In struct page we have enough space to fit long-size page->ptl there,
but we use dynamically-allocated page->ptl if size(spinlock_t) is larger
than sizeof(int).
It hurts 64-bit architectures with CONFIG_GENERIC_LOCKBREAK, where
sizeof(spinlock_t) == 8, but it easily fits into struct page.
Signed-off-by: Kirill A. Shutemov <kirill.shutemov@linux.intel.com>
Acked-by: Hugh Dickins <hughd@google.com>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
According to documentation on barriers, stores issued before a LOCK can
complete after the lock implying that it's possible tlb_flush_pending
can be visible after a page table update. As per revised documentation,
this patch adds a smp_mb__before_spinlock to guarantee the correct
ordering.
Signed-off-by: Mel Gorman <mgorman@suse.de>
Acked-by: Paul E. McKenney <paulmck@linux.vnet.ibm.com>
Reviewed-by: Rik van Riel <riel@redhat.com>
Cc: <stable@vger.kernel.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
There are a few subtle races, between change_protection_range (used by
mprotect and change_prot_numa) on one side, and NUMA page migration and
compaction on the other side.
The basic race is that there is a time window between when the PTE gets
made non-present (PROT_NONE or NUMA), and the TLB is flushed.
During that time, a CPU may continue writing to the page.
This is fine most of the time, however compaction or the NUMA migration
code may come in, and migrate the page away.
When that happens, the CPU may continue writing, through the cached
translation, to what is no longer the current memory location of the
process.
This only affects x86, which has a somewhat optimistic pte_accessible.
All other architectures appear to be safe, and will either always flush,
or flush whenever there is a valid mapping, even with no permissions
(SPARC).
The basic race looks like this:
CPU A CPU B CPU C
load TLB entry
make entry PTE/PMD_NUMA
fault on entry
read/write old page
start migrating page
change PTE/PMD to new page
read/write old page [*]
flush TLB
reload TLB from new entry
read/write new page
lose data
[*] the old page may belong to a new user at this point!
The obvious fix is to flush remote TLB entries, by making sure that
pte_accessible aware of the fact that PROT_NONE and PROT_NUMA memory may
still be accessible if there is a TLB flush pending for the mm.
This should fix both NUMA migration and compaction.
[mgorman@suse.de: fix build]
Signed-off-by: Rik van Riel <riel@redhat.com>
Signed-off-by: Mel Gorman <mgorman@suse.de>
Cc: Alex Thorlton <athorlton@sgi.com>
Cc: <stable@vger.kernel.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Pull SLAB changes from Pekka Enberg:
"The patches from Joonsoo Kim switch mm/slab.c to use 'struct page' for
slab internals similar to mm/slub.c. This reduces memory usage and
improves performance:
https://lkml.org/lkml/2013/10/16/155
Rest of the changes are bug fixes from various people"
* 'slab/next' of git://git.kernel.org/pub/scm/linux/kernel/git/penberg/linux: (21 commits)
mm, slub: fix the typo in mm/slub.c
mm, slub: fix the typo in include/linux/slub_def.h
slub: Handle NULL parameter in kmem_cache_flags
slab: replace non-existing 'struct freelist *' with 'void *'
slab: fix to calm down kmemleak warning
slub: proper kmemleak tracking if CONFIG_SLUB_DEBUG disabled
slab: rename slab_bufctl to slab_freelist
slab: remove useless statement for checking pfmemalloc
slab: use struct page for slab management
slab: replace free and inuse in struct slab with newly introduced active
slab: remove SLAB_LIMIT
slab: remove kmem_bufctl_t
slab: change the management method of free objects of the slab
slab: use __GFP_COMP flag for allocating slab pages
slab: use well-defined macro, virt_to_slab()
slab: overloading the RCU head over the LRU for RCU free
slab: remove cachep in struct slab_rcu
slab: remove nodeid in struct slab
slab: remove colouroff in struct slab
slab: change return type of kmem_getpages() to struct page
...
I don't know what went wrong, mis-merge or something, but ->pmd_huge_pte
placed in wrong union within struct page.
In original patch[1] it's placed to union with ->lru and ->slab, but in
commit e009bb30c8 ("mm: implement split page table lock for PMD
level") it's in union with ->index and ->freelist.
That union seems also unused for pages with table tables and safe to
re-use, but it's not what I've tested.
Let's move it to original place. It fixes indentation at least. :)
[1] https://lkml.org/lkml/2013/10/7/288
Signed-off-by: Kirill A. Shutemov <kirill.shutemov@linux.intel.com>
Reviewed-by: Naoya Horiguchi <n-horiguchi@ah.jp.nec.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Use kernel/bounds.c to convert build-time spinlock_t size check into a
preprocessor symbol and apply that to properly separate the page::ptl
situation.
Signed-off-by: Peter Zijlstra <peterz@infradead.org>
Signed-off-by: Kirill A. Shutemov <kirill.shutemov@linux.intel.com>
Cc: Ingo Molnar <mingo@elte.hu>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
If split page table lock is in use, we embed the lock into struct page
of table's page. We have to disable split lock, if spinlock_t is too
big be to be embedded, like when DEBUG_SPINLOCK or DEBUG_LOCK_ALLOC
enabled.
This patch add support for dynamic allocation of split page table lock
if we can't embed it to struct page.
page->ptl is unsigned long now and we use it as spinlock_t if
sizeof(spinlock_t) <= sizeof(long), otherwise it's pointer to spinlock_t.
The spinlock_t allocated in pgtable_page_ctor() for PTE table and in
pgtable_pmd_page_ctor() for PMD table. All other helpers converted to
support dynamically allocated page->ptl.
Signed-off-by: Kirill A. Shutemov <kirill.shutemov@linux.intel.com>
Reviewed-by: Peter Zijlstra <peterz@infradead.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
The basic idea is the same as with PTE level: the lock is embedded into
struct page of table's page.
We can't use mm->pmd_huge_pte to store pgtables for THP, since we don't
take mm->page_table_lock anymore. Let's reuse page->lru of table's page
for that.
pgtable_pmd_page_ctor() returns true, if initialization is successful
and false otherwise. Current implementation never fails, but assumption
that constructor can fail will help to port it to -rt where spinlock_t
is rather huge and cannot be embedded into struct page -- dynamic
allocation is required.
Signed-off-by: Naoya Horiguchi <n-horiguchi@ah.jp.nec.com>
Signed-off-by: Kirill A. Shutemov <kirill.shutemov@linux.intel.com>
Tested-by: Alex Thorlton <athorlton@sgi.com>
Cc: Ingo Molnar <mingo@redhat.com>
Cc: "Eric W . Biederman" <ebiederm@xmission.com>
Cc: "Paul E . McKenney" <paulmck@linux.vnet.ibm.com>
Cc: Al Viro <viro@zeniv.linux.org.uk>
Cc: Andi Kleen <ak@linux.intel.com>
Cc: Andrea Arcangeli <aarcange@redhat.com>
Cc: Dave Hansen <dave.hansen@intel.com>
Cc: Dave Jones <davej@redhat.com>
Cc: David Howells <dhowells@redhat.com>
Cc: Frederic Weisbecker <fweisbec@gmail.com>
Cc: Johannes Weiner <hannes@cmpxchg.org>
Cc: Kees Cook <keescook@chromium.org>
Cc: Mel Gorman <mgorman@suse.de>
Cc: Michael Kerrisk <mtk.manpages@gmail.com>
Cc: Oleg Nesterov <oleg@redhat.com>
Cc: Peter Zijlstra <peterz@infradead.org>
Cc: Rik van Riel <riel@redhat.com>
Cc: Robin Holt <robinmholt@gmail.com>
Cc: Sedat Dilek <sedat.dilek@gmail.com>
Cc: Srikar Dronamraju <srikar@linux.vnet.ibm.com>
Cc: Thomas Gleixner <tglx@linutronix.de>
Cc: Hugh Dickins <hughd@google.com>
Reviewed-by: Steven Rostedt <rostedt@goodmis.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Now, there are a few field in struct slab, so we can overload these
over struct page. This will save some memory and reduce cache footprint.
After this change, slabp_cache and slab_size no longer related to
a struct slab, so rename them as freelist_cache and freelist_size.
These changes are just mechanical ones and there is no functional change.
Acked-by: Andi Kleen <ak@linux.intel.com>
Acked-by: Christoph Lameter <cl@linux.com>
Signed-off-by: Joonsoo Kim <iamjoonsoo.kim@lge.com>
Signed-off-by: Pekka Enberg <penberg@iki.fi>
With build-time size checking, we can overload the RCU head over the LRU
of struct page to free pages of a slab in rcu context. This really help to
implement to overload the struct slab over the struct page and this
eventually reduce memory usage and cache footprint of the SLAB.
Acked-by: Andi Kleen <ak@linux.intel.com>
Acked-by: Christoph Lameter <cl@linux.com>
Signed-off-by: Joonsoo Kim <iamjoonsoo.kim@lge.com>
Signed-off-by: Pekka Enberg <penberg@iki.fi>
With scan rate adaptions based on whether the workload has properly
converged or not there should be no need for the scan period reset
hammer. Get rid of it.
Signed-off-by: Mel Gorman <mgorman@suse.de>
Reviewed-by: Rik van Riel <riel@redhat.com>
Cc: Andrea Arcangeli <aarcange@redhat.com>
Cc: Johannes Weiner <hannes@cmpxchg.org>
Cc: Srikar Dronamraju <srikar@linux.vnet.ibm.com>
Signed-off-by: Peter Zijlstra <peterz@infradead.org>
Link: http://lkml.kernel.org/r/1381141781-10992-60-git-send-email-mgorman@suse.de
Signed-off-by: Ingo Molnar <mingo@kernel.org>
Change the per page last fault tracking to use cpu,pid instead of
nid,pid. This will allow us to try and lookup the alternate task more
easily. Note that even though it is the cpu that is store in the page
flags that the mpol_misplaced decision is still based on the node.
Signed-off-by: Peter Zijlstra <peterz@infradead.org>
Signed-off-by: Mel Gorman <mgorman@suse.de>
Reviewed-by: Rik van Riel <riel@redhat.com>
Cc: Andrea Arcangeli <aarcange@redhat.com>
Cc: Johannes Weiner <hannes@cmpxchg.org>
Cc: Srikar Dronamraju <srikar@linux.vnet.ibm.com>
Link: http://lkml.kernel.org/r/1381141781-10992-43-git-send-email-mgorman@suse.de
[ Fixed build failure on 32-bit systems. ]
Signed-off-by: Ingo Molnar <mingo@kernel.org>
Ideally it would be possible to distinguish between NUMA hinting faults that
are private to a task and those that are shared. If treated identically
there is a risk that shared pages bounce between nodes depending on
the order they are referenced by tasks. Ultimately what is desirable is
that task private pages remain local to the task while shared pages are
interleaved between sharing tasks running on different nodes to give good
average performance. This is further complicated by THP as even
applications that partition their data may not be partitioning on a huge
page boundary.
To start with, this patch assumes that multi-threaded or multi-process
applications partition their data and that in general the private accesses
are more important for cpu->memory locality in the general case. Also,
no new infrastructure is required to treat private pages properly but
interleaving for shared pages requires additional infrastructure.
To detect private accesses the pid of the last accessing task is required
but the storage requirements are a high. This patch borrows heavily from
Ingo Molnar's patch "numa, mm, sched: Implement last-CPU+PID hash tracking"
to encode some bits from the last accessing task in the page flags as
well as the node information. Collisions will occur but it is better than
just depending on the node information. Node information is then used to
determine if a page needs to migrate. The PID information is used to detect
private/shared accesses. The preferred NUMA node is selected based on where
the maximum number of approximately private faults were measured. Shared
faults are not taken into consideration for a few reasons.
First, if there are many tasks sharing the page then they'll all move
towards the same node. The node will be compute overloaded and then
scheduled away later only to bounce back again. Alternatively the shared
tasks would just bounce around nodes because the fault information is
effectively noise. Either way accounting for shared faults the same as
private faults can result in lower performance overall.
The second reason is based on a hypothetical workload that has a small
number of very important, heavily accessed private pages but a large shared
array. The shared array would dominate the number of faults and be selected
as a preferred node even though it's the wrong decision.
The third reason is that multiple threads in a process will race each
other to fault the shared page making the fault information unreliable.
Signed-off-by: Mel Gorman <mgorman@suse.de>
[ Fix complication error when !NUMA_BALANCING. ]
Reviewed-by: Rik van Riel <riel@redhat.com>
Cc: Andrea Arcangeli <aarcange@redhat.com>
Cc: Johannes Weiner <hannes@cmpxchg.org>
Cc: Srikar Dronamraju <srikar@linux.vnet.ibm.com>
Signed-off-by: Peter Zijlstra <peterz@infradead.org>
Link: http://lkml.kernel.org/r/1381141781-10992-30-git-send-email-mgorman@suse.de
Signed-off-by: Ingo Molnar <mingo@kernel.org>
PTE scanning and NUMA hinting fault handling is expensive so commit
5bca2303 ("mm: sched: numa: Delay PTE scanning until a task is scheduled
on a new node") deferred the PTE scan until a task had been scheduled on
another node. The problem is that in the purely shared memory case that
this may never happen and no NUMA hinting fault information will be
captured. We are not ruling out the possibility that something better
can be done here but for now, this patch needs to be reverted and depend
entirely on the scan_delay to avoid punishing short-lived processes.
Signed-off-by: Mel Gorman <mgorman@suse.de>
Reviewed-by: Rik van Riel <riel@redhat.com>
Cc: Andrea Arcangeli <aarcange@redhat.com>
Cc: Johannes Weiner <hannes@cmpxchg.org>
Cc: Srikar Dronamraju <srikar@linux.vnet.ibm.com>
Signed-off-by: Peter Zijlstra <peterz@infradead.org>
Link: http://lkml.kernel.org/r/1381141781-10992-16-git-send-email-mgorman@suse.de
Signed-off-by: Ingo Molnar <mingo@kernel.org>
Pull aio changes from Ben LaHaise:
"First off, sorry for this pull request being late in the merge window.
Al had raised a couple of concerns about 2 items in the series below.
I addressed the first issue (the race introduced by Gu's use of
mm_populate()), but he has not provided any further details on how he
wants to rework the anon_inode.c changes (which were sent out months
ago but have yet to be commented on).
The bulk of the changes have been sitting in the -next tree for a few
months, with all the issues raised being addressed"
* git://git.kvack.org/~bcrl/aio-next: (22 commits)
aio: rcu_read_lock protection for new rcu_dereference calls
aio: fix race in ring buffer page lookup introduced by page migration support
aio: fix rcu sparse warnings introduced by ioctx table lookup patch
aio: remove unnecessary debugging from aio_free_ring()
aio: table lookup: verify ctx pointer
staging/lustre: kiocb->ki_left is removed
aio: fix error handling and rcu usage in "convert the ioctx list to table lookup v3"
aio: be defensive to ensure request batching is non-zero instead of BUG_ON()
aio: convert the ioctx list to table lookup v3
aio: double aio_max_nr in calculations
aio: Kill ki_dtor
aio: Kill ki_users
aio: Kill unneeded kiocb members
aio: Kill aio_rw_vect_retry()
aio: Don't use ctx->tail unnecessarily
aio: io_cancel() no longer returns the io_event
aio: percpu ioctx refcount
aio: percpu reqs_available
aio: reqs_active -> reqs_available
aio: fix build when migration is disabled
...
