- 03 2月, 2010 1 次提交
-
-
由 anfei zhou 提交于
The cache alias problem will happen if the changes of user shared mapping is not flushed before copying, then user and kernel mapping may be mapped into two different cache line, it is impossible to guarantee the coherence after iov_iter_copy_from_user_atomic. So the right steps should be: flush_dcache_page(page); kmap_atomic(page); write to page; kunmap_atomic(page); flush_dcache_page(page); More precisely, we might create two new APIs flush_dcache_user_page and flush_dcache_kern_page to replace the two flush_dcache_page accordingly. Here is a snippet tested on omap2430 with VIPT cache, and I think it is not ARM-specific: int val = 0x11111111; fd = open("abc", O_RDWR); addr = mmap(NULL, 4096, PROT_READ|PROT_WRITE, MAP_SHARED, fd, 0); *(addr+0) = 0x44444444; tmp = *(addr+0); *(addr+1) = 0x77777777; write(fd, &val, sizeof(int)); close(fd); The results are not always 0x11111111 0x77777777 at the beginning as expected. Sometimes we see 0x44444444 0x77777777. Signed-off-by: NAnfei <anfei.zhou@gmail.com> Cc: Russell King <rmk@arm.linux.org.uk> Cc: Miklos Szeredi <miklos@szeredi.hu> Cc: Nick Piggin <nickpiggin@yahoo.com.au> Cc: <linux-arch@vger.kernel.org> Cc: <stable@kernel.org> Signed-off-by: NAndrew Morton <akpm@linux-foundation.org> Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
-
- 28 1月, 2010 1 次提交
-
-
由 Linus Torvalds 提交于
It's a simplified 'read_cache_page()' which takes a page allocation flag, so that different paths can control how aggressive the memory allocations are that populate a address space. In particular, the intel GPU object mapping code wants to be able to do a certain amount of own internal memory management by automatically shrinking the address space when memory starts getting tight. This allows it to dynamically use different memory allocation policies on a per-allocation basis, rather than depend on the (static) address space gfp policy. The actual new function is a one-liner, but re-organizing the helper functions to the point where you can do this with a single line of code is what most of the patch is all about. Tested-by: NChris Wilson <chris@chris-wilson.co.uk> Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
-
- 17 12月, 2009 1 次提交
-
-
由 Christoph Hellwig 提交于
In the case of direct I/O falling back to buffered I/O we sync data twice currently: once at the end of generic_file_buffered_write using filemap_write_and_wait_range and once a little later in __generic_file_aio_write using do_sync_mapping_range with all flags set. The wait before write of the do_sync_mapping_range call does not make any sense, so just keep the filemap_write_and_wait_range call and move it to the right spot. Signed-off-by: NChristoph Hellwig <hch@lst.de> Signed-off-by: NAl Viro <viro@zeniv.linux.org.uk>
-
- 10 12月, 2009 1 次提交
-
-
由 Christoph Hellwig 提交于
All callers really want the more logical filemap_fdatawait_range interface, so convert them to use it and merge wait_on_page_writeback_range into filemap_fdatawait_range. Signed-off-by: NChristoph Hellwig <hch@lst.de> Signed-off-by: NJan Kara <jack@suse.cz>
-
- 04 12月, 2009 1 次提交
-
-
由 André Goddard Rosa 提交于
That is "success", "unknown", "through", "performance", "[re|un]mapping" , "access", "default", "reasonable", "[con]currently", "temperature" , "channel", "[un]used", "application", "example","hierarchy", "therefore" , "[over|under]flow", "contiguous", "threshold", "enough" and others. Signed-off-by: NAndré Goddard Rosa <andre.goddard@gmail.com> Signed-off-by: NJiri Kosina <jkosina@suse.cz>
-
- 28 9月, 2009 1 次提交
-
-
由 Alexey Dobriyan 提交于
* mark struct vm_area_struct::vm_ops as const * mark vm_ops in AGP code But leave TTM code alone, something is fishy there with global vm_ops being used. Signed-off-by: NAlexey Dobriyan <adobriyan@gmail.com> Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
-
- 24 9月, 2009 1 次提交
-
-
由 npiggin@suse.de 提交于
Introduce new truncate helpers truncate_pagecache and inode_newsize_ok. vmtruncate is also consolidated from mm/memory.c and mm/nommu.c and into mm/truncate.c. Reviewed-by: NChristoph Hellwig <hch@lst.de> Signed-off-by: NNick Piggin <npiggin@suse.de> Signed-off-by: NAl Viro <viro@zeniv.linux.org.uk>
