- 07 Jul, 2020 10 commits
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Dave Chinner authored
When we attempt to reclaim an inode, the first thing we do is take the inode lock. This is blocking right now, so if the inode being accessed by something else (e.g. being flushed to the cluster buffer) we will block here. Change this to a trylock so that we do not block inode reclaim unnecessarily here. Signed-off-by:
Dave Chinner <dchinner@redhat.com> Reviewed-by:
Darrick J. Wong <darrick.wong@oracle.com> Reviewed-by:
Brian Foster <bfoster@redhat.com> Signed-off-by:
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Dave Chinner authored
Inode reclaim will still throttle direct reclaim on the per-ag reclaim locks. This is no longer necessary as reclaim can run non-blocking now. Hence we can remove these locks so that we don't arbitrarily block reclaimers just because there are more direct reclaimers than there are AGs. This can result in multiple reclaimers working on the same range of an AG, but this doesn't cause any apparent issues. Optimising the spread of concurrent reclaimers for best efficiency can be done in a future patchset. Signed-off-by:
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Dave Chinner authored
We no longer need to issue IO from shrinker based inode reclaim to prevent spurious OOM killer invocation. This leaves only the global filesystem management operations such as unmount needing to writeback dirty inodes and reclaim them. Instead of using the reclaim pass to write dirty inodes before reclaiming them, use the AIL to push all the dirty inodes before we try to reclaim them. This allows us to remove all the conditional SYNC_WAIT locking and the writeback code from xfs_reclaim_inode() and greatly simplify the checks we need to do to reclaim an inode. Signed-off-by:
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Dave Chinner authored
Now that dirty inode writeback doesn't cause read-modify-write cycles on the inode cluster buffer under memory pressure, the need to throttle memory reclaim to the rate at which we can clean dirty inodes goes away. That is due to the fact that we no longer thrash inode cluster buffers under memory pressure to clean dirty inodes. This means inode writeback no longer stalls on memory allocation or read IO, and hence can be done asynchronously without generating memory pressure. As a result, blocking inode writeback in reclaim is no longer necessary to prevent reclaim priority windup as cleaning dirty inodes is no longer dependent on having memory reserves available for the filesystem to make progress reclaiming inodes. Hence we can convert inode reclaim to be non-blocking for shrinker callouts, both for direct reclaim and kswapd. On a vanilla kernel, running a 16-way fsmark create workload on a 4 node/16p/16GB RAM machine, I can reliably pin 14.75GB of RAM via userspace mlock(). The OOM killer gets invoked at 15GB of pinned RAM. Without the inode cluster pinning, this non-blocking reclaim patch triggers premature OOM killer invocation with the same memory pinning, sometimes with as much as 45% of RAM being free. It's trivially easy to trigger the OOM killer when reclaim does not block. With pinning inode clusters in RAM and then adding this patch, I can reliably pin 14.5GB of RAM and still have the fsmark workload run to completion. The OOM killer gets invoked 14.75GB of pinned RAM, which is only a small amount of memory less than the vanilla kernel. It is much more reliable than just with async reclaim alone. simoops shows that allocation stalls go away when async reclaim is used. Vanilla kernel: Run time: 1924 seconds Read latency (p50: 3,305,472) (p95: 3,723,264) (p99: 4,001,792) Write latency (p50: 184,064) (p95: 553,984) (p99: 807,936) Allocation latency (p50: 2,641,920) (p95: 3,911,680) (p99: 4,464,640) work rate = 13.45/sec (avg 13.44/sec) (p50: 13.46) (p95: 13.58) (p99: 13.70) alloc stall rate = 3.80/sec (avg: 2.59) (p50: 2.54) (p95: 2.96) (p99: 3.02) With inode cluster pinning and async reclaim: Run time: 1924 seconds Read latency (p50: 3,305,472) (p95: 3,715,072) (p99: 3,977,216) Write latency (p50: 187,648) (p95: 553,984) (p99: 789,504) Allocation latency (p50: 2,748,416) (p95: 3,919,872) (p99: 4,448,256) work rate = 13.28/sec (avg 13.32/sec) (p50: 13.26) (p95: 13.34) (p99: 13.34) alloc stall rate = 0.02/sec (avg: 0.02) (p50: 0.01) (p95: 0.03) (p99: 0.03) Latencies don't really change much, nor does the work rate. However, allocation almost never stalls with these changes, whilst the vanilla kernel is sometimes reporting 20 stalls/s over a 60s sample period. This difference is due to inode reclaim being largely non-blocking now. IOWs, once we have pinned inode cluster buffers, we can make inode reclaim non-blocking without a major risk of premature and/or spurious OOM killer invocation, and without any changes to memory reclaim infrastructure. Signed-off-by:
Amir Goldstein <amir73il@gmail.com> Reviewed-by:
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Dave Chinner authored