This is the updated version of df54d6fa54 ("x86 get_unmapped_area():
use proper mmap base for bottom-up direction") that only randomizes the
mmap base address once.
Signed-off-by: Radu Caragea <sinaelgl@gmail.com>
Reported-and-tested-by: Jeff Shorey <shoreyjeff@gmail.com>
Cc: Andrew Morton <akpm@linux-foundation.org>
Cc: Michel Lespinasse <walken@google.com>
Cc: Oleg Nesterov <oleg@redhat.com>
Cc: Rik van Riel <riel@redhat.com>
Cc: Ingo Molnar <mingo@elte.hu>
Cc: Adrian Sendroiu <molecula2788@gmail.com>
Cc: Greg KH <greg@kroah.com>
Cc: Kamal Mostafa <kamal@canonical.com>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
On Wed, Jun 12, 2013 at 11:14:40AM -0700, Kent Overstreet wrote:
> On Mon, Apr 15, 2013 at 02:40:55PM +0300, Octavian Purdila wrote:
> > When using a large number of threads performing AIO operations the
> > IOCTX list may get a significant number of entries which will cause
> > significant overhead. For example, when running this fio script:
> >
> > rw=randrw; size=256k ;directory=/mnt/fio; ioengine=libaio; iodepth=1
> > blocksize=1024; numjobs=512; thread; loops=100
> >
> > on an EXT2 filesystem mounted on top of a ramdisk we can observe up to
> > 30% CPU time spent by lookup_ioctx:
> >
> > 32.51% [guest.kernel] [g] lookup_ioctx
> > 9.19% [guest.kernel] [g] __lock_acquire.isra.28
> > 4.40% [guest.kernel] [g] lock_release
> > 4.19% [guest.kernel] [g] sched_clock_local
> > 3.86% [guest.kernel] [g] local_clock
> > 3.68% [guest.kernel] [g] native_sched_clock
> > 3.08% [guest.kernel] [g] sched_clock_cpu
> > 2.64% [guest.kernel] [g] lock_release_holdtime.part.11
> > 2.60% [guest.kernel] [g] memcpy
> > 2.33% [guest.kernel] [g] lock_acquired
> > 2.25% [guest.kernel] [g] lock_acquire
> > 1.84% [guest.kernel] [g] do_io_submit
> >
> > This patchs converts the ioctx list to a radix tree. For a performance
> > comparison the above FIO script was run on a 2 sockets 8 core
> > machine. This are the results (average and %rsd of 10 runs) for the
> > original list based implementation and for the radix tree based
> > implementation:
> >
> > cores 1 2 4 8 16 32
> > list 109376 ms 69119 ms 35682 ms 22671 ms 19724 ms 16408 ms
> > %rsd 0.69% 1.15% 1.17% 1.21% 1.71% 1.43%
> > radix 73651 ms 41748 ms 23028 ms 16766 ms 15232 ms 13787 ms
> > %rsd 1.19% 0.98% 0.69% 1.13% 0.72% 0.75%
> > % of radix
> > relative 66.12% 65.59% 66.63% 72.31% 77.26% 83.66%
> > to list
> >
> > To consider the impact of the patch on the typical case of having
> > only one ctx per process the following FIO script was run:
> >
> > rw=randrw; size=100m ;directory=/mnt/fio; ioengine=libaio; iodepth=1
> > blocksize=1024; numjobs=1; thread; loops=100
> >
> > on the same system and the results are the following:
> >
> > list 58892 ms
> > %rsd 0.91%
> > radix 59404 ms
> > %rsd 0.81%
> > % of radix
> > relative 100.87%
> > to list
>
> So, I was just doing some benchmarking/profiling to get ready to send
> out the aio patches I've got for 3.11 - and it looks like your patch is
> causing a ~1.5% throughput regression in my testing :/
... <snip>
I've got an alternate approach for fixing this wart in lookup_ioctx()...
Instead of using an rbtree, just use the reserved id in the ring buffer
header to index an array pointing the ioctx. It's not finished yet, and
it needs to be tidied up, but is most of the way there.
-ben
--
"Thought is the essence of where you are now."
--
kmo> And, a rework of Ben's code, but this was entirely his idea
kmo> -Kent
bcrl> And fix the code to use the right mm_struct in kill_ioctx(), actually
free memory.
Signed-off-by: Benjamin LaHaise <bcrl@kvack.org>
Since all architectures have been converted to use vm_unmapped_area(),
there is no remaining use for the free_area_cache.
Signed-off-by: Michel Lespinasse <walken@google.com>
Acked-by: Rik van Riel <riel@redhat.com>
Cc: "James E.J. Bottomley" <jejb@parisc-linux.org>
Cc: "Luck, Tony" <tony.luck@intel.com>
Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org>
Cc: David Howells <dhowells@redhat.com>
Cc: Helge Deller <deller@gmx.de>
Cc: Ivan Kokshaysky <ink@jurassic.park.msu.ru>
Cc: Matt Turner <mattst88@gmail.com>
Cc: Paul Mackerras <paulus@samba.org>
Cc: Richard Henderson <rth@twiddle.net>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
page->_last_nid fits into page->flags on 64-bit. The unlikely 32-bit
NUMA configuration with NUMA Balancing will still need an extra page
field. As Peter notes "Completely dropping 32bit support for
CONFIG_NUMA_BALANCING would simplify things, but it would also remove
the warning if we grow enough 64bit only page-flags to push the last-cpu
out."
[mgorman@suse.de: minor modifications]
Signed-off-by: Mel Gorman <mgorman@suse.de>
Cc: Peter Zijlstra <a.p.zijlstra@chello.nl>
Cc: Andrea Arcangeli <aarcange@redhat.com>
Cc: Ingo Molnar <mingo@kernel.org>
Cc: Simon Jeons <simon.jeons@gmail.com>
Cc: Wanpeng Li <liwanp@linux.vnet.ibm.com>
Cc: Hugh Dickins <hughd@google.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
This is a preparation patch for moving page->_last_nid into page->flags
that moves page flag layout information to a separate header. This
patch is necessary because otherwise there would be a circular
dependency between mm_types.h and mm.h.
Signed-off-by: Mel Gorman <mgorman@suse.de>
Cc: Peter Zijlstra <a.p.zijlstra@chello.nl>
Cc: Andrea Arcangeli <aarcange@redhat.com>
Cc: Ingo Molnar <mingo@kernel.org>
Cc: Simon Jeons <simon.jeons@gmail.com>
Cc: Wanpeng Li <liwanp@linux.vnet.ibm.com>
Cc: Hugh Dickins <hughd@google.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
s/me/be/ and clarify the comment a bit when we're changing it anyway.
Signed-off-by: Mel Gorman <mgorman@suse.de>
Suggested-by: Simon Jeons <simon.jeons@gmail.com>
Cc: Peter Zijlstra <a.p.zijlstra@chello.nl>
Cc: Andrea Arcangeli <aarcange@redhat.com>
Cc: Ingo Molnar <mingo@kernel.org>
Cc: Wanpeng Li <liwanp@linux.vnet.ibm.com>
Cc: Hugh Dickins <hughd@google.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Pull SLAB changes from Pekka Enberg:
"This contains preparational work from Christoph Lameter and Glauber
Costa for SLAB memcg and cleanups and improvements from Ezequiel
Garcia and Joonsoo Kim.
Please note that the SLOB cleanup commit from Arnd Bergmann already
appears in your tree but I had also merged it myself which is why it
shows up in the shortlog."
* 'slab/for-linus' of git://git.kernel.org/pub/scm/linux/kernel/git/penberg/linux:
mm/sl[aou]b: Common alignment code
slab: Use the new create_boot_cache function to simplify bootstrap
slub: Use statically allocated kmem_cache boot structure for bootstrap
mm, sl[au]b: create common functions for boot slab creation
slab: Simplify bootstrap
slub: Use correct cpu_slab on dead cpu
mm: fix slab.c kernel-doc warnings
mm/slob: use min_t() to compare ARCH_SLAB_MINALIGN
slab: Ignore internal flags in cache creation
mm/slob: Use free_page instead of put_page for page-size kmalloc allocations
mm/sl[aou]b: Move common kmem_cache_size() to slab.h
mm/slob: Use object_size field in kmem_cache_size()
mm/slob: Drop usage of page->private for storing page-sized allocations
slub: Commonize slab_cache field in struct page
sl[au]b: Process slabinfo_show in common code
mm/sl[au]b: Move print_slabinfo_header to slab_common.c
mm/sl[au]b: Move slabinfo processing to slab_common.c
slub: remove one code path and reduce lock contention in __slab_free()
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Merge tag 'balancenuma-v11' of git://git.kernel.org/pub/scm/linux/kernel/git/mel/linux-balancenuma
Pull Automatic NUMA Balancing bare-bones from Mel Gorman:
"There are three implementations for NUMA balancing, this tree
(balancenuma), numacore which has been developed in tip/master and
autonuma which is in aa.git.
In almost all respects balancenuma is the dumbest of the three because
its main impact is on the VM side with no attempt to be smart about
scheduling. In the interest of getting the ball rolling, it would be
desirable to see this much merged for 3.8 with the view to building
scheduler smarts on top and adapting the VM where required for 3.9.
The most recent set of comparisons available from different people are
mel: https://lkml.org/lkml/2012/12/9/108
mingo: https://lkml.org/lkml/2012/12/7/331
tglx: https://lkml.org/lkml/2012/12/10/437
srikar: https://lkml.org/lkml/2012/12/10/397
The results are a mixed bag. In my own tests, balancenuma does
reasonably well. It's dumb as rocks and does not regress against
mainline. On the other hand, Ingo's tests shows that balancenuma is
incapable of converging for this workloads driven by perf which is bad
but is potentially explained by the lack of scheduler smarts. Thomas'
results show balancenuma improves on mainline but falls far short of
numacore or autonuma. Srikar's results indicate we all suffer on a
large machine with imbalanced node sizes.
My own testing showed that recent numacore results have improved
dramatically, particularly in the last week but not universally.
We've butted heads heavily on system CPU usage and high levels of
migration even when it shows that overall performance is better.
There are also cases where it regresses. Of interest is that for
specjbb in some configurations it will regress for lower numbers of
warehouses and show gains for higher numbers which is not reported by
the tool by default and sometimes missed in treports. Recently I
reported for numacore that the JVM was crashing with
NullPointerExceptions but currently it's unclear what the source of
this problem is. Initially I thought it was in how numacore batch
handles PTEs but I'm no longer think this is the case. It's possible
numacore is just able to trigger it due to higher rates of migration.
These reports were quite late in the cycle so I/we would like to start
with this tree as it contains much of the code we can agree on and has
not changed significantly over the last 2-3 weeks."
* tag 'balancenuma-v11' of git://git.kernel.org/pub/scm/linux/kernel/git/mel/linux-balancenuma: (50 commits)
mm/rmap, migration: Make rmap_walk_anon() and try_to_unmap_anon() more scalable
mm/rmap: Convert the struct anon_vma::mutex to an rwsem
mm: migrate: Account a transhuge page properly when rate limiting
mm: numa: Account for failed allocations and isolations as migration failures
mm: numa: Add THP migration for the NUMA working set scanning fault case build fix
mm: numa: Add THP migration for the NUMA working set scanning fault case.
mm: sched: numa: Delay PTE scanning until a task is scheduled on a new node
mm: sched: numa: Control enabling and disabling of NUMA balancing if !SCHED_DEBUG
mm: sched: numa: Control enabling and disabling of NUMA balancing
mm: sched: Adapt the scanning rate if a NUMA hinting fault does not migrate
mm: numa: Use a two-stage filter to restrict pages being migrated for unlikely task<->node relationships
mm: numa: migrate: Set last_nid on newly allocated page
mm: numa: split_huge_page: Transfer last_nid on tail page
mm: numa: Introduce last_nid to the page frame
sched: numa: Slowly increase the scanning period as NUMA faults are handled
mm: numa: Rate limit setting of pte_numa if node is saturated
mm: numa: Rate limit the amount of memory that is migrated between nodes
mm: numa: Structures for Migrate On Fault per NUMA migration rate limiting
mm: numa: Migrate pages handled during a pmd_numa hinting fault
mm: numa: Migrate on reference policy
...
The kernel walks the VMA rbtree in various places, including the page
fault path. However, the vm_rb node spanned two cache lines, on 64 bit
systems with 64 byte cache lines (most x86 systems).
Rearrange vm_area_struct a little, so all the information we need to do a
VMA tree walk is in the first cache line.
[akpm@linux-foundation.org: checkpatch fixes]
Signed-off-by: Michel Lespinasse <walken@google.com>
Signed-off-by: Rik van Riel <riel@redhat.com>
Cc: Hugh Dickins <hughd@google.com>
Cc: Russell King <linux@arm.linux.org.uk>
Cc: Ralf Baechle <ralf@linux-mips.org>
Cc: Paul Mundt <lethal@linux-sh.org>
Cc: "David S. Miller" <davem@davemloft.net>
Cc: Chris Metcalf <cmetcalf@tilera.com>
Cc: Ingo Molnar <mingo@elte.hu>
Cc: Thomas Gleixner <tglx@linutronix.de>
Cc: "H. Peter Anvin" <hpa@zytor.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Define vma->rb_subtree_gap as the largest gap between any vma in the
subtree rooted at that vma, and their predecessor. Or, for a recursive
definition, vma->rb_subtree_gap is the max of:
- vma->vm_start - vma->vm_prev->vm_end
- rb_subtree_gap fields of the vmas pointed by vma->rb.rb_left and
vma->rb.rb_right
This will allow get_unmapped_area_* to find a free area of the right
size in O(log(N)) time, instead of potentially having to do a linear
walk across all the VMAs.
Also define mm->highest_vm_end as the vm_end field of the highest vma,
so that we can easily check if the following gap is suitable.
This does have the potential to make unmapping VMAs more expensive,
especially for processes with very large numbers of VMAs, where the VMA
rbtree can grow quite deep.