-
- 22 9月, 2009 1 次提交
-
-
由 KOSAKI Motohiro 提交于
Recently we encountered OOM problems due to memory use of the GEM cache. Generally a large amuont of Shmem/Tmpfs pages tend to create a memory shortage problem. We often use the following calculation to determine the amount of shmem pages: shmem = NR_ACTIVE_ANON + NR_INACTIVE_ANON - NR_ANON_PAGES however the expression does not consider isolated and mlocked pages. This patch adds explicit accounting for pages used by shmem and tmpfs. Signed-off-by: NKOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Acked-by: NRik van Riel <riel@redhat.com> Reviewed-by: NChristoph Lameter <cl@linux-foundation.org> Acked-by: NWu Fengguang <fengguang.wu@intel.com> Cc: David Rientjes <rientjes@google.com> Cc: Hugh Dickins <hugh.dickins@tiscali.co.uk> Signed-off-by: NAndrew Morton <akpm@linux-foundation.org> Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
-
- 16 9月, 2009 1 次提交
-
-
由 Andi Kleen 提交于
Add the high level memory handler that poisons pages that got corrupted by hardware (typically by a two bit flip in a DIMM or a cache) on the Linux level. The goal is to prevent everyone from accessing these pages in the future. This done at the VM level by marking a page hwpoisoned and doing the appropriate action based on the type of page it is. The code that does this is portable and lives in mm/memory-failure.c To quote the overview comment: High level machine check handler. Handles pages reported by the hardware as being corrupted usually due to a 2bit ECC memory or cache failure. This focuses on pages detected as corrupted in the background. When the current CPU tries to consume corruption the currently running process can just be killed directly instead. This implies that if the error cannot be handled for some reason it's safe to just ignore it because no corruption has been consumed yet. Instead when that happens another machine check will happen. Handles page cache pages in various states. The tricky part here is that we can access any page asynchronous to other VM users, because memory failures could happen anytime and anywhere, possibly violating some of their assumptions. This is why this code has to be extremely careful. Generally it tries to use normal locking rules, as in get the standard locks, even if that means the error handling takes potentially a long time. Some of the operations here are somewhat inefficient and have non linear algorithmic complexity, because the data structures have not been optimized for this case. This is in particular the case for the mapping from a vma to a process. Since this case is expected to be rare we hope we can get away with this. There are in principle two strategies to kill processes on poison: - just unmap the data and wait for an actual reference before killing - kill as soon as corruption is detected. Both have advantages and disadvantages and should be used in different situations. Right now both are implemented and can be switched with a new sysctl vm.memory_failure_early_kill The default is early kill. The patch does some rmap data structure walking on its own to collect processes to kill. This is unusual because normally all rmap data structure knowledge is in rmap.c only. I put it here for now to keep everything together and rmap knowledge has been seeping out anyways Includes contributions from Johannes Weiner, Chris Mason, Fengguang Wu, Nick Piggin (who did a lot of great work) and others. Cc: npiggin@suse.de Cc: riel@redhat.com Signed-off-by: NAndi Kleen <ak@linux.intel.com> Acked-by: NRik van Riel <riel@redhat.com> Reviewed-by: NHidehiro Kawai <hidehiro.kawai.ez@hitachi.com>
-
- 14 9月, 2009 6 次提交
-
-
由 Jan Kara 提交于
Remove these three functions since nobody uses them anymore. Signed-off-by: NJan Kara <jack@suse.cz>
-
由 Jan Kara 提交于