When we dirty an inode, we are going to have to write it disk at some point in the near future. This requires the inode cluster backing buffer to be present in memory. Unfortunately, under severe memory pressure we can reclaim the inode backing buffer while the inode is dirty in memory, resulting in stalling the AIL pushing because it has to do a read-modify-write cycle on the cluster buffer. When we have no memory available, the read of the cluster buffer blocks the AIL pushing process, and this causes all sorts of issues for memory reclaim as it requires inode writeback to make forwards progress. Allocating a cluster buffer causes more memory pressure, and results in more cluster buffers to be reclaimed, resulting in more RMW cycles to be done in the AIL context and everything then backs up on AIL progress. Only the synchronous inode cluster writeback in the the inode reclaim code provides some level of forwards progress guarantees that prevent OOM-killer rampages in this situation. Fix this by pinning the inode backing buffer to the inode log item when the inode is first dirtied (i.e. in xfs_trans_log_inode()). This may mean the first modification of an inode that has been held in cache for a long time may block on a cluster buffer read, but we can do that in transaction context and block safely until the buffer has been allocated and read. Once we have the cluster buffer, the inode log item takes a reference to it, pinning it in memory, and attaches it to the log item for future reference. This means we can always grab the cluster buffer from the inode log item when we need it. When the inode is finally cleaned and removed from the AIL, we can drop the reference the inode log item holds on the cluster buffer. Once all inodes on the cluster buffer are clean, the cluster buffer will be unpinned and it will be available for memory reclaim to reclaim again. This avoids the issues with needing to do RMW cycles in the AIL pushing context, and hence allows complete non-blocking inode flushing to be performed by the AIL pushing context. Signed-off-by:
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Dave Chinner authored
xfs_ail_delete_one() is called directly from dquot and inode IO completion, as well as from the generic xfs_trans_ail_delete() function. Inodes are about to have their own failure handling, and dquots will in future, too. Pull the clearing of the LI_FAILED flag up into the callers so we can customise the code appropriately. Signed-off-by:
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Dave Chinner authored
When an buffer IO error occurs, we want to mark all the log items attached to the buffer as failed. Open code the error handling loop so that we can modify the flagging for the different types of objects directly and independently of each other. This also allows us to remove the ->iop_error method from the log item operations. Signed-off-by:
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Dave Chinner authored
Currently when a buffer with attached log items has an IO error it called ->iop_error for each attched log item. These all call xfs_set_li_failed() to handle the error, but we are about to change the way log items manage buffers. hence we first need to remove the per-item dependency on buffer handling done by xfs_set_li_failed(). We already have specific buffer type IO completion routines, so move the log item error handling out of the generic error handling and into the log item specific functions so we can implement per-type error handling easily. This requires a more complex return value from the error handling code so that we can take the correct action the failure handling requires. This results in some repeated boilerplate in the functions, but that can be cleaned up later once all the changes cascade through this code. Signed-off-by:
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Dave Chinner authored
They are not used anymore, so remove them from the log item and the buffer iodone attachment interfaces. Signed-off-by:
Christoph Hellwig <hch@lst.de> Reviewed-by:
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Dave Chinner authored
Now that we've sorted inode and dquot buffers, we can apply the same cleanups to dirty buffers with buffer log items. They only have one callback, too, so we don't need the log item callback. Collapse the iodone functions and remove all the now unnecessary infrastructure around callback processing. Signed-off-by:
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- 06 Jul, 2020 24 commits
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Dave Chinner authored
Similar to inodes, we can call the dquot IO completion functions directly from the buffer completion code, removing another user of log item callbacks for IO completion processing. Signed-off-by:
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Dave Chinner authored