Signed-off-by: Michel Lespinasse <walken@google.com>
Reviewed-by: Rik van Riel <riel@redhat.com>
Cc: Hugh Dickins <hughd@google.com>
Cc: Russell King <linux@arm.linux.org.uk>
Cc: Ralf Baechle <ralf@linux-mips.org>
Cc: Paul Mundt <lethal@linux-sh.org>
Cc: "David S. Miller" <davem@davemloft.net>
Cc: Chris Metcalf <cmetcalf@tilera.com>
Cc: Ingo Molnar <mingo@elte.hu>
Cc: Thomas Gleixner <tglx@linutronix.de>
Cc: "H. Peter Anvin" <hpa@zytor.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Due to the fact that migrations are driven by the CPU a task is running
on there is no point tracking NUMA faults until one task runs on a new
node. This patch tracks the first node used by an address space. Until
it changes, PTE scanning is disabled and no NUMA hinting faults are
trapped. This should help workloads that are short-lived, do not care
about NUMA placement or have bound themselves to a single node.
This takes advantage of the logic in "mm: sched: numa: Implement slow
start for working set sampling" to delay when the checks are made. This
will take advantage of processes that set their CPU and node bindings
early in their lifetime. It will also potentially allow any initial load
balancing to take place.
Signed-off-by: Mel Gorman <mgorman@suse.de>
The PTE scanning rate and fault rates are two of the biggest sources of
system CPU overhead with automatic NUMA placement. Ideally a proper policy
would detect if a workload was properly placed, schedule and adjust the
PTE scanning rate accordingly. We do not track the necessary information
to do that but we at least know if we migrated or not.
This patch scans slower if a page was not migrated as the result of a
NUMA hinting fault up to sysctl_numa_balancing_scan_period_max which is
now higher than the previous default. Once every minute it will reset
the scanner in case of phase changes.
This is hilariously crude and the numbers are arbitrary. Workloads will
converge quite slowly in comparison to what a proper policy should be able
to do. On the plus side, we will chew up less CPU for workloads that have
no need for automatic balancing.
Signed-off-by: Mel Gorman <mgorman@suse.de>
This patch introduces a last_nid field to the page struct. This is used
to build a two-stage filter in the next patch that is aimed at
mitigating a problem whereby pages migrate to the wrong node when
referenced by a process that was running off its home node.
Signed-off-by: Mel Gorman <mgorman@suse.de>
Previously, to probe the working set of a task, we'd use
a very simple and crude method: mark all of its address
space PROT_NONE.
That method has various (obvious) disadvantages:
- it samples the working set at dissimilar rates,
giving some tasks a sampling quality advantage
over others.
- creates performance problems for tasks with very
large working sets
- over-samples processes with large address spaces but
which only very rarely execute
Improve that method by keeping a rotating offset into the
address space that marks the current position of the scan,
and advance it by a constant rate (in a CPU cycles execution
proportional manner). If the offset reaches the last mapped
address of the mm then it then it starts over at the first
address.
The per-task nature of the working set sampling functionality in this tree
allows such constant rate, per task, execution-weight proportional sampling
of the working set, with an adaptive sampling interval/frequency that
goes from once per 100ms up to just once per 8 seconds. The current
sampling volume is 256 MB per interval.
As tasks mature and converge their working set, so does the
sampling rate slow down to just a trickle, 256 MB per 8
seconds of CPU time executed.
This, beyond being adaptive, also rate-limits rarely
executing systems and does not over-sample on overloaded
systems.
[ In AutoNUMA speak, this patch deals with the effective sampling
rate of the 'hinting page fault'. AutoNUMA's scanning is
currently rate-limited, but it is also fundamentally
single-threaded, executing in the knuma_scand kernel thread,
so the limit in AutoNUMA is global and does not scale up with
the number of CPUs, nor does it scan tasks in an execution
proportional manner.
So the idea of rate-limiting the scanning was first implemented
in the AutoNUMA tree via a global rate limit. This patch goes
beyond that by implementing an execution rate proportional
working set sampling rate that is not implemented via a single
global scanning daemon. ]
[ Dan Carpenter pointed out a possible NULL pointer dereference in the
first version of this patch. ]
Based-on-idea-by: Andrea Arcangeli <aarcange@redhat.com>
Bug-Found-By: Dan Carpenter <dan.carpenter@oracle.com>
Signed-off-by: Peter Zijlstra <a.p.zijlstra@chello.nl>
Cc: Linus Torvalds <torvalds@linux-foundation.org>
Cc: Andrew Morton <akpm@linux-foundation.org>
Cc: Peter Zijlstra <a.p.zijlstra@chello.nl>
Cc: Andrea Arcangeli <aarcange@redhat.com>
Cc: Rik van Riel <riel@redhat.com>
[ Wrote changelog and fixed bug. ]
Signed-off-by: Ingo Molnar <mingo@kernel.org>
Signed-off-by: Mel Gorman <mgorman@suse.de>
Reviewed-by: Rik van Riel <riel@redhat.com>
NOTE: This patch is based on "sched, numa, mm: Add fault driven
placement and migration policy" but as it throws away all the policy
to just leave a basic foundation I had to drop the signed-offs-by.
This patch creates a bare-bones method for setting PTEs pte_numa in the
context of the scheduler that when faulted later will be faulted onto the
node the CPU is running on. In itself this does nothing useful but any
placement policy will fundamentally depend on receiving hints on placement
from fault context and doing something intelligent about it.
Signed-off-by: Mel Gorman <mgorman@suse.de>
Acked-by: Rik van Riel <riel@redhat.com>
Right now, slab and slub have fields in struct page to derive which
cache a page belongs to, but they do it slightly differently.
slab uses a field called slab_cache, that lives in the third double
word. slub, uses a field called "slab", living outside of the
doublewords area.
Ideally, we could use the same field for this. Since slub heavily makes
use of the doubleword region, there isn't really much room to move
slub's slab_cache field around. Since slab does not have such strict
placement restrictions, we can move it outside the doubleword area.
The naming used by slab, "slab_cache", is less confusing, and it is
preferred over slub's generic "slab".
Signed-off-by: Glauber Costa <glommer@parallels.com>
Acked-by: Christoph Lameter <cl@linux.com>
CC: David Rientjes <rientjes@google.com>
Signed-off-by: Pekka Enberg <penberg@kernel.org>
Implement an interval tree as a replacement for the VMA prio_tree. The
algorithms are similar to lib/interval_tree.c; however that code can't be
directly reused as the interval endpoints are not explicitly stored in the
VMA. So instead, the common algorithm is moved into a template and the
details (node type, how to get interval endpoints from the node, etc) are
filled in using the C preprocessor.
Once the interval tree functions are available, using them as a
replacement to the VMA prio tree is a relatively simple, mechanical job.
Signed-off-by: Michel Lespinasse <walken@google.com>
Cc: Rik van Riel <riel@redhat.com>
Cc: Hillf Danton <dhillf@gmail.com>
Cc: Peter Zijlstra <a.p.zijlstra@chello.nl>
Cc: Catalin Marinas <catalin.marinas@arm.com>
Cc: Andrea Arcangeli <aarcange@redhat.com>
Cc: David Woodhouse <dwmw2@infradead.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
A long time ago, in v2.4, VM_RESERVED kept swapout process off VMA,
currently it lost original meaning but still has some effects:
| effect | alternative flags
-+------------------------+---------------------------------------------
1| account as reserved_vm | VM_IO
2| skip in core dump | VM_IO, VM_DONTDUMP
3| do not merge or expand | VM_IO, VM_DONTEXPAND, VM_HUGETLB, VM_PFNMAP
4| do not mlock | VM_IO, VM_DONTEXPAND, VM_HUGETLB, VM_PFNMAP
This patch removes reserved_vm counter from mm_struct. Seems like nobody
cares about it, it does not exported into userspace directly, it only
reduces total_vm showed in proc.
Thus VM_RESERVED can be replaced with VM_IO or pair VM_DONTEXPAND | VM_DONTDUMP.
remap_pfn_range() and io_remap_pfn_range() set VM_IO|VM_DONTEXPAND|VM_DONTDUMP.
remap_vmalloc_range() set VM_DONTEXPAND | VM_DONTDUMP.
[akpm@linux-foundation.org: drivers/vfio/pci/vfio_pci.c fixup]
Signed-off-by: Konstantin Khlebnikov <khlebnikov@openvz.org>
Cc: Alexander Viro <viro@zeniv.linux.org.uk>
Cc: Carsten Otte <cotte@de.ibm.com>
Cc: Chris Metcalf <cmetcalf@tilera.com>
Cc: Cyrill Gorcunov <gorcunov@openvz.org>
Cc: Eric Paris <eparis@redhat.com>
Cc: H. Peter Anvin <hpa@zytor.com>
Cc: Hugh Dickins <hughd@google.com>
Cc: Ingo Molnar <mingo@redhat.com>
Cc: James Morris <james.l.morris@oracle.com>
Cc: Jason Baron <jbaron@redhat.com>
Cc: Kentaro Takeda <takedakn@nttdata.co.jp>
Cc: Matt Helsley <matthltc@us.ibm.com>
Cc: Nick Piggin <npiggin@kernel.dk>
Cc: Oleg Nesterov <oleg@redhat.com>
Cc: Peter Zijlstra <a.p.zijlstra@chello.nl>
Cc: Robert Richter <robert.richter@amd.com>
Cc: Suresh Siddha <suresh.b.siddha@intel.com>
Cc: Tetsuo Handa <penguin-kernel@I-love.SAKURA.ne.jp>
Cc: Venkatesh Pallipadi <venki@google.com>
Acked-by: Linus Torvalds <torvalds@linux-foundation.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Currently the kernel sets mm->exe_file during sys_execve() and then tracks
number of vmas with VM_EXECUTABLE flag in mm->num_exe_file_vmas, as soon
as this counter drops to zero kernel resets mm->exe_file to NULL. Plus it
resets mm->exe_file at last mmput() when mm->mm_users drops to zero.
VMA with VM_EXECUTABLE flag appears after mapping file with flag
MAP_EXECUTABLE, such vmas can appears only at sys_execve() or after vma
splitting, because sys_mmap ignores this flag. Usually binfmt module sets
mm->exe_file and mmaps executable vmas with this file, they hold
mm->exe_file while task is running.
comment from v2.6.25-6245-g925d1c4 ("procfs task exe symlink"),
where all this stuff was introduced:
> The kernel implements readlink of /proc/pid/exe by getting the file from
> the first executable VMA. Then the path to the file is reconstructed and
> reported as the result.
>
> Because of the VMA walk the code is slightly different on nommu systems.
> This patch avoids separate /proc/pid/exe code on nommu systems. Instead of
> walking the VMAs to find the first executable file-backed VMA we store a
> reference to the exec'd file in the mm_struct.
>
> That reference would prevent the filesystem holding the executable file
> from being unmounted even after unmapping the VMAs. So we track the number
> of VM_EXECUTABLE VMAs and drop the new reference when the last one is
> unmapped. This avoids pinning the mounted filesystem.
exe_file's vma accounting is hooked into every file mmap/unmmap and vma
split/merge just to fix some hypothetical pinning fs from umounting by mm,
which already unmapped all its executable files, but still alive.
Seems like currently nobody depends on this behaviour. We can try to
remove this logic and keep mm->exe_file until final mmput().
mm->exe_file is still protected with mm->mmap_sem, because we want to
change it via new sys_prctl(PR_SET_MM_EXE_FILE). Also via this syscall
task can change its mm->exe_file and unpin mountpoint explicitly.
Signed-off-by: Konstantin Khlebnikov <khlebnikov@openvz.org>
Cc: Alexander Viro <viro@zeniv.linux.org.uk>
Cc: Carsten Otte <cotte@de.ibm.com>
Cc: Chris Metcalf <cmetcalf@tilera.com>
Cc: Cyrill Gorcunov <gorcunov@openvz.org>
Cc: Eric Paris <eparis@redhat.com>
Cc: H. Peter Anvin <hpa@zytor.com>
Cc: Hugh Dickins <hughd@google.com>
Cc: Ingo Molnar <mingo@redhat.com>
Cc: James Morris <james.l.morris@oracle.com>
Cc: Jason Baron <jbaron@redhat.com>
Cc: Kentaro Takeda <takedakn@nttdata.co.jp>
Cc: Matt Helsley <matthltc@us.ibm.com>
Cc: Nick Piggin <npiggin@kernel.dk>
Cc: Oleg Nesterov <oleg@redhat.com>
Cc: Peter Zijlstra <a.p.zijlstra@chello.nl>
Cc: Robert Richter <robert.richter@amd.com>
Cc: Suresh Siddha <suresh.b.siddha@intel.com>
Cc: Tetsuo Handa <penguin-kernel@I-love.SAKURA.ne.jp>
Cc: Venkatesh Pallipadi <venki@google.com>
Acked-by: Linus Torvalds <torvalds@linux-foundation.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
__GFP_MEMALLOC will allow the allocation to disregard the watermarks, much
like PF_MEMALLOC. It allows one to pass along the memalloc state in
object related allocation flags as opposed to task related flags, such as
sk->sk_allocation. This removes the need for ALLOC_PFMEMALLOC as callers
using __GFP_MEMALLOC can get the ALLOC_NO_WATERMARK flag which is now
enough to identify allocations related to page reclaim.
Signed-off-by: Peter Zijlstra <a.p.zijlstra@chello.nl>
Signed-off-by: Mel Gorman <mgorman@suse.de>
Cc: David Miller <davem@davemloft.net>
Cc: Neil Brown <neilb@suse.de>
Cc: Mike Christie <michaelc@cs.wisc.edu>
Cc: Eric B Munson <emunson@mgebm.net>
Cc: Eric Dumazet <eric.dumazet@gmail.com>
Cc: Sebastian Andrzej Siewior <sebastian@breakpoint.cc>
Cc: Mel Gorman <mgorman@suse.de>
Cc: Christoph Lameter <cl@linux.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
When a user or administrator requires swap for their application, they
create a swap partition and file, format it with mkswap and activate it
with swapon. Swap over the network is considered as an option in diskless
systems. The two likely scenarios are when blade servers are used as part
of a cluster where the form factor or maintenance costs do not allow the
use of disks and thin clients.
The Linux Terminal Server Project recommends the use of the Network Block
Device (NBD) for swap according to the manual at
https://sourceforge.net/projects/ltsp/files/Docs-Admin-Guide/LTSPManual.pdf/download
There is also documentation and tutorials on how to setup swap over NBD at
places like https://help.ubuntu.com/community/UbuntuLTSP/EnableNBDSWAP The
nbd-client also documents the use of NBD as swap. Despite this, the fact
is that a machine using NBD for swap can deadlock within minutes if swap
is used intensively. This patch series addresses the problem.
The core issue is that network block devices do not use mempools like
normal block devices do. As the host cannot control where they receive
packets from, they cannot reliably work out in advance how much memory
they might need. Some years ago, Peter Zijlstra developed a series of
patches that supported swap over an NFS that at least one distribution is
carrying within their kernels. This patch series borrows very heavily
from Peter's work to support swapping over NBD as a pre-requisite to
supporting swap-over-NFS. The bulk of the complexity is concerned with
preserving memory that is allocated from the PFMEMALLOC reserves for use
by the network layer which is needed for both NBD and NFS.