Introduce new function for generic inode syncing (vfs_fsync_range) and use it from fsync() path. Introduce also new helper for syncing after a sync write (generic_write_sync) using the generic function. Use these new helpers for syncing from generic VFS functions. This makes O_SYNC writes to block devices acquire i_mutex for syncing. If we really care about this, we can make block_fsync() drop the i_mutex and reacquire it before it returns. CC: Evgeniy Polyakov <zbr@ioremap.net> CC: ocfs2-devel@oss.oracle.com CC: Joel Becker <joel.becker@oracle.com> CC: Felix Blyakher <felixb@sgi.com> CC: xfs@oss.sgi.com CC: Anton Altaparmakov <aia21@cantab.net> CC: linux-ntfs-dev@lists.sourceforge.net CC: OGAWA Hirofumi <hirofumi@mail.parknet.co.jp> CC: linux-ext4@vger.kernel.org CC: tytso@mit.edu Acked-by: NChristoph Hellwig <hch@lst.de> Signed-off-by: NJan Kara <jack@suse.cz>
-
由 Christoph Hellwig 提交于
generic_file_aio_write_nolock() is now used only by block devices and raw character device. Filesystems should use __generic_file_aio_write() in case generic_file_aio_write() doesn't suit them. So rename the function to blkdev_aio_write() and move it to fs/blockdev.c. Signed-off-by: NChristoph Hellwig <hch@lst.de> Signed-off-by: NJan Kara <jack@suse.cz>
-
由 Jan Kara 提交于
generic_file_direct_write() and generic_file_buffered_write() called generic_osync_inode() if it was called on O_SYNC file or IS_SYNC inode. But this is superfluous since generic_file_aio_write() does the syncing as well. Also XFS and OCFS2 which call these functions directly handle syncing themselves. So let's have a single place where syncing happens: generic_file_aio_write(). We slightly change the behavior by syncing only the range of file to which the write happened for buffered writes but that should be all that is required. CC: ocfs2-devel@oss.oracle.com CC: Joel Becker <joel.becker@oracle.com> CC: Felix Blyakher <felixb@sgi.com> CC: xfs@oss.sgi.com Signed-off-by: NJan Kara <jack@suse.cz>
-
由 Jan Kara 提交于
Rename __generic_file_aio_write_nolock() to __generic_file_aio_write(), add comments to write helpers explaining how they should be used and export __generic_file_aio_write() since it will be used by some filesystems. CC: ocfs2-devel@oss.oracle.com CC: Joel Becker <joel.becker@oracle.com> Acked-by: NEvgeniy Polyakov <zbr@ioremap.net> Reviewed-by: NChristoph Hellwig <hch@lst.de> Signed-off-by: NJan Kara <jack@suse.cz>
-
由 Jan Kara 提交于
This simple helper saves some filesystems conversion from byte offset to page numbers and also makes the fdata* interface more complete. Reviewed-by: NChristoph Hellwig <hch@lst.de> Signed-off-by: NJan Kara <jack@suse.cz>
-
- 07 7月, 2009 1 次提交
-
-
由 Josef Bacik 提交于
In testing a backport of the write_begin/write_end AOPs, a 10% re-read regression was noticed when running iozone. This regression was introduced because the old AOPs would always do a mark_page_accessed(page) after the commit_write, but when the new AOPs where introduced, the only place this was kept was in pagecache_write_end(). This patch does the same thing in the generic case as what is done in pagecache_write_end(), which is just to mark the page accessed before we do write_end(). Signed-off-by: NJosef Bacik <jbacik@redhat.com> Acked-by: NNick Piggin <npiggin@suse.de> Signed-off-by: NAndrew Morton <akpm@linux-foundation.org> Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
-
- 17 6月, 2009 8 次提交
-
-
由 Mel Gorman 提交于
Callers of alloc_pages_node() can optionally specify -1 as a node to mean "allocate from the current node". However, a number of the callers in fast paths know for a fact their node is valid. To avoid a comparison and branch, this patch adds alloc_pages_exact_node() that only checks the nid with VM_BUG_ON(). Callers that know their node is valid are then converted. Signed-off-by: NMel Gorman <mel@csn.ul.ie> Reviewed-by: NChristoph Lameter <cl@linux-foundation.org> Reviewed-by: NKOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Reviewed-by: NPekka Enberg <penberg@cs.helsinki.fi> Acked-by: Paul Mundt <lethal@linux-sh.org> [for the SLOB NUMA bits] Cc: Peter Zijlstra <a.p.zijlstra@chello.nl> Cc: Nick Piggin <nickpiggin@yahoo.com.au> Cc: Dave Hansen <dave@linux.vnet.ibm.com> Cc: Lee Schermerhorn <Lee.Schermerhorn@hp.com> Signed-off-by: NAndrew Morton <akpm@linux-foundation.org> Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