Having different io completion callbacks for different inode states makes things complex. We can detect if the inode is stale via the XFS_ISTALE flag in IO completion, so we don't need a special callback just for this. This means inodes only have a single iodone callback, and inode IO completion is entirely buffer centric at this point. Hence we no longer need to use a log item callback at all as we can just call xfs_iflush_done() directly from the buffer completions and walk the buffer log item list to complete the all inodes under IO. Signed-off-by:
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Dave Chinner authored
When we've emptied the buffer log item list, it does a list_del_init on itself to reset it's pointers to itself. This is unnecessary as the list is already empty at this point - it was a left-over fragment from the list_head conversion of the buffer log item list. Remove them. Signed-off-by:
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Dave Chinner authored
All unmarked dirty buffers should be in the AIL and have log items attached to them. Hence when they are written, we will run a callback to remove the item from the AIL if appropriate. Now that we've handled inode and dquot buffers, all remaining calls are to xfs_buf_iodone() and so we can hard code this rather than use an indirect call. Signed-off-by:
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Dave Chinner authored
Log recovery has it's own buffer write completion handler for buffers that it directly recovers. Convert these to direct calls by flagging these buffers as being log recovery buffers. The flag will get cleared by the log recovery IO completion routine, so it will never leak out of log recovery. Signed-off-by:
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Dave Chinner authored
dquot buffers always have write IO callbacks, so by marking them directly we can avoid needing to attach ->b_iodone functions to them. This avoids an indirect call, and makes future modifications much simpler. This is largely a rearrangement of the code at this point - no IO completion functionality changes at this point, just how the code is run is modified. Signed-off-by:
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Dave Chinner authored
Inode buffers always have write IO callbacks, so by marking them directly we can avoid needing to attach ->b_iodone functions to them. This avoids an indirect call, and makes future modifications much simpler. While this is largely a refactor of existing functionality, we broaden the scope of the flag to beyond where inodes are explicitly attached because future changes need to know what type of log items are attached to the buffer. Adding this buffer flag may invoke the inode iodone callback in cases where it wouldn't have been previously, but this is not a functional change because the callback is identical to the normal buffer write iodone callback when inodes are not attached. Signed-off-by:
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Dave Chinner authored
The inode log item is kind of special in that it can be aggregating new changes in memory at the same time time existing changes are being written back to disk. This means there are fields in the log item that are accessed concurrently from contexts that don't share any locking at all. e.g. updating ili_last_fields occurs at flush time under the ILOCK_EXCL and flush lock at flush time, under the flush lock at IO completion time, and is read under the ILOCK_EXCL when the inode is logged. Hence there is no actual serialisation between reading the field during logging of the inode in transactions vs clearing the field in IO completion. We currently get away with this by the fact that we are only clearing fields in IO completion, and nothing bad happens if we accidentally log more of the inode than we actually modify. Worst case is we consume a tiny bit more memory and log bandwidth. However, if we want to do more complex state manipulations on the log item that requires updates at all three of these potential locations, we need to have some mechanism of serialising those operations. To do this, introduce a spinlock into the log item to serialise internal state. This could be done via the xfs_inode i_flags_lock, but this then leads to potential lock inversion issues where inode flag updates need to occur inside locks that best nest inside the inode log item locks (e.g. marking inodes stale during inode cluster freeing). Using a separate spinlock avoids these sorts of problems and simplifies future code. This does not touch the use of ili_fields in the item formatting code - that is entirely protected by the ILOCK_EXCL at this point in time, so it remains untouched. Signed-off-by:
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Dave Chinner authored
This was used to track if the item had logged fields being flushed to disk. We log everything in the inode these days, so this logic is no longer needed. Remove it. Signed-off-by:
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Dave Chinner authored