Patch 1 adds knowledge of the PFMEMALLOC reserves to SLAB and SLUB to
preserve access to pages allocated under low memory situations
to callers that are freeing memory.
Patch 2 optimises the SLUB fast path to avoid pfmemalloc checks
Patch 3 introduces __GFP_MEMALLOC to allow access to the PFMEMALLOC
reserves without setting PFMEMALLOC.
Patch 4 opens the possibility for softirqs to use PFMEMALLOC reserves
for later use by network packet processing.
Patch 5 only sets page->pfmemalloc when ALLOC_NO_WATERMARKS was required
Patch 6 ignores memory policies when ALLOC_NO_WATERMARKS is set.
Patches 7-12 allows network processing to use PFMEMALLOC reserves when
the socket has been marked as being used by the VM to clean pages. If
packets are received and stored in pages that were allocated under
low-memory situations and are unrelated to the VM, the packets
are dropped.
Patch 11 reintroduces __skb_alloc_page which the networking
folk may object to but is needed in some cases to propogate
pfmemalloc from a newly allocated page to an skb. If there is a
strong objection, this patch can be dropped with the impact being
that swap-over-network will be slower in some cases but it should
not fail.
Patch 13 is a micro-optimisation to avoid a function call in the
common case.
Patch 14 tags NBD sockets as being SOCK_MEMALLOC so they can use
PFMEMALLOC if necessary.
Patch 15 notes that it is still possible for the PFMEMALLOC reserve
to be depleted. To prevent this, direct reclaimers get throttled on
a waitqueue if 50% of the PFMEMALLOC reserves are depleted. It is
expected that kswapd and the direct reclaimers already running
will clean enough pages for the low watermark to be reached and
the throttled processes are woken up.
Patch 16 adds a statistic to track how often processes get throttled
Some basic performance testing was run using kernel builds, netperf on
loopback for UDP and TCP, hackbench (pipes and sockets), iozone and
sysbench. Each of them were expected to use the sl*b allocators
reasonably heavily but there did not appear to be significant performance
variances.
For testing swap-over-NBD, a machine was booted with 2G of RAM with a
swapfile backed by NBD. 8*NUM_CPU processes were started that create
anonymous memory mappings and read them linearly in a loop. The total
size of the mappings were 4*PHYSICAL_MEMORY to use swap heavily under
memory pressure.
Without the patches and using SLUB, the machine locks up within minutes
and runs to completion with them applied. With SLAB, the story is
different as an unpatched kernel run to completion. However, the patched
kernel completed the test 45% faster.
MICRO
3.5.0-rc2 3.5.0-rc2
vanilla swapnbd
Unrecognised test vmscan-anon-mmap-write
MMTests Statistics: duration
Sys Time Running Test (seconds) 197.80 173.07
User+Sys Time Running Test (seconds) 206.96 182.03
Total Elapsed Time (seconds) 3240.70 1762.09
This patch: mm: sl[au]b: add knowledge of PFMEMALLOC reserve pages
Allocations of pages below the min watermark run a risk of the machine
hanging due to a lack of memory. To prevent this, only callers who have
PF_MEMALLOC or TIF_MEMDIE set and are not processing an interrupt are
allowed to allocate with ALLOC_NO_WATERMARKS. Once they are allocated to
a slab though, nothing prevents other callers consuming free objects
within those slabs. This patch limits access to slab pages that were
alloced from the PFMEMALLOC reserves.
When this patch is applied, pages allocated from below the low watermark
are returned with page->pfmemalloc set and it is up to the caller to
determine how the page should be protected. SLAB restricts access to any
page with page->pfmemalloc set to callers which are known to able to
access the PFMEMALLOC reserve. If one is not available, an attempt is
made to allocate a new page rather than use a reserve. SLUB is a bit more
relaxed in that it only records if the current per-CPU page was allocated
from PFMEMALLOC reserve and uses another partial slab if the caller does
not have the necessary GFP or process flags. This was found to be
sufficient in tests to avoid hangs due to SLUB generally maintaining
smaller lists than SLAB.
In low-memory conditions it does mean that !PFMEMALLOC allocators can fail
a slab allocation even though free objects are available because they are
being preserved for callers that are freeing pages.
[a.p.zijlstra@chello.nl: Original implementation]
[sebastian@breakpoint.cc: Correct order of page flag clearing]
Signed-off-by: Mel Gorman <mgorman@suse.de>
Cc: David Miller <davem@davemloft.net>
Cc: Neil Brown <neilb@suse.de>
Cc: Peter Zijlstra <a.p.zijlstra@chello.nl>
Cc: Mike Christie <michaelc@cs.wisc.edu>
Cc: Eric B Munson <emunson@mgebm.net>
Cc: Eric Dumazet <eric.dumazet@gmail.com>
Cc: Sebastian Andrzej Siewior <sebastian@breakpoint.cc>
Cc: Mel Gorman <mgorman@suse.de>
Cc: Christoph Lameter <cl@linux.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Pull SLAB changes from Pekka Enberg:
"Most of the changes included are from Christoph Lameter's "common
slab" patch series that unifies common parts of SLUB, SLAB, and SLOB
allocators. The unification is needed for Glauber Costa's "kmem
memcg" work that will hopefully appear for v3.7.
The rest of the changes are fixes and speedups by various people."
* 'slab/next' of git://git.kernel.org/pub/scm/linux/kernel/git/penberg/linux: (32 commits)
mm: Fix build warning in kmem_cache_create()
slob: Fix early boot kernel crash
mm, slub: ensure irqs are enabled for kmemcheck
mm, sl[aou]b: Move kmem_cache_create mutex handling to common code
mm, sl[aou]b: Use a common mutex definition
mm, sl[aou]b: Common definition for boot state of the slab allocators
mm, sl[aou]b: Extract common code for kmem_cache_create()
slub: remove invalid reference to list iterator variable
mm: Fix signal SIGFPE in slabinfo.c.
slab: move FULL state transition to an initcall
slab: Fix a typo in commit 8c138b "slab: Get rid of obj_size macro"
mm, slab: Build fix for recent kmem_cache changes
slab: rename gfpflags to allocflags
slub: refactoring unfreeze_partials()
slub: use __cmpxchg_double_slab() at interrupt disabled place
slab/mempolicy: always use local policy from interrupt context
slab: Get rid of obj_size macro
mm, sl[aou]b: Extract common fields from struct kmem_cache
slab: Remove some accessors
slab: Use page struct fields instead of casting
...
On arches that do not support this_cpu_cmpxchg_double() slab_lock is used
to do atomic cmpxchg() on double word which contains page->_count. The
page count can be changed from get_page() or put_page() without taking
slab_lock. That corrupts page counter.
Fix it by moving page->_count out of cmpxchg_double data. So that slub
does no change it while updating slub meta-data in struct page.
[akpm@linux-foundation.org: use standard comment layout, tweak comment text]
Reported-by: Amey Bhide <abhide@nicira.com>
Signed-off-by: Pravin B Shelar <pshelar@nicira.com>
Acked-by: Christoph Lameter <cl@linux.com>
Cc: Pekka Enberg <penberg@cs.helsinki.fi>
Cc: Andrea Arcangeli <aarcange@redhat.com>
Cc: <stable@vger.kernel.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Add fields to the page struct so that it is properly documented that
slab overlays the lru fields.
This cleans up some casts in slab.
Reviewed-by: Glauber Costa <glommer@parallels.com>
Reviewed-by: Joonsoo Kim <js1304@gmail.com>
Signed-off-by: Christoph Lameter <cl@linux.com>
Signed-off-by: Pekka Enberg <penberg@kernel.org>
Define the fields used by slob in mm_types.h and use struct page instead
of struct slob_page in slob. This cleans up numerous of typecasts in slob.c and
makes readers aware of slob's use of page struct fields.
[Also cleans up some bitrot in slob.c. The page struct field layout
in slob.c is an old layout and does not match the one in mm_types.h]
Reviewed-by: Glauber Costa <glommer@parallels.com>
Acked-by: David Rientjes <rientjes@google.com>
Reviewed-by: Joonsoo Kim <js1304@gmail.com>
Signed-off-by: Christoph Lameter <cl@linux.com>
Signed-off-by: Pekka Enberg <penberg@kernel.org>
The swap token code no longer fits in with the current VM model. It
does not play well with cgroups or the better NUMA placement code in
development, since we have only one swap token globally.
It also has the potential to mess with scalability of the system, by
increasing the number of non-reclaimable pages on the active and
inactive anon LRU lists.
Last but not least, the swap token code has been broken for a year
without complaints, as reported by Konstantin Khlebnikov. This suggests
we no longer have much use for it.
The days of sub-1G memory systems with heavy use of swap are over. If
we ever need thrashing reducing code in the future, we will have to
implement something that does scale.
Signed-off-by: Rik van Riel <riel@redhat.com>
Cc: Konstantin Khlebnikov <khlebnikov@openvz.org>
Acked-by: Johannes Weiner <hannes@cmpxchg.org>
Cc: Mel Gorman <mel@csn.ul.ie>
Cc: Hugh Dickins <hughd@google.com>
Acked-by: Bob Picco <bpicco@meloft.net>
Acked-by: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Uprobes executes the original instruction at a probed location
out of line. For this, we allocate a page (per mm) upon the
first uprobe hit, in the process user address space, divide it
into slots that are used to store the actual instructions to be
singlestepped. These slots are known as xol (execution out of
line) slots.
Care is taken to ensure that the allocation is in an unmapped
area as close to the top of the user address space as possible,
with appropriate permission settings to keep selinux like
frameworks happy.
Upon a uprobe hit, a free slot is acquired, and is released
after the singlestep completes.
Lots of improvements courtesy suggestions/inputs from Peter and
Oleg.
[ Folded a fix for build issue on powerpc fixed and reported by
Stephen Rothwell. ]
Signed-off-by: Srikar Dronamraju <srikar@linux.vnet.ibm.com>
Cc: Linus Torvalds <torvalds@linux-foundation.org>
Cc: Ananth N Mavinakayanahalli <ananth@in.ibm.com>
Cc: Jim Keniston <jkenisto@linux.vnet.ibm.com>
Cc: Linux-mm <linux-mm@kvack.org>
Cc: Oleg Nesterov <oleg@redhat.com>
Cc: Andi Kleen <andi@firstfloor.org>
Cc: Christoph Hellwig <hch@infradead.org>
Cc: Steven Rostedt <rostedt@goodmis.org>
Cc: Arnaldo Carvalho de Melo <acme@infradead.org>
Cc: Masami Hiramatsu <masami.hiramatsu.pt@hitachi.com>
Cc: Anton Arapov <anton@redhat.com>
Cc: Peter Zijlstra <peterz@infradead.org>
Link: http://lkml.kernel.org/r/20120330182631.10018.48175.sendpatchset@srdronam.in.ibm.com
Signed-off-by: Ingo Molnar <mingo@kernel.org>
While implementing cmpxchg_double() on s390 I realized that we don't set
CONFIG_CMPXCHG_LOCAL despite the fact that we have support for it.
However setting that option will increase the size of struct page by
eight bytes on 64 bit, which we certainly do not want. Also, it doesn't
make sense that a present cpu feature should increase the size of struct
page.
Besides that it looks like the dependency to CMPXCHG_LOCAL is wrong and
that it should depend on CMPXCHG_DOUBLE instead.
This patch:
If an architecture supports CMPXCHG_LOCAL this shouldn't result
automatically in larger struct pages if the SLUB allocator is used.
Instead introduce a new config option "HAVE_ALIGNED_STRUCT_PAGE" which
can be selected if a double word aligned struct page is required. Also
update x86 Kconfig so that it should work as before.
Signed-off-by: Heiko Carstens <heiko.carstens@de.ibm.com>
Acked-by: Christoph Lameter <cl@linux.com>
Cc: Pekka Enberg <penberg@kernel.org>
Cc: Ingo Molnar <mingo@elte.hu>
Cc: Thomas Gleixner <tglx@linutronix.de>
Cc: "H. Peter Anvin" <hpa@zytor.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Michel while working on the working set estimation code, noticed that
calling get_page_unless_zero() on a random pfn_to_page(random_pfn)
wasn't safe, if the pfn ended up being a tail page of a transparent
hugepage under splitting by __split_huge_page_refcount().
He then found the problem could also theoretically materialize with
page_cache_get_speculative() during the speculative radix tree lookups
that uses get_page_unless_zero() in SMP if the radix tree page is freed
and reallocated and get_user_pages is called on it before
page_cache_get_speculative has a chance to call get_page_unless_zero().
So the best way to fix the problem is to keep page_tail->_count zero at
all times. This will guarantee that get_page_unless_zero() can never
succeed on any tail page. page_tail->_mapcount is guaranteed zero and
is unused for all tail pages of a compound page, so we can simply
account the tail page references there and transfer them to
tail_page->_count in __split_huge_page_refcount() (in addition to the
head_page->_mapcount).
While debugging this s/_count/_mapcount/ change I also noticed get_page is
called by direct-io.c on pages returned by get_user_pages. That wasn't
entirely safe because the two atomic_inc in get_page weren't atomic. As
opposed to other get_user_page users like secondary-MMU page fault to
establish the shadow pagetables would never call any superflous get_page
after get_user_page returns. It's safer to make get_page universally safe
for tail pages and to use get_page_foll() within follow_page (inside
get_user_pages()). get_page_foll() is safe to do the refcounting for tail
pages without taking any locks because it is run within PT lock protected
critical sections (PT lock for pte and page_table_lock for
pmd_trans_huge).
The standard get_page() as invoked by direct-io instead will now take
the compound_lock but still only for tail pages. The direct-io paths
are usually I/O bound and the compound_lock is per THP so very
finegrined, so there's no risk of scalability issues with it. A simple
direct-io benchmarks with all lockdep prove locking and spinlock
debugging infrastructure enabled shows identical performance and no
overhead. So it's worth it. Ideally direct-io should stop calling
get_page() on pages returned by get_user_pages(). The spinlock in
get_page() is already optimized away for no-THP builds but doing
get_page() on tail pages returned by GUP is generally a rare operation
and usually only run in I/O paths.
This new refcounting on page_tail->_mapcount in addition to avoiding new
RCU critical sections will also allow the working set estimation code to
work without any further complexity associated to the tail page
refcounting with THP.
Signed-off-by: Andrea Arcangeli <aarcange@redhat.com>
Reported-by: Michel Lespinasse <walken@google.com>
Reviewed-by: Michel Lespinasse <walken@google.com>
Reviewed-by: Minchan Kim <minchan.kim@gmail.com>
Cc: Peter Zijlstra <a.p.zijlstra@chello.nl>
Cc: Hugh Dickins <hughd@google.com>
Cc: Johannes Weiner <jweiner@redhat.com>
Cc: Rik van Riel <riel@redhat.com>
Cc: Mel Gorman <mgorman@suse.de>
Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com>
Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org>
Cc: David Gibson <david@gibson.dropbear.id.au>
Cc: <stable@kernel.org>
Cc: <stable@vger.kernel.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Add comments to explain the page statistics field in the mm_struct.