-
由 Wu Fengguang 提交于
Now that we do readahead for sequential mmap reads, here is a simple evaluation of the impacts, and one further optimization. It's an NFS-root debian desktop system, readahead size = 60 pages. The numbers are grabbed after a fresh boot into console. approach pgmajfault RA miss ratio mmap IO count avg IO size(pages) A 383 31.6% 383 11 B 225 32.4% 390 11 C 224 32.6% 307 13 case A: mmap sync/async readahead disabled case B: mmap sync/async readahead enabled, with enforced full async readahead size case C: mmap sync/async readahead enabled, with enforced full sync/async readahead size or: A = vanilla 2.6.30-rc1 B = A plus mmap readahead C = B plus this patch The numbers show that - there are good possibilities for random mmap reads to trigger readahead - 'pgmajfault' is reduced by 1/3, due to the _async_ nature of readahead - case C can further reduce IO count by 1/4 - readahead miss ratios are not quite affected The theory is - readahead is _good_ for clustered random reads, and can perform _better_ than readaround because they could be _async_. - async readahead size is guaranteed to be larger than readaround size, and they are _async_, hence will mostly behave better However for B - sync readahead size could be smaller than readaround size, hence may make things worse by produce more smaller IOs which will be fixed by this patch. Final conclusion: - mmap readahead reduced major faults by 1/3 and no obvious overheads; - mmap io can be further reduced by 1/4 with this patch. Signed-off-by: NWu Fengguang <fengguang.wu@intel.com> Signed-off-by: NAndrew Morton <akpm@linux-foundation.org> Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
-
由 Wu Fengguang 提交于
Signed-off-by: NWu Fengguang <fengguang.wu@intel.com> Signed-off-by: NAndrew Morton <akpm@linux-foundation.org> Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
-
由 Wu Fengguang 提交于
Mmap read-around now shares the same code style and data structure with readahead code. This also removes do_page_cache_readahead(). Its last user, mmap read-around, has been changed to call ra_submit(). The no-readahead-if-congested logic is dumped by the way. Users will be pretty sensitive about the slow loading of executables. So it's unfavorable to disabled mmap read-around on a congested queue. [akpm@linux-foundation.org: coding-style fixes] Cc: Nick Piggin <npiggin@suse.de> Signed-off-by: NFengguang Wu <wfg@mail.ustc.edu.cn> Cc: Ying Han <yinghan@google.com> Signed-off-by: NAndrew Morton <akpm@linux-foundation.org> Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
-
由 Wu Fengguang 提交于
We need this in one particular case and two more general ones. Now we do async readahead for sequential mmap reads, and do it with the help of PG_readahead. For normal reads, PG_readahead is the sufficient condition to do a sequential readahead. But unfortunately, for mmap reads, there is a tiny nuisance: [11736.998347] readahead-init0(process: sh/23926, file: sda1/w3m, offset=0:4503599627370495, ra=0+4-3) = 4 [11737.014985] readahead-around(process: w3m/23926, file: sda1/w3m, offset=0:0, ra=290+32-0) = 17 [11737.019488] readahead-around(process: w3m/23926, file: sda1/w3m, offset=0:0, ra=118+32-0) = 32 [11737.024921] readahead-interleaved(process: w3m/23926, file: sda1/w3m, offset=0:2, ra=4+6-6) = 6 ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ ~~~~~~~~~~~~~ An unfavorably small readahead. The original dumb read-around size could be more efficient. That happened because ld-linux.so does a read(832) in L1 before mmap(), which triggers a 4-page readahead, with the second page tagged PG_readahead. L0: open("/lib/libc.so.6", O_RDONLY) = 3 L1: read(3, "\177ELF\2\1\1\0\0\0\0\0\0\0\0\0\3\0>\0\1\0\0\0\340\342"..., 832) = 832 ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ L2: fstat(3, {st_mode=S_IFREG|0755, st_size=1420624, ...}) = 0 L3: mmap(NULL, 3527256, PROT_READ|PROT_EXEC, MAP_PRIVATE|MAP_DENYWRITE, 3, 0) = 0x7fac6e51d000 L4: mprotect(0x7fac6e671000, 2097152, PROT_NONE) = 0 L5: mmap(0x7fac6e871000, 20480, PROT_READ|PROT_WRITE, MAP_PRIVATE|MAP_FIXED|MAP_DENYWRITE, 3, 0x154000) = 0x7fac6e871000 L6: mmap(0x7fac6e876000, 16984, PROT_READ|PROT_WRITE, MAP_PRIVATE|MAP_FIXED|MAP_ANONYMOUS, -1, 0) = 0x7fac6e876000 L7: close(3) = 0 In general, the PG_readahead flag will also be hit in cases - sequential reads - clustered random reads A full readahead size is desirable in both cases. Cc: Nick Piggin <npiggin@suse.de> Signed-off-by: NWu Fengguang <fengguang.wu@intel.com> Cc: Ying Han <yinghan@google.com> Signed-off-by: NAndrew Morton <akpm@linux-foundation.org> Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
