lock and ->go_unlock for rgrps, however these maybe also be removed in future in favour of calling the functions directly where required in the code. That will allow locking of resource groups without needing to actually read them in - something that could be useful in speeding up statfs. In the mean time though it is valid to dereference ->bi_bh only when the rgrp is locked. This is basically the same rule as before, modulo the references not being valid until the following journal flush. Signed-off-by: Steven Whitehouse <swhiteho@redhat.com> Signed-off-by: Bob Peterson <rpeterso@redhat.com> Cc: Benjamin Marzinski <bmarzins@redhat.com>
2011-08-31 01:53:19 -07:00
struct rb_root sd_rindex_tree;
unsigned int sd_rgrps;
unsigned int sd_max_rg_data;
/* Journal index stuff */
struct list_head sd_jindex_list;
spinlock_t sd_jindex_spin;
struct mutex sd_jindex_mutex;
unsigned int sd_journals;
struct gfs2_jdesc *sd_jdesc;
struct gfs2_holder sd_journal_gh;
struct gfs2_holder sd_jinode_gh;
gfs2: Force withdraw to replay journals and wait for it to finish When a node withdraws from a file system, it often leaves its journal in an incomplete state. This is especially true when the withdraw is caused by io errors writing to the journal. Before this patch, a withdraw would try to write a "shutdown" record to the journal, tell dlm it's done with the file system, and none of the other nodes know about the problem. Later, when the problem is fixed and the withdrawn node is rebooted, it would then discover that its own journal was incomplete, and replay it. However, replaying it at this point is almost guaranteed to introduce corruption because the other nodes are likely to have used affected resource groups that appeared in the journal since the time of the withdraw. Replaying the journal later will overwrite any changes made, and not through any fault of dlm, which was instructed during the withdraw to release those resources. This patch makes file system withdraws seen by the entire cluster. Withdrawing nodes dequeue their journal glock to allow recovery. The remaining nodes check all the journals to see if they are clean or in need of replay. They try to replay dirty journals, but only the journals of withdrawn nodes will be "not busy" and therefore available for replay. Until the journal replay is complete, no i/o related glocks may be given out, to ensure that the replay does not cause the aforementioned corruption: We cannot allow any journal replay to overwrite blocks associated with a glock once it is held. The "live" glock which is now used to signal when a withdraw occurs. When a withdraw occurs, the node signals its withdraw by dequeueing the "live" glock and trying to enqueue it in EX mode, thus forcing the other nodes to all see a demote request, by way of a "1CB" (one callback) try lock. The "live" glock is not granted in EX; the callback is only just used to indicate a withdraw has occurred. Note that all nodes in the cluster must wait for the recovering node to finish replaying the withdrawing node's journal before continuing. To this end, it checks that the journals are clean multiple times in a retry loop. Also note that the withdraw function may be called from a wide variety of situations, and therefore, we need to take extra precautions to make sure pointers are valid before using them in many circumstances. We also need to take care when glocks decide to withdraw, since the withdraw code now uses glocks. Also, before this patch, if a process encountered an error and decided to withdraw, if another process was already withdrawing, the second withdraw would be silently ignored, which set it free to unlock its glocks. That's correct behavior if the original withdrawer encounters further errors down the road. But if secondary waiters don't wait for the journal replay, unlocking glocks will allow other nodes to use them, despite the fact that the journal containing those blocks is being replayed. The replay needs to finish before our glocks are released to other nodes. IOW, secondary withdraws need to wait for the first withdraw to finish. For example, if an rgrp glock is unlocked by a process that didn't wait for the first withdraw, a journal replay could introduce file system corruption by replaying a rgrp block that has already been granted to a different cluster node. Signed-off-by: Bob Peterson <rpeterso@redhat.com>
2020-01-28 12:23:45 -07:00
struct gfs2_glock *sd_jinode_gl;
struct gfs2_holder sd_sc_gh;
struct buffer_head *sd_sc_bh;
struct gfs2_holder sd_qc_gh;
struct completion sd_journal_ready;
/* Workqueue stuff */
struct workqueue_struct *sd_glock_wq;
struct workqueue_struct *sd_delete_wq;
/* Daemon stuff */
struct task_struct *sd_logd_process;
struct task_struct *sd_quotad_process;
/* Quota stuff */
struct list_head sd_quota_list;
atomic_t sd_quota_count;
struct mutex sd_quota_sync_mutex;
wait_queue_head_t sd_quota_wait;
unsigned int sd_quota_slots;