[akpm@linux-foundation.org: add missing ;]
Signed-off-by: Christoph Lameter <cl@linux.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Some kernel components pin user space memory (infiniband and perf) (by
increasing the page count) and account that memory as "mlocked".
The difference between mlocking and pinning is:
A. mlocked pages are marked with PG_mlocked and are exempt from
swapping. Page migration may move them around though.
They are kept on a special LRU list.
B. Pinned pages cannot be moved because something needs to
directly access physical memory. They may not be on any
LRU list.
I recently saw an mlockalled process where mm->locked_vm became
bigger than the virtual size of the process (!) because some
memory was accounted for twice:
Once when the page was mlocked and once when the Infiniband
layer increased the refcount because it needt to pin the RDMA
memory.
This patch introduces a separate counter for pinned pages and
accounts them seperately.
Signed-off-by: Christoph Lameter <cl@linux.com>
Cc: Mike Marciniszyn <infinipath@qlogic.com>
Cc: Roland Dreier <roland@kernel.org>
Cc: Sean Hefty <sean.hefty@intel.com>
Cc: Hugh Dickins <hughd@google.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
This removes mm->oom_disable_count entirely since it's unnecessary and
currently buggy. The counter was intended to be per-process but it's
currently decremented in the exit path for each thread that exits, causing
it to underflow.
The count was originally intended to prevent oom killing threads that
share memory with threads that cannot be killed since it doesn't lead to
future memory freeing. The counter could be fixed to represent all
threads sharing the same mm, but it's better to remove the count since:
- it is possible that the OOM_DISABLE thread sharing memory with the
victim is waiting on that thread to exit and will actually cause
future memory freeing, and
- there is no guarantee that a thread is disabled from oom killing just
because another thread sharing its mm is oom disabled.
Signed-off-by: David Rientjes <rientjes@google.com>
Reported-by: Oleg Nesterov <oleg@redhat.com>
Reviewed-by: Oleg Nesterov <oleg@redhat.com>
Cc: Ying Han <yinghan@google.com>
Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
A few network drivers currently use skb_frag_struct for this purpose but I have
patches which add additional fields and semantics there which these other uses
do not want.
A structure for reference sub-page regions seems like a generally useful thing
so do so instead of adding a network subsystem specific structure.
Signed-off-by: Ian Campbell <ian.campbell@citrix.com>
Acked-by: Jens Axboe <jaxboe@fusionio.com>
Acked-by: David Rientjes <rientjes@google.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
Allow filling out the rest of the kmem_cache_cpu cacheline with pointers to
partial pages. The partial page list is used in slab_free() to avoid
per node lock taking.
In __slab_alloc() we can then take multiple partial pages off the per
node partial list in one go reducing node lock pressure.
We can also use the per cpu partial list in slab_alloc() to avoid scanning
partial lists for pages with free objects.
The main effect of a per cpu partial list is that the per node list_lock
is taken for batches of partial pages instead of individual ones.
Potential future enhancements:
1. The pickup from the partial list could be perhaps be done without disabling
interrupts with some work. The free path already puts the page into the
per cpu partial list without disabling interrupts.
2. __slab_free() may have some code paths that could use optimization.
Performance:
Before After
./hackbench 100 process 200000
Time: 1953.047 1564.614
./hackbench 100 process 20000
Time: 207.176 156.940
./hackbench 100 process 20000
Time: 204.468 156.940
./hackbench 100 process 20000
Time: 204.879 158.772
./hackbench 10 process 20000
Time: 20.153 15.853
./hackbench 10 process 20000
Time: 20.153 15.986
./hackbench 10 process 20000
Time: 19.363 16.111
./hackbench 1 process 20000
Time: 2.518 2.307
./hackbench 1 process 20000
Time: 2.258 2.339
./hackbench 1 process 20000
Time: 2.864 2.163
Signed-off-by: Christoph Lameter <cl@linux.com>
Signed-off-by: Pekka Enberg <penberg@kernel.org>
Restructure the union / struct cascade in struct page so that
we only have one definition of _count.
Tested-by: Hugh Dickins <hughd@google.com>
Signed-off-by: Christoph Lameter <cl@linux.com>
Signed-off-by: Pekka Enberg <penberg@kernel.org>
On Wed, 6 Jul 2011, Jonathan Cameron wrote:
> Getting:
>
> CHK include/linux/version.h
> CHK include/generated/utsrelease.h
> make[1]: `include/generated/mach-types.h' is up to date.
> CC arch/arm/kernel/asm-offsets.s
> In file included from include/linux/sched.h:64:0,
> from arch/arm/kernel/asm-offsets.c:13:
> include/linux/mm_types.h:74:15: error: duplicate member '_count'
> make[1]: *** [arch/arm/kernel/asm-offsets.s] Error 1
> make: *** [prepare0] Error 2
>
> Issue looks to have been introduced by
>
> mm: Rearrange struct page
>
> fc9bb8c768
>
> Guessing it's a known issue, but just thought I'd flag it up in case
> it's something very specific about my build.
>
> gcc-2.6 armv7a
>
> Reverting that patch works, but given I don't know the history, I'm
> not proposing doing that in general!
Well _count exists in two unionized structs but always has the same offset
within the larger struct. Maybe ARM creates different offsets there for
some reason?
The following is a patch to restructure the union / structs combo in such
a way that only a single definition of _count
Reported-by: Jonathan Cameron <jic23@cam.ac.uk>
Tested-by: Piotr Hosowicz <piotr@hosowicz.com>
Signed-off-by: Christoph Lameter <cl@linux.com>
Signed-off-by: Pekka Enberg <penberg@kernel.org>
We need to be able to use cmpxchg_double on the freelist and object count
field in struct page. Rearrange the fields in struct page according to
doubleword entities so that the freelist pointer comes before the counters.
Do the rearranging with a future in mind where we use more doubleword
atomics to avoid locking of updates to flags/mapping or lru pointers.
Create another union to allow access to counters in struct page as a
single unsigned long value.
The doublewords must be properly aligned for cmpxchg_double to work.
Sadly this increases the size of page struct by one word on some architectures.
But as a resultpage structs are now cacheline aligned on x86_64.
Signed-off-by: Christoph Lameter <cl@linux.com>
Signed-off-by: Pekka Enberg <penberg@kernel.org>
Do not use a page flag for the frozen bit. It needs to be part
of the state that is handled with cmpxchg_double(). So use a bit
in the counter struct in the page struct for that purpose.
Signed-off-by: Christoph Lameter <cl@linux.com>
Signed-off-by: Pekka Enberg <penberg@kernel.org>
Thomas Gleixner reports that we now have a boot crash triggered by
CONFIG_CPUMASK_OFFSTACK=y:
BUG: unable to handle kernel NULL pointer dereference at (null)
IP: [<c11ae035>] find_next_bit+0x55/0xb0
Call Trace:
[<c11addda>] cpumask_any_but+0x2a/0x70
[<c102396b>] flush_tlb_mm+0x2b/0x80
[<c1022705>] pud_populate+0x35/0x50
[<c10227ba>] pgd_alloc+0x9a/0xf0
[<c103a3fc>] mm_init+0xec/0x120
[<c103a7a3>] mm_alloc+0x53/0xd0
which was introduced by commit de03c72cfc ("mm: convert
mm->cpu_vm_cpumask into cpumask_var_t"), and is due to wrong ordering of
mm_init() vs mm_init_cpumask
Thomas wrote a patch to just fix the ordering of initialization, but I
hate the new double allocation in the fork path, so I ended up instead
doing some more radical surgery to clean it all up.
Reported-by: Thomas Gleixner <tglx@linutronix.de>
Reported-by: Ingo Molnar <mingo@elte.hu>
Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com>
Cc: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Setup and cleanup of mm_struct->exe_file is currently done in fs/proc/.
This was because exe_file was needed only for /proc/<pid>/exe. Since we
will need the exe_file functionality also for core dumps (so core name can
contain full binary path), built this functionality always into the
kernel.
To achieve that move that out of proc FS to the kernel/ where in fact it
should belong. By doing that we can make dup_mm_exe_file static. Also we
can drop linux/proc_fs.h inclusion in fs/exec.c and kernel/fork.c.
Signed-off-by: Jiri Slaby <jslaby@suse.cz>
Cc: Alexander Viro <viro@zeniv.linux.org.uk>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
The type of vma->vm_flags is 'unsigned long'. Neither 'int' nor
'unsigned int'. This patch fixes such misuse.
Signed-off-by: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com>
[ Changed to use a typedef - we'll extend it to cover more cases
later, since there has been discussion about making it a 64-bit
type.. - Linus ]
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
The problem with having two different types of counters is that developers
adding new code need to keep in mind whether it's safe to use both the
atomic and non-atomic implementations. For example, when adding new
callers of the *_mm_counter() functions a developer needs to ensure that
those paths are always executed with page_table_lock held, in case we're
using the non-atomic implementation of mm counters.
Hugh Dickins introduced the atomic mm counters in commit f412ac08c9
("[PATCH] mm: fix rss and mmlist locking"). When asked why he left the
non-atomic counters around he said,
| The only reason was to avoid adding costly atomic operations into a
| configuration that had no need for them there: the page_table_lock
| sufficed.
|
| Certainly it would be simpler just to delete the non-atomic variant.
|
| And I think it's fair to say that any configuration on which we're
| measuring performance to that degree (rather than "does it boot fast?"
| type measurements), would already be going the split ptlocks route.
Removing the non-atomic counters eases the maintenance burden because
developers no longer have to mindful of the two implementations when using
*_mm_counter().
Note that all architectures provide a means of atomically updating
atomic_long_t variables, even if they have to revert to the generic
spinlock implementation because they don't support 64-bit atomic
instructions (see lib/atomic64.c).
Signed-off-by: Matt Fleming <matt.fleming@linux.intel.com>
Acked-by: Dave Hansen <dave@linux.vnet.ibm.com>
Acked-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Cc: Hugh Dickins <hughd@google.com>
Cc: Christoph Lameter <cl@linux-foundation.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
cpumask_t is very big struct and cpu_vm_mask is placed wrong position.
It might lead to reduce cache hit ratio.
This patch has two change.
1) Move the place of cpumask into last of mm_struct. Because usually cpumask
is accessed only front bits when the system has cpu-hotplug capability
2) Convert cpu_vm_mask into cpumask_var_t. It may help to reduce memory
footprint if cpumask_size() will use nr_cpumask_bits properly in future.
In addition, this patch change the name of cpu_vm_mask with cpu_vm_mask_var.
It may help to detect out of tree cpu_vm_mask users.
This patch has no functional change.
[akpm@linux-foundation.org: build fix]
[akpm@linux-foundation.org: coding-style fixes]
Signed-off-by: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com>
Cc: David Howells <dhowells@redhat.com>
Cc: Koichi Yasutake <yasutake.koichi@jp.panasonic.com>
Cc: Hugh Dickins <hughd@google.com>
Cc: Chris Metcalf <cmetcalf@tilera.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Hugh says:
"The only significant loser, I think, would be page reclaim (when
concurrent with truncation): could spin for a long time waiting for
the i_mmap_mutex it expects would soon be dropped? "
Counter points:
- cpu contention makes the spin stop (need_resched())
- zap pages should be freeing pages at a higher rate than reclaim
ever can
I think the simplification of the truncate code is definitely worth it.
Effectively reverts: 2aa15890f3 ("mm: prevent concurrent
unmap_mapping_range() on the same inode") and takes out the code that
caused its problem.
Signed-off-by: Peter Zijlstra <a.p.zijlstra@chello.nl>
Reviewed-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Cc: Hugh Dickins <hughd@google.com>
Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org>
Cc: David Miller <davem@davemloft.net>
Cc: Martin Schwidefsky <schwidefsky@de.ibm.com>
Cc: Russell King <rmk@arm.linux.org.uk>
Cc: Paul Mundt <lethal@linux-sh.org>
Cc: Jeff Dike <jdike@addtoit.com>
Cc: Richard Weinberger <richard@nod.at>
Cc: Tony Luck <tony.luck@intel.com>
Cc: Mel Gorman <mel@csn.ul.ie>
Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com>
Cc: Nick Piggin <npiggin@kernel.dk>
Cc: Namhyung Kim <namhyung@gmail.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Reorder mm_struct to remove 16 bytes of alignment padding on 64 bit
builds. On my config this shrinks mm_struct by enough to fit in one
fewer cache lines and allows more objects per slab in mm_struct
kmem_cache under SLUB.
slabinfo before patch :-
Sizes (bytes) Slabs
--------------------------------
Object : 848 Total : 9
SlabObj: 896 Full : 2
SlabSiz: 16384 Partial: 5
Loss : 48 CpuSlab: 2
Align : 64 Objects: 18
slabinfo after :-
Sizes (bytes) Slabs
--------------------------------
Object : 832 Total : 7
SlabObj: 832 Full : 2
SlabSiz: 16384 Partial: 3
Loss : 0 CpuSlab: 2
Align : 64 Objects: 19
Signed-off-by: Richard Kennedy <richard@rsk.demon.co.uk>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
This increase the size of the mm struct a bit but it is needed to
preallocate one pte for each hugepage so that split_huge_page will not
require a fail path. Guarantee of success is a fundamental property of
split_huge_page to avoid decrasing swapping reliability and to avoid
adding -ENOMEM fail paths that would otherwise force the hugepage-unaware
VM code to learn rolling back in the middle of its pte mangling operations
(if something we need it to learn handling pmd_trans_huge natively rather
being capable of rollback). When split_huge_page runs a pte is needed to
succeed the split, to map the newly splitted regular pages with a regular
pte. This way all existing VM code remains backwards compatible by just
adding a split_huge_page* one liner. The memory waste of those
preallocated ptes is negligible and so it is worth it.
Signed-off-by: Andrea Arcangeli <aarcange@redhat.com>
Acked-by: Rik van Riel <riel@redhat.com>
Acked-by: Mel Gorman <mel@csn.ul.ie>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
It's pointless to kill a task if another thread sharing its mm cannot be
killed to allow future memory freeing. A subsequent patch will prevent
kills in such cases, but first it's necessary to have a way to flag a task
that shares memory with an OOM_DISABLE task that doesn't incur an
additional tasklist scan, which would make select_bad_process() an O(n^2)
function.
This patch adds an atomic counter to struct mm_struct that follows how
many threads attached to it have an oom_score_adj of OOM_SCORE_ADJ_MIN.
They cannot be killed by the kernel, so their memory cannot be freed in
oom conditions.
This only requires task_lock() on the task that we're operating on, it
does not require mm->mmap_sem since task_lock() pins the mm and the
operation is atomic.