-
由 Wu Fengguang 提交于
Auto-detect sequential mmap reads and do readahead for them. The sequential mmap readahead will be triggered when - sync readahead: it's a major fault and (prev_offset == offset-1); - async readahead: minor fault on PG_readahead page with valid readahead state. The benefits of doing readahead instead of read-around: - less I/O wait thanks to async readahead - double real I/O size and no more cache hits The single stream case is improved a little. For 100,000 sequential mmap reads: user system cpu total (1-1) plain -mm, 128KB readaround: 3.224 2.554 48.40% 11.838 (1-2) plain -mm, 256KB readaround: 3.170 2.392 46.20% 11.976 (2) patched -mm, 128KB readahead: 3.117 2.448 47.33% 11.607 The patched (2) has smallest total time, since it has no cache hit overheads and less I/O block time(thanks to async readahead). Here the I/O size makes no much difference, since there's only one single stream. Note that (1-1)'s real I/O size is 64KB and (1-2)'s real I/O size is 128KB, since the half of the read-around pages will be readahead cache hits. This is going to make _real_ differences for _concurrent_ IO streams. Cc: Nick Piggin <npiggin@suse.de> Signed-off-by: NWu Fengguang <fengguang.wu@intel.com> Cc: Ying Han <yinghan@google.com> Signed-off-by: NAndrew Morton <akpm@linux-foundation.org> Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
-
由 Linus Torvalds 提交于
This shouldn't really change behavior all that much, but the single rather complex function with read-ahead inside a loop etc is broken up into more manageable pieces. The behaviour is also less subtle, with the read-ahead being done up-front rather than inside some subtle loop and thus avoiding the now unnecessary extra state variables (ie "did_readaround" is gone). Fengguang: the code split in fact fixed a bug reported by Pavel Levshin: the PGMAJFAULT accounting used to be bypassed when MADV_RANDOM is set, in which case the original code will directly jump to no_cached_page reading. Cc: Pavel Levshin <lpk@581.spb.su> Cc: <wli@movementarian.org> Cc: Nick Piggin <npiggin@suse.de> Signed-off-by: NWu Fengguang <fengguang.wu@intel.com> Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org> Signed-off-by: NAndrew Morton <akpm@linux-foundation.org> Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
-
由 Wu Fengguang 提交于
Impact: code simplification. Cc: Nick Piggin <npiggin@suse.de> Signed-off-by: NWu Fengguang <fengguang.wu@intel.com> Cc: Ying Han <yinghan@google.com> Signed-off-by: NAndrew Morton <akpm@linux-foundation.org> Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
-
- 29 5月, 2009 1 次提交
-
-
由 Daisuke Nishimura 提交于
mapping->tree_lock can be acquired from interrupt context. Then, following dead lock can occur. Assume "A" as a page. CPU0: lock_page_cgroup(A) interrupted -> take mapping->tree_lock. CPU1: take mapping->tree_lock -> lock_page_cgroup(A) This patch tries to fix above deadlock by moving memcg's hook to out of mapping->tree_lock. charge/uncharge of pagecache/swapcache is protected by page lock, not tree_lock. After this patch, lock_page_cgroup() is not called under mapping->tree_lock. Signed-off-by: NKAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Signed-off-by: NDaisuke Nishimura <nishimura@mxp.nes.nec.co.jp> Cc: Balbir Singh <balbir@in.ibm.com> Cc: Daisuke Nishimura <nishimura@mxp.nes.nec.co.jp> Signed-off-by: NAndrew Morton <akpm@linux-foundation.org> Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
-
- 16 4月, 2009 1 次提交
-
-
由 Chris Mason 提交于