unsigned long *sd_quota_bitmap;
spinlock_t sd_bitmap_lock;
u64 sd_quota_sync_gen;
/* Log stuff */
struct address_space sd_aspace;
spinlock_t sd_log_lock;
struct gfs2_trans *sd_log_tr;
unsigned int sd_log_blks_reserved;
atomic_t sd_log_pinned;
unsigned int sd_log_num_revoke;
struct list_head sd_log_revokes;
struct list_head sd_log_ordered;
spinlock_t sd_ordered_lock;
atomic_t sd_log_thresh1;
atomic_t sd_log_thresh2;
atomic_t sd_log_blks_free;
atomic_t sd_log_blks_needed;
atomic_t sd_log_revokes_available;
wait_queue_head_t sd_log_waitq;
wait_queue_head_t sd_logd_waitq;
u64 sd_log_sequence;
int sd_log_idle;
struct rw_semaphore sd_log_flush_lock;
atomic_t sd_log_in_flight;
wait_queue_head_t sd_log_flush_wait;
int sd_log_error; /* First log error */
gfs2: Force withdraw to replay journals and wait for it to finish When a node withdraws from a file system, it often leaves its journal in an incomplete state. This is especially true when the withdraw is caused by io errors writing to the journal. Before this patch, a withdraw would try to write a "shutdown" record to the journal, tell dlm it's done with the file system, and none of the other nodes know about the problem. Later, when the problem is fixed and the withdrawn node is rebooted, it would then discover that its own journal was incomplete, and replay it. However, replaying it at this point is almost guaranteed to introduce corruption because the other nodes are likely to have used affected resource groups that appeared in the journal since the time of the withdraw. Replaying the journal later will overwrite any changes made, and not through any fault of dlm, which was instructed during the withdraw to release those resources. This patch makes file system withdraws seen by the entire cluster. Withdrawing nodes dequeue their journal glock to allow recovery. The remaining nodes check all the journals to see if they are clean or in need of replay. They try to replay dirty journals, but only the journals of withdrawn nodes will be "not busy" and therefore available for replay. Until the journal replay is complete, no i/o related glocks may be given out, to ensure that the replay does not cause the aforementioned corruption: We cannot allow any journal replay to overwrite blocks associated with a glock once it is held. The "live" glock which is now used to signal when a withdraw occurs. When a withdraw occurs, the node signals its withdraw by dequeueing the "live" glock and trying to enqueue it in EX mode, thus forcing the other nodes to all see a demote request, by way of a "1CB" (one callback) try lock. The "live" glock is not granted in EX; the callback is only just used to indicate a withdraw has occurred. Note that all nodes in the cluster must wait for the recovering node to finish replaying the withdrawing node's journal before continuing. To this end, it checks that the journals are clean multiple times in a retry loop. Also note that the withdraw function may be called from a wide variety of situations, and therefore, we need to take extra precautions to make sure pointers are valid before using them in many circumstances. We also need to take care when glocks decide to withdraw, since the withdraw code now uses glocks. Also, before this patch, if a process encountered an error and decided to withdraw, if another process was already withdrawing, the second withdraw would be silently ignored, which set it free to unlock its glocks. That's correct behavior if the original withdrawer encounters further errors down the road. But if secondary waiters don't wait for the journal replay, unlocking glocks will allow other nodes to use them, despite the fact that the journal containing those blocks is being replayed. The replay needs to finish before our glocks are released to other nodes. IOW, secondary withdraws need to wait for the first withdraw to finish. For example, if an rgrp glock is unlocked by a process that didn't wait for the first withdraw, a journal replay could introduce file system corruption by replaying a rgrp block that has already been granted to a different cluster node. Signed-off-by: Bob Peterson <rpeterso@redhat.com>
2020-01-28 12:23:45 -07:00
wait_queue_head_t sd_withdraw_wait;
unsigned int sd_log_tail;
unsigned int sd_log_flush_tail;
unsigned int sd_log_head;
unsigned int sd_log_flush_head;
spinlock_t sd_ail_lock;
struct list_head sd_ail1_list;
struct list_head sd_ail2_list;
/* For quiescing the filesystem */
struct gfs2_holder sd_freeze_gh;
struct mutex sd_freeze_mutex;
struct list_head sd_dead_glocks;
char sd_fsname[GFS2_FSNAME_LEN + 3 * sizeof(int) + 2];
char sd_table_name[GFS2_FSNAME_LEN];
char sd_proto_name[GFS2_FSNAME_LEN];
/* Debugging crud */
unsigned long sd_last_warning;
struct dentry *debugfs_dir; /* debugfs directory */