[rientjes@google.com: changelog and sys_unshare() code]
[rientjes@google.com: protect oom_disable_count with task_lock in fork]
[rientjes@google.com: use old_mm for oom_disable_count in exec]
Signed-off-by: Ying Han <yinghan@google.com>
Signed-off-by: David Rientjes <rientjes@google.com>
Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com>
Cc: Rik van Riel <riel@redhat.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
It's a really simple list, and several of the users want to go backwards
in it to find the previous vma. So rather than have to look up the
previous entry with 'find_vma_prev()' or something similar, just make it
doubly linked instead.
Tested-by: Ian Campbell <ijc@hellion.org.uk>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
This adds annotations for RCU operations in core kernel components
Signed-off-by: Arnd Bergmann <arnd@arndb.de>
Signed-off-by: Paul E. McKenney <paulmck@linux.vnet.ibm.com>
Cc: Al Viro <viro@ZenIV.linux.org.uk>
Cc: Jens Axboe <jens.axboe@oracle.com>
Cc: Andrew Morton <akpm@linux-foundation.org>
Reviewed-by: Josh Triplett <josh@joshtriplett.org>
Commit 34e55232e5 ("mm: avoid false sharing
of mm_counter") added sync_mm_rss() for syncing loosely accounted rss
counters. It's for CONFIG_MMU but sync_mm_rss is called even in NOMMU
enviroment (kerne/exit.c, fs/exec.c). Above commit doesn't handle it
well.
This patch changes
SPLIT_RSS_COUNTING depends on SPLIT_PTLOCKS && CONFIG_MMU
And for avoid unnecessary function calls, sync_mm_rss changed to be inlined
noop function in header file.
Reported-by: David Howells <dhowells@redhat.com>
Signed-off-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Signed-off-by: Mike Frysinger <vapier@gentoo.org>
Signed-off-by: Michal Simek <monstr@monstr.eu>
Signed-off-by: David Howells <dhowells@redhat.com>
Cc: Greg Ungerer <gerg@snapgear.com>
Cc: Geert Uytterhoeven <geert@linux-m68k.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
The old anon_vma code can lead to scalability issues with heavily forking
workloads. Specifically, each anon_vma will be shared between the parent
process and all its child processes.
In a workload with 1000 child processes and a VMA with 1000 anonymous
pages per process that get COWed, this leads to a system with a million
anonymous pages in the same anon_vma, each of which is mapped in just one
of the 1000 processes. However, the current rmap code needs to walk them
all, leading to O(N) scanning complexity for each page.
This can result in systems where one CPU is walking the page tables of
1000 processes in page_referenced_one, while all other CPUs are stuck on
the anon_vma lock. This leads to catastrophic failure for a benchmark
like AIM7, where the total number of processes can reach in the tens of
thousands. Real workloads are still a factor 10 less process intensive
than AIM7, but they are catching up.
This patch changes the way anon_vmas and VMAs are linked, which allows us
to associate multiple anon_vmas with a VMA. At fork time, each child
process gets its own anon_vmas, in which its COWed pages will be
instantiated. The parents' anon_vma is also linked to the VMA, because
non-COWed pages could be present in any of the children.
This reduces rmap scanning complexity to O(1) for the pages of the 1000
child processes, with O(N) complexity for at most 1/N pages in the system.
This reduces the average scanning cost in heavily forking workloads from
O(N) to 2.
The only real complexity in this patch stems from the fact that linking a
VMA to anon_vmas now involves memory allocations. This means vma_adjust
can fail, if it needs to attach a VMA to anon_vma structures. This in
turn means error handling needs to be added to the calling functions.
A second source of complexity is that, because there can be multiple
anon_vmas, the anon_vma linking in vma_adjust can no longer be done under
"the" anon_vma lock. To prevent the rmap code from walking up an
incomplete VMA, this patch introduces the VM_LOCK_RMAP VMA flag. This bit
flag uses the same slot as the NOMMU VM_MAPPED_COPY, with an ifdef in mm.h
to make sure it is impossible to compile a kernel that needs both symbolic
values for the same bitflag.
Some test results:
Without the anon_vma changes, when AIM7 hits around 9.7k users (on a test
box with 16GB RAM and not quite enough IO), the system ends up running
>99% in system time, with every CPU on the same anon_vma lock in the
pageout code.
With these changes, AIM7 hits the cross-over point around 29.7k users.
This happens with ~99% IO wait time, there never seems to be any spike in
system time. The anon_vma lock contention appears to be resolved.
[akpm@linux-foundation.org: cleanups]
Signed-off-by: Rik van Riel <riel@redhat.com>
Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com>
Cc: Larry Woodman <lwoodman@redhat.com>
Cc: Lee Schermerhorn <Lee.Schermerhorn@hp.com>
Cc: Minchan Kim <minchan.kim@gmail.com>
Cc: Andrea Arcangeli <aarcange@redhat.com>
Cc: Hugh Dickins <hugh.dickins@tiscali.co.uk>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
A frequent questions from users about memory management is what numbers of
swap ents are user for processes. And this information will give some
hints to oom-killer.
Besides we can count the number of swapents per a process by scanning
/proc/<pid>/smaps, this is very slow and not good for usual process
information handler which works like 'ps' or 'top'. (ps or top is now
enough slow..)
This patch adds a counter of swapents to mm_counter and update is at each
swap events. Information is exported via /proc/<pid>/status file as
[kamezawa@bluextal memory]$ cat /proc/self/status
Name: cat
State: R (running)
Tgid: 2910
Pid: 2910
PPid: 2823
TracerPid: 0
Uid: 500 500 500 500
Gid: 500 500 500 500
FDSize: 256
Groups: 500
VmPeak: 82696 kB
VmSize: 82696 kB
VmLck: 0 kB
VmHWM: 432 kB
VmRSS: 432 kB
VmData: 172 kB
VmStk: 84 kB
VmExe: 48 kB
VmLib: 1568 kB
VmPTE: 40 kB
VmSwap: 0 kB <=============== this.
[akpm@linux-foundation.org: coding-style fixes]
Signed-off-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Reviewed-by: Minchan Kim <minchan.kim@gmail.com>
Reviewed-by: Christoph Lameter <cl@linux-foundation.org>
Cc: Lee Schermerhorn <lee.schermerhorn@hp.com>
Cc: David Rientjes <rientjes@google.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Considering the nature of per mm stats, it's the shared object among
threads and can be a cache-miss point in the page fault path.
This patch adds per-thread cache for mm_counter. RSS value will be
counted into a struct in task_struct and synchronized with mm's one at
events.
Now, in this patch, the event is the number of calls to handle_mm_fault.
Per-thread value is added to mm at each 64 calls.
rough estimation with small benchmark on parallel thread (2threads) shows
[before]
4.5 cache-miss/faults
[after]
4.0 cache-miss/faults
Anyway, the most contended object is mmap_sem if the number of threads grows.
[akpm@linux-foundation.org: coding-style fixes]
Signed-off-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Cc: Minchan Kim <minchan.kim@gmail.com>
Cc: Christoph Lameter <cl@linux-foundation.org>
Cc: Lee Schermerhorn <lee.schermerhorn@hp.com>
Cc: David Rientjes <rientjes@google.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Presently, per-mm statistics counter is defined by macro in sched.h
This patch modifies it to
- defined in mm.h as inlinf functions
- use array instead of macro's name creation.
This patch is for reducing patch size in future patch to modify
implementation of per-mm counter.
Signed-off-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Reviewed-by: Minchan Kim <minchan.kim@gmail.com>
Cc: Christoph Lameter <cl@linux-foundation.org>
Cc: Lee Schermerhorn <lee.schermerhorn@hp.com>
Cc: David Rientjes <rientjes@google.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
get_unmapped_area() is unnecessary for NOMMU as no-one calls it.
Signed-off-by: David Howells <dhowells@redhat.com>
Acked-by: Al Viro <viro@zeniv.linux.org.uk>
Cc: Greg Ungerer <gerg@snapgear.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
The vm_usage count field in struct vm_region does not need to be atomic as
it's only even modified whilst nommu_region_sem is write locked.
Signed-off-by: David Howells <dhowells@redhat.com>
Acked-by: Al Viro <viro@zeniv.linux.org.uk>
Cc: Greg Ungerer <gerg@snapgear.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
When working with FDPIC, there are many shared mappings of read-only
code regions between applications (the C library, applet packages like
busybox, etc.), but the current do_mmap_pgoff() function will issue an
icache flush whenever a VMA is added to an MM instead of only doing it
when the map is initially created.
The flush can instead be done when a region is first mmapped PROT_EXEC.
Note that we may not rely on the first mapping of a region being
executable - it's possible for it to be PROT_READ only, so we have to
remember whether we've flushed the region or not, and then flush the
entire region when a bit of it is made executable.
However, this also affects the brk area. That will no longer be
executable. We can mprotect() it to PROT_EXEC on MPU-mode kernels, but
for NOMMU mode kernels, when it increases the brk allocation, making
sys_brk() flush the extra from the icache should suffice. The brk area
probably isn't used by NOMMU programs since the brk area can only use up
the leavings from the stack allocation, where the stack allocation is
larger than requested.
Signed-off-by: David Howells <dhowells@redhat.com>
Signed-off-by: Mike Frysinger <vapier@gentoo.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
* 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: Alexey Dobriyan <adobriyan@gmail.com>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Because the binfmt is not different between threads in the same process,
it can be moved from task_struct to mm_struct. And binfmt moudle is
handled per mm_struct instead of task_struct.
Signed-off-by: Hiroshi Shimamoto <h-shimamoto@ct.jp.nec.com>
Acked-by: Oleg Nesterov <oleg@redhat.com>
Cc: Rusty Russell <rusty@rustcorp.com.au>
Acked-by: Roland McGrath <roland@redhat.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
->ioctx_lock and ->ioctx_list are used only under CONFIG_AIO.
Signed-off-by: Alexey Dobriyan <adobriyan@gmail.com>
Cc: Zach Brown <zach.brown@oracle.com>
Cc: Benjamin LaHaise <bcrl@kvack.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
The commit 2ff05b2b (oom: move oom_adj value) moveed the oom_adj value to
the mm_struct. It was a very good first step for sanitize OOM.
However Paul Menage reported the commit makes regression to his job
scheduler. Current OOM logic can kill OOM_DISABLED process.
Why? His program has the code of similar to the following.
...
set_oom_adj(OOM_DISABLE); /* The job scheduler never killed by oom */
...
if (vfork() == 0) {
set_oom_adj(0); /* Invoked child can be killed */
execve("foo-bar-cmd");
}
....
vfork() parent and child are shared the same mm_struct. then above
set_oom_adj(0) doesn't only change oom_adj for vfork() child, it's also
change oom_adj for vfork() parent. Then, vfork() parent (job scheduler)
lost OOM immune and it was killed.
Actually, fork-setting-exec idiom is very frequently used in userland program.
We must not break this assumption.
Then, this patch revert commit 2ff05b2b and related commit.
Reverted commit list
---------------------
- commit 2ff05b2b4e (oom: move oom_adj value from task_struct to mm_struct)
- commit 4d8b9135c3 (oom: avoid unnecessary mm locking and scanning for OOM_DISABLE)
- commit 8123681022 (oom: only oom kill exiting tasks with attached memory)
- commit 933b787b57 (mm: copy over oom_adj value at fork time)
Signed-off-by: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com>
Cc: Paul Menage <menage@google.com>
Cc: David Rientjes <rientjes@google.com>
Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Cc: Rik van Riel <riel@redhat.com>
Cc: Linus Torvalds <torvalds@linux-foundation.org>
Cc: Oleg Nesterov <oleg@redhat.com>
Cc: Nick Piggin <npiggin@suse.de>
Cc: Mel Gorman <mel@csn.ul.ie>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
* akpm: (182 commits)
fbdev: bf54x-lq043fb: use kzalloc over kmalloc/memset
fbdev: *bfin*: fix __dev{init,exit} markings
fbdev: *bfin*: drop unnecessary calls to memset
fbdev: bfin-t350mcqb-fb: drop unused local variables
fbdev: blackfin has __raw I/O accessors, so use them in fb.h
fbdev: s1d13xxxfb: add accelerated bitblt functions
tcx: use standard fields for framebuffer physical address and length
fbdev: add support for handoff from firmware to hw framebuffers
intelfb: fix a bug when changing video timing
fbdev: use framebuffer_release() for freeing fb_info structures
radeon: P2G2CLK_ALWAYS_ONb tested twice, should 2nd be P2G2CLK_DAC_ALWAYS_ONb?
s3c-fb: CPUFREQ frequency scaling support
s3c-fb: fix resource releasing on error during probing
carminefb: fix possible access beyond end of carmine_modedb[]
acornfb: remove fb_mmap function
mb862xxfb: use CONFIG_OF instead of CONFIG_PPC_OF
mb862xxfb: restrict compliation of platform driver to PPC
Samsung SoC Framebuffer driver: add Alpha Channel support
atmel-lcdc: fix pixclock upper bound detection
offb: use framebuffer_alloc() to allocate fb_info struct
...
Manually fix up conflicts due to kmemcheck in mm/slab.c
The per-task oom_adj value is a characteristic of its mm more than the
task itself since it's not possible to oom kill any thread that shares the
mm. If a task were to be killed while attached to an mm that could not be
freed because another thread were set to OOM_DISABLE, it would have
needlessly been terminated since there is no potential for future memory
freeing.
This patch moves oomkilladj (now more appropriately named oom_adj) from
struct task_struct to struct mm_struct. This requires task_lock() on a
task to check its oom_adj value to protect against exec, but it's already
necessary to take the lock when dereferencing the mm to find the total VM
size for the badness heuristic.
This fixes a livelock if the oom killer chooses a task and another thread
sharing the same memory has an oom_adj value of OOM_DISABLE. This occurs
because oom_kill_task() repeatedly returns 1 and refuses to kill the
chosen task while select_bad_process() will repeatedly choose the same
task during the next retry.
Taking task_lock() in select_bad_process() to check for OOM_DISABLE and in
oom_kill_task() to check for threads sharing the same memory will be
removed in the next patch in this series where it will no longer be
necessary.
Writing to /proc/pid/oom_adj for a kthread will now return -EINVAL since
these threads are immune from oom killing already. They simply report an
oom_adj value of OOM_DISABLE.
Cc: Nick Piggin <npiggin@suse.de>
Cc: Rik van Riel <riel@redhat.com>
Cc: Mel Gorman <mel@csn.ul.ie>
Signed-off-by: David Rientjes <rientjes@google.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
General description: kmemcheck is a patch to the linux kernel that
detects use of uninitialized memory. It does this by trapping every
read and write to memory that was allocated dynamically (e.g. using
kmalloc()). If a memory address is read that has not previously been
written to, a message is printed to the kernel log.
Thanks to Andi Kleen for the set_memory_4k() solution.
Andrew Morton suggested documenting the shadow member of struct page.