This wasn't exported before and is useful (used by the experimental ext3 data=guarded code) Signed-off-by: NChris Mason <chris.mason@oracle.com> Acked-by: NTheodore Tso <tytso@mit.edu> Acked-by: NJan Kara <jack@suse.cz> Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
-
- 14 4月, 2009 1 次提交
-
-
由 Randy Dunlap 提交于
Fix filemap.c kernel-doc warnings: Warning(mm/filemap.c:575): No description found for parameter 'page' Warning(mm/filemap.c:575): No description found for parameter 'waiter' Signed-off-by: NRandy Dunlap <randy.dunlap@oracle.com> Signed-off-by: NAndrew Morton <akpm@linux-foundation.org> Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
-
- 04 4月, 2009 1 次提交
-
-
由 Evgeniy Polyakov 提交于
This patch adds Kconfig and Makefile entries and exports to VFS functions to be used by POHMELFS. Signed-off-by: NEvgeniy Polyakov <zbr@ioremap.net> Signed-off-by: NGreg Kroah-Hartman <gregkh@suse.de>
-
- 03 4月, 2009 2 次提交
-
-
由 David Howells 提交于
Add a function to install a monitor on the page lock waitqueue for a particular page, thus allowing the page being unlocked to be detected. This is used by CacheFiles to detect read completion on a page in the backing filesystem so that it can then copy the data to the waiting netfs page. Signed-off-by: NDavid Howells <dhowells@redhat.com> Acked-by: NSteve Dickson <steved@redhat.com> Acked-by: NTrond Myklebust <Trond.Myklebust@netapp.com> Acked-by: NRik van Riel <riel@redhat.com> Acked-by: NAl Viro <viro@zeniv.linux.org.uk> Tested-by: NDaire Byrne <Daire.Byrne@framestore.com>
-
由 David Howells 提交于
Recruit a page flag to aid in cache management. The following extra flag is defined: (1) PG_fscache (PG_private_2) The marked page is backed by a local cache and is pinning resources in the cache driver. If PG_fscache is set, then things that checked for PG_private will now also check for that. This includes things like truncation and page invalidation. The function page_has_private() had been added to make the checks for both PG_private and PG_private_2 at the same time. Signed-off-by: NDavid Howells <dhowells@redhat.com> Acked-by: NSteve Dickson <steved@redhat.com> Acked-by: NTrond Myklebust <Trond.Myklebust@netapp.com> Acked-by: NRik van Riel <riel@redhat.com> Acked-by: NAl Viro <viro@zeniv.linux.org.uk> Tested-by: NDaire Byrne <Daire.Byrne@framestore.com>
-
- 02 3月, 2009 1 次提交
-
-
由 Ingo Molnar 提交于
Impact: standardize IO on cached ops On modern CPUs it is almost always a bad idea to use non-temporal stores, as the regression in this commit has shown it: 30d697fa: x86: fix performance regression in write() syscall The kernel simply has no good information about whether using non-temporal stores is a good idea or not - and trying to add heuristics only increases complexity and inserts fragility. The regression on cached write()s took very long to be found - over two years. So dont take any chances and let the hardware decide how it makes use of its caches. The only exception is drivers/gpu/drm/i915/i915_gem.c: there were we are absolutely sure that another entity (the GPU) will pick up the dirty data immediately and that the CPU will not touch that data before the GPU will. Also, keep the _nocache() primitives to make it easier for people to experiment with these details. There may be more clear-cut cases where non-cached copies can be used, outside of filemap.c. Cc: Salman Qazi <sqazi@google.com> Cc: Nick Piggin <npiggin@suse.de> Cc: Linus Torvalds <torvalds@linux-foundation.org> Signed-off-by: NIngo Molnar <mingo@elte.hu>
-
- 25 2月, 2009 1 次提交
-
-
由 Ingo Molnar 提交于
Impact: cleanup, enable future change Add a 'total bytes copied' parameter to __copy_from_user_*nocache(), and update all the callsites. The parameter is not used yet - architecture code can use it to more intelligently decide whether the copy should be cached or non-temporal. Cc: Salman Qazi <sqazi@google.com> Cc: Nick Piggin <npiggin@suse.de> Cc: Linus Torvalds <torvalds@linux-foundation.org> Signed-off-by: NIngo Molnar <mingo@elte.hu>
-
- 14 1月, 2009 2 次提交
-
-
由 Heiko Carstens 提交于