In tracking down a problem in this patchset, I discovered we are reclaiming dirty stale inodes. This wasn't discovered until inodes were always attached to the cluster buffer and then the rcu callback that freed inodes was assert failing because the inode still had an active pointer to the cluster buffer after it had been reclaimed. Debugging the issue indicated that this was a pre-existing issue resulting from the way the inodes are handled in xfs_inactive_ifree. When we free a cluster buffer from xfs_ifree_cluster, all the inodes in cache are marked XFS_ISTALE. Those that are clean have nothing else done to them and so eventually get cleaned up by background reclaim. i.e. it is assumed we'll never dirty/relog an inode marked XFS_ISTALE. On journal commit dirty stale inodes as are handled by both buffer and inode log items to run though xfs_istale_done() and removed from the AIL (buffer log item commit) or the log item will simply unpin it because the buffer log item will clean it. What happens to any specific inode is entirely dependent on which log item wins the commit race, but the result is the same - stale inodes are clean, not attached to the cluster buffer, and not in the AIL. Hence inode reclaim can just free these inodes without further care. However, if the stale inode is relogged, it gets dirtied again and relogged into the CIL. Most of the time this isn't an issue, because relogging simply changes the inode's location in the current checkpoint. Problems arise, however, when the CIL checkpoints between two transactions in the xfs_inactive_ifree() deferops processing. This results in the XFS_ISTALE inode being redirtied and inserted into the CIL without any of the other stale cluster buffer infrastructure being in place. Hence on journal commit, it simply gets unpinned, so it remains dirty in memory. Everything in inode writeback avoids XFS_ISTALE inodes so it can't be written back, and it is not tracked in the AIL so there's not even a trigger to attempt to clean the inode. Hence the inode just sits dirty in memory until inode reclaim comes along, sees that it is XFS_ISTALE, and goes to reclaim it. This reclaiming of a dirty inode caused use after free, list corruptions and other nasty issues later in this patchset. Hence this patch addresses a violation of the "never log XFS_ISTALE inodes" caused by the deferops processing rolling a transaction and relogging a stale inode in xfs_inactive_free. It also adds a bunch of asserts to catch this problem in debug kernels so that we don't reintroduce this problem in future. Reproducer for this issue was generic/558 on a v4 filesystem. Signed-off-by:
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Yafang Shao authored
Remove current_pid(), current_test_flags() and current_clear_flags_nested(), because they are useless. Signed-off-by:
Yafang Shao <laoar.shao@gmail.com> Reviewed-by:
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Dave Chinner authored
The page faultround path ->map_pages is implemented in XFS via filemap_map_pages(). This function checks that pages found in page cache lookups have not raced with truncate based invalidation by checking page->mapping is correct and page->index is within EOF. However, we've known for a long time that this is not sufficient to protect against races with invalidations done by operations that do not change EOF. e.g. hole punching and other fallocate() based direct extent manipulations. The way we protect against these races is we wrap the page fault operations in a XFS_MMAPLOCK_SHARED lock so they serialise against fallocate and truncate before calling into the filemap function that processes the fault. Do the same for XFS's ->map_pages implementation to close this potential data corruption issue. Signed-off-by:
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Darrick J. Wong authored
Move the double-inode locking helpers to xfs_inode.c since they're not specific to reflink. Signed-off-by:
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Darrick J. Wong authored
Refactor the two functions that we use to lock and unlock two inodes to block userspace from initiating IO against a file, whether via system calls or mmap activity. Signed-off-by:
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Darrick J. Wong authored
Fix the return value of xfs_reflink_remap_prep so that its return value conventions match the rest of xfs. Signed-off-by:
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Darrick J. Wong authored
If the source and destination map are identical, we can skip the remap step to save some time. Signed-off-by:
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Darrick J. Wong authored
When logging quota block count updates during a reflink operation, we only log the /delta/ of the block count changes to the dquot. Since we now know ahead of time the extent type of both dmap and smap (and that they have the same length), we know that we only need to reserve quota blocks for dmap's blockcount if we're mapping it into a hole. Signed-off-by:
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Darrick J. Wong authored