gfs2: Force withdraw to replay journals and wait for it to finish When a node withdraws from a file system, it often leaves its journal in an incomplete state. This is especially true when the withdraw is caused by io errors writing to the journal. Before this patch, a withdraw would try to write a "shutdown" record to the journal, tell dlm it's done with the file system, and none of the other nodes know about the problem. Later, when the problem is fixed and the withdrawn node is rebooted, it would then discover that its own journal was incomplete, and replay it. However, replaying it at this point is almost guaranteed to introduce corruption because the other nodes are likely to have used affected resource groups that appeared in the journal since the time of the withdraw. Replaying the journal later will overwrite any changes made, and not through any fault of dlm, which was instructed during the withdraw to release those resources. This patch makes file system withdraws seen by the entire cluster. Withdrawing nodes dequeue their journal glock to allow recovery. The remaining nodes check all the journals to see if they are clean or in need of replay. They try to replay dirty journals, but only the journals of withdrawn nodes will be "not busy" and therefore available for replay. Until the journal replay is complete, no i/o related glocks may be given out, to ensure that the replay does not cause the aforementioned corruption: We cannot allow any journal replay to overwrite blocks associated with a glock once it is held. The "live" glock which is now used to signal when a withdraw occurs. When a withdraw occurs, the node signals its withdraw by dequeueing the "live" glock and trying to enqueue it in EX mode, thus forcing the other nodes to all see a demote request, by way of a "1CB" (one callback) try lock. The "live" glock is not granted in EX; the callback is only just used to indicate a withdraw has occurred. Note that all nodes in the cluster must wait for the recovering node to finish replaying the withdrawing node's journal before continuing. To this end, it checks that the journals are clean multiple times in a retry loop. Also note that the withdraw function may be called from a wide variety of situations, and therefore, we need to take extra precautions to make sure pointers are valid before using them in many circumstances. We also need to take care when glocks decide to withdraw, since the withdraw code now uses glocks. Also, before this patch, if a process encountered an error and decided to withdraw, if another process was already withdrawing, the second withdraw would be silently ignored, which set it free to unlock its glocks. That's correct behavior if the original withdrawer encounters further errors down the road. But if secondary waiters don't wait for the journal replay, unlocking glocks will allow other nodes to use them, despite the fact that the journal containing those blocks is being replayed. The replay needs to finish before our glocks are released to other nodes. IOW, secondary withdraws need to wait for the first withdraw to finish. For example, if an rgrp glock is unlocked by a process that didn't wait for the first withdraw, a journal replay could introduce file system corruption by replaying a rgrp block that has already been granted to a different cluster node. Signed-off-by: Bob Peterson <rpeterso@redhat.com>
2020-01-28 12:23:45 -07:00
unsigned long sd_glock_dqs_held;
};
GFS2: glock statistics gathering The stats are divided into two sets: those relating to the super block and those relating to an individual glock. The super block stats are done on a per cpu basis in order to try and reduce the overhead of gathering them. They are also further divided by glock type. In the case of both the super block and glock statistics, the same information is gathered in each case. The super block statistics are used to provide default values for most of the glock statistics, so that newly created glocks should have, as far as possible, a sensible starting point. The statistics are divided into three pairs of mean and variance, plus two counters. The mean/variance pairs are smoothed exponential estimates and the algorithm used is one which will be very familiar to those used to calculation of round trip times in network code. The three pairs of mean/variance measure the following things: 1. DLM lock time (non-blocking requests) 2. DLM lock time (blocking requests) 3. Inter-request time (again to the DLM) A non-blocking request is one which will complete right away, whatever the state of the DLM lock in question. That currently means any requests when (a) the current state of the lock is exclusive (b) the requested state is either null or unlocked or (c) the "try lock" flag is set. A blocking request covers all the other lock requests. There are two counters. The first is there primarily to show how many lock requests have been made, and thus how much data has gone into the mean/variance calculations. The other counter is counting queueing of holders at the top layer of the glock code. Hopefully that number will be