Signed-off-by: Vegard Nossum <vegardno@ifi.uio.no>
Signed-off-by: Pekka Enberg <penberg@cs.helsinki.fi>
[export kmemcheck_mark_initialized]
[build fix for setup_max_cpus]
Signed-off-by: Ingo Molnar <mingo@elte.hu>
[rebased for mainline inclusion]
Signed-off-by: Vegard Nossum <vegardno@ifi.uio.no>
This fixes a build failure with generic debug pagealloc:
mm/debug-pagealloc.c: In function 'set_page_poison':
mm/debug-pagealloc.c:8: error: 'struct page' has no member named 'debug_flags'
mm/debug-pagealloc.c: In function 'clear_page_poison':
mm/debug-pagealloc.c:13: error: 'struct page' has no member named 'debug_flags'
mm/debug-pagealloc.c: In function 'page_poison':
mm/debug-pagealloc.c:18: error: 'struct page' has no member named 'debug_flags'
mm/debug-pagealloc.c: At top level:
mm/debug-pagealloc.c:120: error: redefinition of 'kernel_map_pages'
include/linux/mm.h:1278: error: previous definition of 'kernel_map_pages' was here
mm/debug-pagealloc.c: In function 'kernel_map_pages':
mm/debug-pagealloc.c:122: error: 'debug_pagealloc_enabled' undeclared (first use in this function)
by fixing
- debug_flags should be in struct page
- define DEBUG_PAGEALLOC config option for all architectures
Signed-off-by: Akinobu Mita <akinobu.mita@gmail.com>
Reported-by: Alexander Beregalov <a.beregalov@gmail.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
CONFIG_DEBUG_PAGEALLOC is now supported by x86, powerpc, sparc64, and
s390. This patch implements it for the rest of the architectures by
filling the pages with poison byte patterns after free_pages() and
verifying the poison patterns before alloc_pages().
This generic one cannot detect invalid page accesses immediately but
invalid read access may cause invalid dereference by poisoned memory and
invalid write access can be detected after a long delay.
Signed-off-by: Akinobu Mita <akinobu.mita@gmail.com>
Cc: <linux-arch@vger.kernel.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
This allows us to change the representation (to a dangling bitmap or
cpumask_var_t) without breaking all the callers: they can use
mm_cpumask() now and won't see a difference as the changes roll into
linux-next.
Signed-off-by: Rusty Russell <rusty@rustcorp.com.au>
NOMMU mmap allocates a piece of memory for an mmap that's rounded up in size to
the nearest power-of-2 number of pages. Currently it then discards the excess
pages back to the page allocator, making that memory available for use by other
things. This can, however, cause greater amount of fragmentation.
To counter this, a sysctl is added in order to fine-tune the trimming
behaviour. The default behaviour remains to trim pages aggressively, while
this can either be disabled completely or set to a higher page-granular
watermark in order to have finer-grained control.
vm region vm_top bits taken from an earlier patch by David Howells.
Signed-off-by: Paul Mundt <lethal@linux-sh.org>
Signed-off-by: David Howells <dhowells@redhat.com>
Tested-by: Mike Frysinger <vapier.adi@gmail.com>
Make VMAs per mm_struct as for MMU-mode linux. This solves two problems:
(1) In SYSV SHM where nattch for a segment does not reflect the number of
shmat's (and forks) done.
(2) In mmap() where the VMA's vm_mm is set to point to the parent mm by an
exec'ing process when VM_EXECUTABLE is specified, regardless of the fact
that a VMA might be shared and already have its vm_mm assigned to another
process or a dead process.
A new struct (vm_region) is introduced to track a mapped region and to remember
the circumstances under which it may be shared and the vm_list_struct structure
is discarded as it's no longer required.
This patch makes the following additional changes:
(1) Regions are now allocated with alloc_pages() rather than kmalloc() and
with no recourse to __GFP_COMP, so the pages are not composite. Instead,
each page has a reference on it held by the region. Anything else that is
interested in such a page will have to get a reference on it to retain it.
When the pages are released due to unmapping, each page is passed to
put_page() and will be freed when the page usage count reaches zero.
(2) Excess pages are trimmed after an allocation as the allocation must be
made as a power-of-2 quantity of pages.
(3) VMAs are added to the parent MM's R/B tree and mmap lists. As an MM may
end up with overlapping VMAs within the tree, the VMA struct address is
appended to the sort key.
(4) Non-anonymous VMAs are now added to the backing inode's prio list.
(5) Holes may be punched in anonymous VMAs with munmap(), releasing parts of
the backing region. The VMA and region structs will be split if
necessary.
(6) sys_shmdt() only releases one attachment to a SYSV IPC shared memory
segment instead of all the attachments at that addresss. Multiple
shmat()'s return the same address under NOMMU-mode instead of different
virtual addresses as under MMU-mode.
(7) Core dumping for ELF-FDPIC requires fewer exceptions for NOMMU-mode.
(8) /proc/maps is now the global list of mapped regions, and may list bits
that aren't actually mapped anywhere.
(9) /proc/meminfo gains a line (tagged "MmapCopy") that indicates the amount
of RAM currently allocated by mmap to hold mappable regions that can't be
mapped directly. These are copies of the backing device or file if not
anonymous.
These changes make NOMMU mode more similar to MMU mode. The downside is that
NOMMU mode requires some extra memory to track things over NOMMU without this
patch (VMAs are no longer shared, and there are now region structs).
Signed-off-by: David Howells <dhowells@redhat.com>
Tested-by: Mike Frysinger <vapier.adi@gmail.com>
Acked-by: Paul Mundt <lethal@linux-sh.org>
The mm->ioctx_list is currently protected by a reader-writer lock,
so we always grab that lock on the read side for doing ioctx
lookups. As the workload is extremely reader biased, turn this into
an rcu hlist so we can make lookup_ioctx() lockless. Get rid of
the rwlock and use a spinlock for providing update side exclusion.
There's usually only 1 entry on this list, so it doesn't make sense
to look into fancier data structures.
Reviewed-by: Jeff Moyer <jmoyer@redhat.com>
Signed-off-by: Jens Axboe <jens.axboe@oracle.com>
Allocate all page_cgroup at boot and remove page_cgroup poitner from
struct page. This patch adds an interface as
struct page_cgroup *lookup_page_cgroup(struct page*)
All FLATMEM/DISCONTIGMEM/SPARSEMEM and MEMORY_HOTPLUG is supported.
Remove page_cgroup pointer reduces the amount of memory by
- 4 bytes per PAGE_SIZE.
- 8 bytes per PAGE_SIZE
if memory controller is disabled. (even if configured.)
On usual 8GB x86-32 server, this saves 8MB of NORMAL_ZONE memory.
On my x86-64 server with 48GB of memory, this saves 96MB of memory.
I think this reduction makes sense.
By pre-allocation, kmalloc/kfree in charge/uncharge are removed.
This means
- we're not necessary to be afraid of kmalloc faiulre.
(this can happen because of gfp_mask type.)
- we can avoid calling kmalloc/kfree.
- we can avoid allocating tons of small objects which can be fragmented.
- we can know what amount of memory will be used for this extra-lru handling.
I added printk message as
"allocated %ld bytes of page_cgroup"
"please try cgroup_disable=memory option if you don't want"
maybe enough informative for users.
Signed-off-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Reviewed-by: Balbir Singh <balbir@linux.vnet.ibm.com>
Cc: Daisuke Nishimura <nishimura@mxp.nes.nec.co.jp>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Define USE_SPLIT_PTLOCKS as a constant expression rather than repeating
"NR_CPUS >= CONFIG_SPLIT_PTLOCK_CPUS" all over the place.
Signed-off-by: Jeremy Fitzhardinge <jeremy.fitzhardinge@citrix.com>
Acked-by: Hugh Dickins <hugh@veritas.com>
Signed-off-by: Ingo Molnar <mingo@elte.hu>
Try to comment away a little of the confusion between mm's vm_area_struct
vm_flags and vmalloc's vm_struct flags: based on an idea by Ulrich Drepper.
Signed-off-by: Hugh Dickins <hugh@veritas.com>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
With KVM/GFP/XPMEM there isn't just the primary CPU MMU pointing to pages.
There are secondary MMUs (with secondary sptes and secondary tlbs) too.
sptes in the kvm case are shadow pagetables, but when I say spte in
mmu-notifier context, I mean "secondary pte". In GRU case there's no
actual secondary pte and there's only a secondary tlb because the GRU
secondary MMU has no knowledge about sptes and every secondary tlb miss
event in the MMU always generates a page fault that has to be resolved by
the CPU (this is not the case of KVM where the a secondary tlb miss will
walk sptes in hardware and it will refill the secondary tlb transparently
to software if the corresponding spte is present). The same way
zap_page_range has to invalidate the pte before freeing the page, the spte
(and secondary tlb) must also be invalidated before any page is freed and
reused.
Currently we take a page_count pin on every page mapped by sptes, but that
means the pages can't be swapped whenever they're mapped by any spte
because they're part of the guest working set. Furthermore a spte unmap
event can immediately lead to a page to be freed when the pin is released
(so requiring the same complex and relatively slow tlb_gather smp safe
logic we have in zap_page_range and that can be avoided completely if the
spte unmap event doesn't require an unpin of the page previously mapped in
the secondary MMU).
The mmu notifiers allow kvm/GRU/XPMEM to attach to the tsk->mm and know
when the VM is swapping or freeing or doing anything on the primary MMU so
that the secondary MMU code can drop sptes before the pages are freed,
avoiding all page pinning and allowing 100% reliable swapping of guest
physical address space. Furthermore it avoids the code that teardown the
mappings of the secondary MMU, to implement a logic like tlb_gather in
zap_page_range that would require many IPI to flush other cpu tlbs, for
each fixed number of spte unmapped.
To make an example: if what happens on the primary MMU is a protection
downgrade (from writeable to wrprotect) the secondary MMU mappings will be
invalidated, and the next secondary-mmu-page-fault will call
get_user_pages and trigger a do_wp_page through get_user_pages if it
called get_user_pages with write=1, and it'll re-establishing an updated
spte or secondary-tlb-mapping on the copied page. Or it will setup a
readonly spte or readonly tlb mapping if it's a guest-read, if it calls
get_user_pages with write=0. This is just an example.
This allows to map any page pointed by any pte (and in turn visible in the
primary CPU MMU), into a secondary MMU (be it a pure tlb like GRU, or an
full MMU with both sptes and secondary-tlb like the shadow-pagetable layer
with kvm), or a remote DMA in software like XPMEM (hence needing of
schedule in XPMEM code to send the invalidate to the remote node, while no
need to schedule in kvm/gru as it's an immediate event like invalidating
primary-mmu pte).
At least for KVM without this patch it's impossible to swap guests
reliably. And having this feature and removing the page pin allows
several other optimizations that simplify life considerably.
Dependencies:
1) mm_take_all_locks() to register the mmu notifier when the whole VM
isn't doing anything with "mm". This allows mmu notifier users to keep
track if the VM is in the middle of the invalidate_range_begin/end
critical section with an atomic counter incraese in range_begin and
decreased in range_end. No secondary MMU page fault is allowed to map
any spte or secondary tlb reference, while the VM is in the middle of
range_begin/end as any page returned by get_user_pages in that critical
section could later immediately be freed without any further
->invalidate_page notification (invalidate_range_begin/end works on
ranges and ->invalidate_page isn't called immediately before freeing
the page). To stop all page freeing and pagetable overwrites the
mmap_sem must be taken in write mode and all other anon_vma/i_mmap
locks must be taken too.
2) It'd be a waste to add branches in the VM if nobody could possibly
run KVM/GRU/XPMEM on the kernel, so mmu notifiers will only enabled if
CONFIG_KVM=m/y. In the current kernel kvm won't yet take advantage of
mmu notifiers, but this already allows to compile a KVM external module
against a kernel with mmu notifiers enabled and from the next pull from
kvm.git we'll start using them. And GRU/XPMEM will also be able to
continue the development by enabling KVM=m in their config, until they
submit all GRU/XPMEM GPLv2 code to the mainline kernel. Then they can
also enable MMU_NOTIFIERS in the same way KVM does it (even if KVM=n).
This guarantees nobody selects MMU_NOTIFIER=y if KVM and GRU and XPMEM
are all =n.
The mmu_notifier_register call can fail because mm_take_all_locks may be
interrupted by a signal and return -EINTR. Because mmu_notifier_reigster
is used when a driver startup, a failure can be gracefully handled. Here
an example of the change applied to kvm to register the mmu notifiers.
Usually when a driver startups other allocations are required anyway and
-ENOMEM failure paths exists already.
struct kvm *kvm_arch_create_vm(void)
{
struct kvm *kvm = kzalloc(sizeof(struct kvm), GFP_KERNEL);
+ int err;
if (!kvm)
return ERR_PTR(-ENOMEM);
INIT_LIST_HEAD(&kvm->arch.active_mmu_pages);
+ kvm->arch.mmu_notifier.ops = &kvm_mmu_notifier_ops;
+ err = mmu_notifier_register(&kvm->arch.mmu_notifier, current->mm);
+ if (err) {
+ kfree(kvm);
+ return ERR_PTR(err);
+ }
+
return kvm;
}
mmu_notifier_unregister returns void and it's reliable.
The patch also adds a few needed but missing includes that would prevent
kernel to compile after these changes on non-x86 archs (x86 didn't need
them by luck).
[akpm@linux-foundation.org: coding-style fixes]
[akpm@linux-foundation.org: fix mm/filemap_xip.c build]
[akpm@linux-foundation.org: fix mm/mmu_notifier.c build]
Signed-off-by: Andrea Arcangeli <andrea@qumranet.com>
Signed-off-by: Nick Piggin <npiggin@suse.de>
Signed-off-by: Christoph Lameter <cl@linux-foundation.org>
Cc: Jack Steiner <steiner@sgi.com>
Cc: Robin Holt <holt@sgi.com>
Cc: Nick Piggin <npiggin@suse.de>
Cc: Peter Zijlstra <a.p.zijlstra@chello.nl>
Cc: Kanoj Sarcar <kanojsarcar@yahoo.com>
Cc: Roland Dreier <rdreier@cisco.com>
Cc: Steve Wise <swise@opengridcomputing.com>
Cc: Avi Kivity <avi@qumranet.com>
Cc: Hugh Dickins <hugh@veritas.com>
Cc: Rusty Russell <rusty@rustcorp.com.au>
Cc: Anthony Liguori <aliguori@us.ibm.com>
Cc: Chris Wright <chrisw@redhat.com>
Cc: Marcelo Tosatti <marcelo@kvack.org>
Cc: Eric Dumazet <dada1@cosmosbay.com>
Cc: "Paul E. McKenney" <paulmck@us.ibm.com>
Cc: Izik Eidus <izike@qumranet.com>
Cc: Anthony Liguori <aliguori@us.ibm.com>
Cc: Rik van Riel <riel@redhat.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Now that we have core_state->dumper list we can use it to wake up the
sub-threads waiting for the coredump completion.
This uglifies the code and .text grows by 47 bytes, but otoh mm_struct
lessens by sizeof(struct completion). Also, with this change we can
decouple exit_mm() from the coredumping code.