Now that we've reworked xfs_reflink_remap_extent to remap only one extent per transaction, we actually know if the extent being removed is an allocated mapping. This means that we now know ahead of time if we're going to be touching the data fork. Since we only need blocks for a bmbt split if we're going to update the data fork, we only need to get quota reservation if we know we're going to touch the data fork. Signed-off-by:
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Darrick J. Wong authored
The existing reflink remapping loop has some structural problems that need addressing: The biggest problem is that we create one transaction for each extent in the source file without accounting for the number of mappings there are for the same range in the destination file. In other words, we don't know the number of remap operations that will be necessary and we therefore cannot guess the block reservation required. On highly fragmented filesystems (e.g. ones with active dedupe) we guess wrong, run out of block reservation, and fail. The second problem is that we don't actually use the bmap intents to their full potential -- instead of calling bunmapi directly and having to deal with its backwards operation, we could call the deferred ops xfs_bmap_unmap_extent and xfs_refcount_decrease_extent instead. This makes the frontend loop much simpler. Solve all of these problems by refactoring the remapping loops so that we only perform one remapping operation per transaction, and each operation only tries to remap a single extent from source to dest. Signed-off-by:
Edwin Török <edwin@etorok.net> Tested-by:
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Darrick J. Wong authored
The name of this predicate is a little misleading -- it decides if the extent mapping is allocated and written. Change the name to be more direct, as we're going to add a new predicate in the next patch. Signed-off-by:
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Darrick J. Wong authored
Quota reservations are supposed to account for the blocks that might be allocated due to a bmap btree split. Reflink doesn't do this, so fix this to make the quota accounting more accurate before we start rearranging things. Fixes: 862bb360 ("xfs: reflink extents from one file to another") Signed-off-by:
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Darrick J. Wong authored
The data fork scrubber calls filemap_write_and_wait to flush dirty pages and delalloc reservations out to disk prior to checking the data fork's extent mappings. Unfortunately, this means that scrub can consume the EIO/ENOSPC errors that would otherwise have stayed around in the address space until (we hope) the writer application calls fsync to persist data and collect errors. The end result is that programs that wrote to a file might never see the error code and proceed as if nothing were wrong. xfs_scrub is not in a position to notify file writers about the writeback failure, and it's only here to check metadata, not file contents. Therefore, if writeback fails, we should stuff the error code back into the address space so that an fsync by the writer application can pick that up. Fixes: 99d9d8d0 ("xfs: scrub inode block mappings") Signed-off-by:
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Brian Foster authored
The rmapbt extent swap algorithm remaps individual extents between the source inode and the target to trigger reverse mapping metadata updates. If either inode straddles a format or other bmap allocation boundary, the individual unmap and map cycles can trigger repeated bmap block allocations and frees as the extent count bounces back and forth across the boundary. While net block usage is bound across the swap operation, this behavior can prematurely exhaust the transaction block reservation because it continuously drains as the transaction rolls. Each allocation accounts against the reservation and each free returns to global free space on transaction roll. The previous workaround to this problem attempted to detect this boundary condition and provide surplus block reservation to acommodate it. This is insufficient because more remaps can occur than implied by the extent counts; if start offset boundaries are not aligned between the two inodes, for example. To address this problem more generically and dynamically, add a transaction accounting mode that returns freed blocks to the transaction reservation instead of the superblock counters on transaction roll and use it when the rmapbt based algorithm is active. This allows the chain of remap transactions to preserve the block reservation based own its own frees and prevent premature exhaustion regardless of the remap pattern. Note that this is only safe for superblocks with lazy sb accounting, but the latter is required for v5 supers and the rmap feature depends on v5. Fixes: b3fed434 ("xfs: account format bouncing into rmapbt swapext tx reservation") Root-caused-by:
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Keyur Patel authored
./xfs/libxfs/xfs_inode_buf.c:56: unnecssary ==> unnecessary ./xfs/libxfs/xfs_inode_buf.c:59: behavour ==> behaviour ./xfs/libxfs/xfs_inode_buf.c:206: unitialized ==> uninitialized Signed-off-by:
Keyur Patel <iamkeyur96@gmail.com> Reviewed-by:
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- 05 Jul, 2020 6 commits
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Linus Torvalds authored
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Linus Torvalds authored
Using a mutex for "print this warning only once" is so overdesigned as to be actively offensive to my sensitive stomach. Just use "pr_info_once()" that already does this, although in a (harmlessly) racy manner that can in theory cause the message to be printed twice if more than one CPU races on that "is this the first time" test. [ If somebody really cares about that harmless data race (which sounds very unlikely indeed), that person can trivially fix printk_once() by using a simple atomic access, preferably with an optimistic non-atomic test first before even bothering to treat the pointless "make sure it is _really_ just once" case. A mutex is most definitely never the right primitive to use for something like this. ] Yes, this is a small and meaningless detail in a code path that hardly matters. But let's keep some code quality standards here, and not accept outrageously bad code. Link: https://lore.kernel.org/lkml/CAHk-=wgV9toS7GU3KmNpj8hCS9SeF+A0voHS8F275_mgLhL4Lw@mail.gmail.com/ Cc: Andy Lutomirski <luto@kernel.org> Cc: Thomas Gleixner <tglx@linutronix.de> Cc: Peter Zijlstra (Intel) <peterz@infradead.org> Signed-off-by:
Linus Torvalds <torvalds@linux-foundation.org>
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git://git.kernel.org/pub/scm/linux/kernel/git/tip/tipLinus Torvalds authored
Pull x86 fixes from Thomas Gleixner: "A series of fixes for x86: - Reset MXCSR in kernel_fpu_begin() to prevent using a stale user space value. - Prevent writing MSR_TEST_CTRL on CPUs which are not explicitly whitelisted for split lock detection. Some CPUs which do not support it crash even when the MSR is written to 0 which is the default value. - Fix the XEN PV fallout of the entry code rework - Fix the 32bit fallout of the entry code rework - Add more selftests to ensure that these entry problems don't come back. - Disable 16 bit segments on XEN PV. It's not supported because XEN PV does not implement ESPFIX64" * tag 'x86-urgent-2020-07-05' of git://git.kernel.org/pub/scm/linux/kernel/git/tip/tip: x86/ldt: Disable 16-bit segments on Xen PV x86/entry/32: Fix #MC and #DB wiring on x86_32 x86/entry/xen: Route #DB correctly on Xen PV x86/entry, selftests: Further improve user entry sanity checks x86/entry/compat: Clear RAX high bits on Xen PV SYSENTER selftests/x86: Consolidate and fix get/set_eflags() helpers selftests/x86/syscall_nt: Clear weird flags after each test selftests/x86/syscall_nt: Add more flag combinations x86/entry/64/compat: Fix Xen PV SYSENTER frame setup x86/entry: Move SYSENTER's regs->sp and regs->flags fixups into C x86/entry: Assert that syscalls are on the right stack x86/split_lock: Don't write MSR_TEST_CTRL on CPUs that aren't whitelisted x86/fpu: Reset MXCSR to default in kernel_fpu_begin() -
git://git.kernel.org/pub/scm/linux/kernel/git/tip/tipLinus Torvalds authored
Pull irq fixes from Thomas Gleixner: "A set of interrupt chip driver fixes: - Ensure the atomicity of affinity updates in the GIC driver - Don't try to sleep in atomic context when waiting for the GICv4.1 to respond. Use polling instead. - Typo fixes in Kconfig and warnings" * tag 'irq-urgent-2020-07-05' of git://git.kernel.org/pub/scm/linux/kernel/git/tip/tip: irqchip/gic: Atomically update affinity irqchip/riscv-intc: Fix a typo in a pr_warn() irqchip/gic-v4.1: Use readx_poll_timeout_atomic() to fix sleep in atomic irqchip/loongson-pci-msi: Fix a typo in Kconfig -
git://git.kernel.org/pub/scm/linux/kernel/git/tip/tipLinus Torvalds authored
Pull rcu fixlet from Thomas Gleixner: "A single fix for a printk format warning in RCU" * tag 'core-urgent-2020-07-05' of git://git.kernel.org/pub/scm/linux/kernel/git/tip/tip: rcuperf: Fix printk format warning
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Linus Torvalds authored
Merge tag 'kbuild-fixes-v5.8-2' of git://git.kernel.org/pub/scm/linux/kernel/git/masahiroy/linux-kbuild Pull Kbuild fixes frin Masahiro Yamada: - fix various bugs in xconfig - fix some issues in cross-compilation using Clang - fix documentation * tag 'kbuild-fixes-v5.8-2' of git://git.kernel.org/pub/scm/linux/kernel/git/masahiroy/linux-kbuild: .gitignore: Do not track `defconfig` from `make savedefconfig` kbuild: make Clang build userprogs for target architecture kbuild: fix CONFIG_CC_CAN_LINK(_STATIC) for cross-compilation with Clang kconfig: qconf: parse newer types at debug info kconfig: qconf: navigate menus on hyperlinks kconfig: qconf: don't show goback button on splitMode kconfig: qconf: simplify the goBack() logic kconfig: qconf: re-implement setSelected() kconfig: qconf: make debug links work again kconfig: qconf: make search fully work again on split mode kconfig: qconf: cleanup includes docs: kbuild: fix ReST formatting gcc-plugins: fix gcc-plugins directory path in documentation
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