a lot larger than the number of dlm lock requests issued. So why gather these statistics? There are several reasons we'd like to get a better idea of these timings: 1. To be able to better set the glock "min hold time" 2. To spot performance issues more easily 3. To improve the algorithm for selecting resource groups for allocation (to base it on lock wait time, rather than blindly using a "try lock") Due to the smoothing action of the updates, a step change in some input quantity being sampled will only fully be taken into account after 8 samples (or 4 for the variance) and this needs to be carefully considered when interpreting the results. Knowing both the time it takes a lock request to complete and the average time between lock requests for a glock means we can compute the total percentage of the time for which the node is able to use a glock vs. time that the rest of the cluster has its share. That will be very useful when setting the lock min hold time. The other point to remember is that all times are in nanoseconds. Great care has been taken to ensure that we measure exactly the quantities that we want, as accurately as possible. There are always inaccuracies in any measuring system, but I hope this is as accurate as we can reasonably make it. Signed-off-by: Steven Whitehouse <swhiteho@redhat.com>
2012-01-20 03:38:36 -07:00
static inline void gfs2_glstats_inc(struct gfs2_glock *gl, int which)
{
gl->gl_stats.stats[which]++;
}
static inline void gfs2_sbstats_inc(const struct gfs2_glock *gl, int which)
{
const struct gfs2_sbd *sdp = gl->gl_name.ln_sbd;
GFS2: glock statistics gathering The stats are divided into two sets: those relating to the super block and those relating to an individual glock. The super block stats are done on a per cpu basis in order to try and reduce the overhead of gathering them. They are also further divided by glock type. In the case of both the super block and glock statistics, the same information is gathered in each case. The super block statistics are used to provide default values for most of the glock statistics, so that newly created glocks should have, as far as possible, a sensible starting point. The statistics are divided into three pairs of mean and variance, plus two counters. The mean/variance pairs are smoothed exponential estimates and the algorithm used is one which will be very familiar to those used to calculation of round trip times in network code. The three pairs of mean/variance measure the following things: 1. DLM lock time (non-blocking requests) 2. DLM lock time (blocking requests) 3. Inter-request time (again to the DLM) A non-blocking request is one which will complete right away, whatever the state of the DLM lock in question. That currently means any requests when (a) the current state of the lock is exclusive (b) the requested state is either null or unlocked or (c) the "try lock" flag is set. A blocking request covers all the other lock requests. There are two counters. The first is there primarily to show how many lock requests have been made, and thus how much data has gone into the mean/variance calculations. The other counter is counting queueing of holders at the top layer of the glock code. Hopefully that number will be a lot larger than the number of dlm lock requests issued. So why gather these statistics? There are several reasons we'd like to get a better idea of these timings: 1. To be able to better set the glock "min hold time" 2. To spot performance issues more easily 3. To improve the algorithm for selecting resource groups for allocation (to base it on lock wait time, rather than blindly using a "try lock") Due to the smoothing action of the updates, a step change in some input quantity being sampled will only fully be taken into account after 8 samples (or 4 for the variance) and this needs to be carefully considered when interpreting the results. Knowing both the time it takes a lock request to complete and the average time between lock requests for a glock means we can compute the total percentage of the time for which the node is able to use a glock vs. time that the rest of the cluster has its share. That will be very useful when setting the lock min hold time. The other point to remember is that all times are in nanoseconds. Great care has been taken to ensure that we measure exactly the quantities that we want, as accurately as possible. There are always inaccuracies in any measuring system, but I hope this is as accurate as we can reasonably make it. Signed-off-by: Steven Whitehouse <swhiteho@redhat.com>
2012-01-20 03:38:36 -07:00
preempt_disable();
this_cpu_ptr(sdp->sd_lkstats)->lkstats[gl->gl_name.ln_type].stats[which]++;
preempt_enable();
}
struct gfs2_rgrpd *gfs2_glock2rgrp(struct gfs2_glock *gl);
static inline unsigned gfs2_max_stuffed_size(const struct gfs2_inode *ip)
{
return GFS2_SB(&ip->i_inode)->sd_sb.sb_bsize - sizeof(struct gfs2_dinode);
}
#endif /* __INCORE_DOT_H__ */