System calls with an unsigned long long argument can't be converted with the standard wrappers since that would include a cast to long, which in turn means that we would lose the upper 32 bit on 32 bit architectures. Also semctl can't use the standard wrapper since it has a 'union' parameter. So we handle them as special case and add some extra wrappers instead. Signed-off-by: NHeiko Carstens <heiko.carstens@de.ibm.com>
-
由 Heiko Carstens 提交于
Convert all system calls to return a long. This should be a NOP since all converted types should have the same size anyway. With the exception of sys_exit_group which returned void. But that doesn't matter since the system call doesn't return. Signed-off-by: NHeiko Carstens <heiko.carstens@de.ibm.com>
-
- 09 1月, 2009 1 次提交
-
-
由 KAMEZAWA Hiroyuki 提交于
My patch, memcg-fix-gfp_mask-of-callers-of-charge.patch changed gfp_mask of callers of charge to be GFP_HIGHUSER_MOVABLE for showing what will happen at memory reclaim. But in recent discussion, it's NACKed because it sounds ugly. This patch is for reverting it and add some clean up to gfp_mask of callers of charge. No behavior change but need review before generating HUNK in deep queue. This patch also adds explanation to meaning of gfp_mask passed to charge functions in memcontrol.h. Signed-off-by: NKAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: Balbir Singh <balbir@in.ibm.com> Cc: Daisuke Nishimura <nishimura@mxp.nes.nec.co.jp> Cc: Hugh Dickins <hugh@veritas.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Signed-off-by: NAndrew Morton <akpm@linux-foundation.org> Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
-
- 07 1月, 2009 4 次提交
-
-
由 Nick Piggin 提交于
Frustratingly, gfp_t is really divided into two classes of flags. One are the context dependent ones (can we sleep? can we enter filesystem? block subsystem? should we use some extra reserves, etc.). The other ones are the type of memory required and depend on how the algorithm is implemented rather than the point at which the memory is allocated (highmem? dma memory? etc). Some of the functions which allocate a page and add it to page cache take a gfp_t, but sometimes those functions or their callers aren't really doing the right thing: when allocating pagecache page, the memory type should be mapping_gfp_mask(mapping). When allocating radix tree nodes, the memory type should be kernel mapped (not highmem) memory. The gfp_t argument should only really be needed for context dependent options. This patch doesn't really solve that tangle in a nice way, but it does attempt to fix a couple of bugs. - find_or_create_page changes its radix-tree allocation to only include the main context dependent flags in order so the pagecache page may be allocated from arbitrary types of memory without affecting the radix-tree. In practice, slab allocations don't come from highmem anyway, and radix-tree only uses slab allocations. So there isn't a practical change (unless some fs uses GFP_DMA for pages). - grab_cache_page_nowait() is changed to allocate radix-tree nodes with GFP_NOFS, because it is not supposed to reenter the filesystem. This bug could cause lock recursion if a filesystem is not expecting the function to reenter the fs (as-per documentation). Filesystems should be careful about exactly what semantics they want and what they get when fiddling with gfp_t masks to allocate pagecache. One should be as liberal as possible with the type of memory that can be used, and same for the the context specific flags. Signed-off-by: NNick Piggin <npiggin@suse.de> Signed-off-by: NAndrew Morton <akpm@linux-foundation.org> Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
-
由 Nick Piggin 提交于
Direct IO can invalidate and sync a lot of pagecache pages in the mapping. A 4K direct IO will actually try to sync and/or invalidate the pagecache of the entire file, for example (which might be many GB or TB large). Improve this by doing range syncs. Also, memory no longer has to be unmapped to catch the dirty bits for syncing, as dirty bits would remain coherent due to dirty mmap accounting. This fixes the immediate DM deadlocks when doing direct IO reads to block device with a mounted filesystem, if only by papering over the problem somewhat rather than addressing the fsync starvation cases. Signed-off-by: NNick Piggin <npiggin@suse.de> Reviewed-by: NJeff Moyer <jmoyer@redhat.com> Signed-off-by: NAndrew Morton <akpm@linux-foundation.org> Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
-
由 Nick Piggin 提交于
In write_cache_pages, nr_to_write is heeded even for data-integrity syncs, so the function will return success after writing out nr_to_write pages, even if that was not sufficient to guarantee data integrity. The callers tend to set it to values that could break data interity semantics easily in practice. For example, nr_to_write can be set to mapping->nr_pages * 2, however if a file has a single, dirty page, then fsync is called, subsequent pages might be concurrently added and dirtied, then write_cache_pages might writeout two of these newly dirty pages, while not writing out the old page that should have been written out. Fix this by ignoring nr_to_write if it is a data integrity sync. This is a data integrity bug. The reason this has been done in the past is to avoid stalling sync operations behind page dirtiers. "If a file has one dirty page at offset 1000000000000000 then someone does an fsync() and someone else gets in first and starts madly writing pages at offset 0, we want to write that page at 1000000000000000. Somehow." What we do today is return success after an arbitrary amount of pages are written, whether or not we have provided the data-integrity semantics that the caller has asked for. Even this doesn't actually fix all stall cases completely: in the above situation, if the file has a huge number of pages in pagecache (but not dirty), then mapping->nrpages is going to be huge, even if pages are being dirtied. This change does indeed make the possibility of long stalls lager, and that's not a good thing, but lying about data integrity is even worse. We have to either perform the sync, or return -ELINUXISLAME so at least the caller knows what has happened. There are subsequent competing approaches in the works to solve the stall problems properly, without compromising data integrity. Signed-off-by: NNick Piggin <npiggin@suse.de> Cc: Chris Mason <chris.mason@oracle.com> Cc: Dave Chinner <david@fromorbit.com> Signed-off-by: NAndrew Morton <akpm@linux-foundation.org> Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
-
由 Nick Piggin 提交于
Doing a mark_page_accessed at fault-time, then doing SetPageReferenced at unmap-time if the pte is young has a number of problems. mark_page_accessed is supposed to be roughly the equivalent of a young pte for unmapped references. Unfortunately it doesn't come with any context: after being called, reclaim doesn't know who or why the page was touched. So calling mark_page_accessed not only adds extra lru or PG_referenced manipulations for pages that are already going to have pte_young ptes anyway, but it also adds these references which are difficult to work with from the context of vma specific references (eg. MADV_SEQUENTIAL pte_young may not wish to contribute to the page being referenced). Then, simply doing SetPageReferenced when zapping a pte and finding it is young, is not a really good solution either. SetPageReferenced does not correctly promote the page to the active list for example. So after removing mark_page_accessed from the fault path, several mmap()+touch+munmap() would have a very different result from several read(2) calls for example, which is not really desirable. Signed-off-by: NNick Piggin <npiggin@suse.de> Acked-by: NJohannes Weiner <hannes@saeurebad.de> Signed-off-by: NAndrew Morton <akpm@linux-foundation.org> Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
-
- 06 1月, 2009 1 次提交
-
-
由 Dmitri Monakhov 提交于
We don't have to do it because it is useless for non regular files. In fact block device may trigger this path without dentry->d_inode->i_mutex. (akpm: concerns were expressed (by me) about S_ISDIR inodes) Signed-off-by: NDmitri Monakhov <dmonakhov@openvz.org> Signed-off-by: NAndrew Morton <akpm@linux-foundation.org> Signed-off-by: NAl Viro <viro@zeniv.linux.org.uk>
-