Signed-off-by: Oleg Nesterov <oleg@tv-sign.ru>
Cc: Roland McGrath <roland@redhat.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
binfmt->core_dump() has to iterate over the all threads in system in order
to find the coredumping threads and construct the list using the
GFP_ATOMIC allocations.
With this patch each thread allocates the list node on exit_mm()'s stack and
adds itself to the list.
This allows us to do further changes:
- simplify ->core_dump()
- change exit_mm() to clear ->mm first, then wait for ->core_done.
this makes the coredumping process visible to oom_kill
- kill mm->core_done
Signed-off-by: Oleg Nesterov <oleg@tv-sign.ru>
Acked-by: Roland McGrath <roland@redhat.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Turn core_state->nr_threads into atomic_t and kill now unneeded
down_write(&mm->mmap_sem) in exit_mm().
Signed-off-by: Oleg Nesterov <oleg@tv-sign.ru>
Cc: Roland McGrath <roland@redhat.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Move mm->core_waiters into "struct core_state" allocated on stack. This
shrinks mm_struct a little bit and allows further changes.
This patch mostly does s/core_waiters/core_state. The only essential
change is that coredump_wait() must clear mm->core_state before return.
The coredump_wait()'s path is uglified and .text grows by 30 bytes, this
is fixed by the next patch.
Signed-off-by: Oleg Nesterov <oleg@tv-sign.ru>
Cc: Roland McGrath <roland@redhat.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
mm->core_startup_done points to "struct completion startup_done" allocated
on the coredump_wait()'s stack. Introduce the new structure, core_state,
which holds this "struct completion". This way we can add more info
visible to the threads participating in coredump without enlarging
mm_struct.
No changes in affected .o files.
Signed-off-by: Oleg Nesterov <oleg@tv-sign.ru>
Cc: Roland McGrath <roland@redhat.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
When mm destruction happens, we should pass mm_update_next_owner() the old mm.
But unfortunately new mm is passed in exec_mmap().
Thus, kernel panic is possible when a multi-threaded process uses exec().
Also, the owner member comment description is wrong. mm->owner does not
necessarily point to the thread group leader.
[akpm@linux-foundation.org: coding-style fixes]
Signed-off-by: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com>
Acked-by: Balbir Singh <balbir@linux.vnet.ibm.com>
Cc: "Paul Menage" <menage@google.com>
Cc: "KAMEZAWA Hiroyuki" <kamezawa.hiroyu@jp.fujitsu.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
The kernel implements readlink of /proc/pid/exe by getting the file from
the first executable VMA. Then the path to the file is reconstructed and
reported as the result.
Because of the VMA walk the code is slightly different on nommu systems.
This patch avoids separate /proc/pid/exe code on nommu systems. Instead of
walking the VMAs to find the first executable file-backed VMA we store a
reference to the exec'd file in the mm_struct.
That reference would prevent the filesystem holding the executable file
from being unmounted even after unmapping the VMAs. So we track the number
of VM_EXECUTABLE VMAs and drop the new reference when the last one is
unmapped. This avoids pinning the mounted filesystem.
[akpm@linux-foundation.org: improve comments]
[yamamoto@valinux.co.jp: fix dup_mmap]
Signed-off-by: Matt Helsley <matthltc@us.ibm.com>
Cc: Oleg Nesterov <oleg@tv-sign.ru>
Cc: David Howells <dhowells@redhat.com>
Cc:"Eric W. Biederman" <ebiederm@xmission.com>
Cc: Christoph Hellwig <hch@lst.de>
Cc: Al Viro <viro@zeniv.linux.org.uk>
Cc: Hugh Dickins <hugh@veritas.com>
Signed-off-by: YAMAMOTO Takashi <yamamoto@valinux.co.jp>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Remove the mem_cgroup member from mm_struct and instead adds an owner.
This approach was suggested by Paul Menage. The advantage of this approach
is that, once the mm->owner is known, using the subsystem id, the cgroup
can be determined. It also allows several control groups that are
virtually grouped by mm_struct, to exist independent of the memory
controller i.e., without adding mem_cgroup's for each controller, to
mm_struct.
A new config option CONFIG_MM_OWNER is added and the memory resource
controller selects this config option.
This patch also adds cgroup callbacks to notify subsystems when mm->owner
changes. The mm_cgroup_changed callback is called with the task_lock() of
the new task held and is called just prior to changing the mm->owner.
I am indebted to Paul Menage for the several reviews of this patchset and
helping me make it lighter and simpler.
This patch was tested on a powerpc box, it was compiled with both the
MM_OWNER config turned on and off.
After the thread group leader exits, it's moved to init_css_state by
cgroup_exit(), thus all future charges from runnings threads would be
redirected to the init_css_set's subsystem.
Signed-off-by: Balbir Singh <balbir@linux.vnet.ibm.com>
Cc: Pavel Emelianov <xemul@openvz.org>
Cc: Hugh Dickins <hugh@veritas.com>
Cc: Sudhir Kumar <skumar@linux.vnet.ibm.com>
Cc: YAMAMOTO Takashi <yamamoto@valinux.co.jp>
Cc: Hirokazu Takahashi <taka@valinux.co.jp>
Cc: David Rientjes <rientjes@google.com>,
Cc: Balbir Singh <balbir@linux.vnet.ibm.com>
Acked-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Acked-by: Pekka Enberg <penberg@cs.helsinki.fi>
Reviewed-by: Paul Menage <menage@google.com>
Cc: Oleg Nesterov <oleg@tv-sign.ru>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
* 'for-linus' of git://git.kernel.org/pub/scm/linux/kernel/git/penberg/slab-2.6:
slub: pack objects denser
slub: Calculate min_objects based on number of processors.
slub: Drop DEFAULT_MAX_ORDER / DEFAULT_MIN_OBJECTS
slub: Simplify any_slab_object checks
slub: Make the order configurable for each slab cache
slub: Drop fallback to page allocator method
slub: Fallback to minimal order during slab page allocation
slub: Update statistics handling for variable order slabs
slub: Add kmem_cache_order_objects struct
slub: for_each_object must be passed the number of objects in a slab
slub: Store max number of objects in the page struct.
slub: Dump list of objects not freed on kmem_cache_close()
slub: free_list() cleanup
slub: improve kmem_cache_destroy() error message
slob: fix bug - when slob allocates "struct kmem_cache", it does not force alignment.
Save some bytes in mm_struct by filling holes
Putting int values together for better packing on 64bit shrinks sizeof(struct
mm_struct) from 776 bytes to 764 bytes.
Signed-off-by: Andi Kleen <ak@suse.de>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Split the inuse field up to be able to store the number of objects in this
page in the page struct as well. Necessary if we want to have pages of
various orders for a slab. Also avoids touching struct kmem_cache cachelines in
__slab_alloc().
Update diagnostic code to check the number of objects and make sure that
the number of objects always stays within the bounds of a 16 bit unsigned
integer.
Signed-off-by: Christoph Lameter <clameter@sgi.com>
Signed-off-by: Pekka Enberg <penberg@cs.helsinki.fi>
Rename Memory Controller to Memory Resource Controller. Reflect the same
changes in the CONFIG definition for the Memory Resource Controller. Group
together the config options for Resource Counters and Memory Resource
Controller.
Signed-off-by: Balbir Singh <balbir@linux.vnet.ibm.com>
Cc: Paul Menage <menage@google.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
This only made sense for the alternate fastpath which was reverted last week.
Mathieu is working on a new version that addresses the fastpath issues but that
new code first needs to go through mm and it is not clear if we need the
unique end pointers with his new scheme.
Reviewed-by: Pekka Enberg <penberg@cs.helsinki.fi>
Signed-off-by: Christoph Lameter <clameter@sgi.com>
We use a NULL pointer on freelists to signal that there are no more objects.
However the NULL pointers of all slabs match in contrast to the pointers to
the real objects which are in different ranges for different slab pages.
Change the end pointer to be a pointer to the first object and set bit 0.
Every slab will then have a different end pointer. This is necessary to ensure
that end markers can be matched to the source slab during cmpxchg_local.
Bring back the use of the mapping field by SLUB since we would otherwise have
to call a relatively expensive function page_address() in __slab_alloc(). Use
of the mapping field allows avoiding a call to page_address() in various other
functions as well.
There is no need to change the page_mapping() function since bit 0 is set on
the mapping as also for anonymous pages. page_mapping(slab_page) will
therefore still return NULL although the mapping field is overloaded.
Signed-off-by: Christoph Lameter <clameter@sgi.com>
Cc: Pekka Enberg <penberg@cs.helsinki.fi>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Basic setup routines, the mm_struct has a pointer to the cgroup that
it belongs to and the the page has a page_cgroup associated with it.
Signed-off-by: Pavel Emelianov <xemul@openvz.org>
Signed-off-by: Balbir Singh <balbir@linux.vnet.ibm.com>
Cc: Paul Menage <menage@google.com>
Cc: Peter Zijlstra <a.p.zijlstra@chello.nl>
Cc: "Eric W. Biederman" <ebiederm@xmission.com>
Cc: Nick Piggin <nickpiggin@yahoo.com.au>
Cc: Kirill Korotaev <dev@sw.ru>
Cc: Herbert Poetzl <herbert@13thfloor.at>
Cc: David Rientjes <rientjes@google.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
include/asm-powerpc/elf.h has 6 entries in ARCH_DLINFO. fs/binfmt_elf.c
has 14 unconditional NEW_AUX_ENT entries and 2 conditional NEW_AUX_ENT
entries. So in the worst case, saved_auxv does not get an AT_NULL entry at
the end.
The saved_auxv array must be terminated with an AT_NULL entry. Make the
size of mm_struct->saved_auxv arch dependend, based on the number of
ARCH_DLINFO entries.
Signed-off-by: Olaf Hering <olh@suse.de>
Cc: Roland McGrath <roland@redhat.com>
Cc: Jakub Jelinek <jakub@redhat.com>
Cc: Richard Henderson <rth@twiddle.net>
Cc: Ivan Kokshaysky <ink@jurassic.park.msu.ru>
Cc: "Luck, Tony" <tony.luck@intel.com>
Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org>
Cc: Paul Mackerras <paulus@samba.org>
Cc: Paul Mundt <lethal@linux-sh.org>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
We need the offset from the page struct during slab_alloc and slab_free. In
both cases we also reference the cacheline of the kmem_cache_cpu structure.
We can therefore move the offset field into the kmem_cache_cpu structure
freeing up 16 bits in the page struct.
Moving the offset allows an allocation from slab_alloc() without touching the
page struct in the hot path.
The only thing left in slab_free() that touches the page struct cacheline for
per cpu freeing is the checking of SlabDebug(page). The next patch deals with
that.
Use the available 16 bits to broaden page->inuse. More than 64k objects per
slab become possible and we can get rid of the checks for that limitation.
No need anymore to shrink the order of slabs if we boot with 2M sized slabs
(slub_min_order=9).
No need anymore to switch off the offset calculation for very large slabs
since the field in the kmem_cache_cpu structure is 32 bits and so the offset
field can now handle slab sizes of up to 8GB.
Signed-off-by: Christoph Lameter <clameter@sgi.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
After moving the lockless_freelist to kmem_cache_cpu we no longer need
page->lockless_freelist. Restructure the use of the struct page fields in
such a way that we never touch the mapping field.
This is turn allows us to remove the special casing of SLUB when determining
the mapping of a page (needed for corner cases of virtual caches machines that
need to flush caches of processors mapping a page).
Signed-off-by: Christoph Lameter <clameter@sgi.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Move the definitions of struct mm_struct and struct vma_area_struct to
include/mm_types.h. This allows to define more function in asm/pgtable.h
and friends with inline assemblies instead of macros. Compile tested on
i386, powerpc, powerpc64, s390-32, s390-64 and x86_64.
[aurelien@aurel32.net: build fix]
Signed-off-by: Martin Schwidefsky <schwidefsky@de.ibm.com>
Signed-off-by: Aurelien Jarno <aurelien@aurel32.net>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
Avoid atomic overhead in slab_alloc and slab_free
SLUB needs to use the slab_lock for the per cpu slabs to synchronize with
potential kfree operations. This patch avoids that need by moving all free
objects onto a lockless_freelist. The regular freelist continues to exist
and will be used to free objects. So while we consume the
lockless_freelist the regular freelist may build up objects.
If we are out of objects on the lockless_freelist then we may check the
regular freelist. If it has objects then we move those over to the
lockless_freelist and do this again. There is a significant savings in
terms of atomic operations that have to be performed.
We can even free directly to the lockless_freelist if we know that we are
running on the same processor. So this speeds up short lived objects.
They may be allocated and freed without taking the slab_lock. This is
particular good for netperf.
In order to maximize the effect of the new faster hotpath we extract the
hottest performance pieces into inlined functions. These are then inlined
into kmem_cache_alloc and kmem_cache_free. So hotpath allocation and
freeing no longer requires a subroutine call within SLUB.
[I am not sure that it is worth doing this because it changes the easy to
read structure of slub just to reduce atomic ops. However, there is
someone out there with a benchmark on 4 way and 8 way processor systems
that seems to show a 5% regression vs. Slab. Seems that the regression is
due to increased atomic operations use vs. SLAB in SLUB). I wonder if
this is applicable or discernable at all in a real workload?]
Signed-off-by: Christoph Lameter <clameter@sgi.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
This is a new slab allocator which was motivated by the complexity of the
existing code in mm/slab.c. It attempts to address a variety of concerns
with the existing implementation.
A. Management of object queues
A particular concern was the complex management of the numerous object
queues in SLAB. SLUB has no such queues. Instead we dedicate a slab for
each allocating CPU and use objects from a slab directly instead of
queueing them up.
B. Storage overhead of object queues
SLAB Object queues exist per node, per CPU. The alien cache queue even
has a queue array that contain a queue for each processor on each
node. For very large systems the number of queues and the number of
objects that may be caught in those queues grows exponentially. On our
systems with 1k nodes / processors we have several gigabytes just tied up
for storing references to objects for those queues This does not include
the objects that could be on those queues. One fears that the whole
memory of the machine could one day be consumed by those queues.
C. SLAB meta data overhead
SLAB has overhead at the beginning of each slab. This means that data
cannot be naturally aligned at the beginning of a slab block. SLUB keeps
all meta data in the corresponding page_struct. Objects can be naturally
aligned in the slab. F.e. a 128 byte object will be aligned at 128 byte
boundaries and can fit tightly into a 4k page with no bytes left over.
SLAB cannot do this.
D. SLAB has a complex cache reaper
SLUB does not need a cache reaper for UP systems. On SMP systems
the per CPU slab may be pushed back into partial list but that
operation is simple and does not require an iteration over a list
of objects. SLAB expires per CPU, shared and alien object queues
during cache reaping which may cause strange hold offs.
E. SLAB has complex NUMA policy layer support
SLUB pushes NUMA policy handling into the page allocator. This means that
allocation is coarser (SLUB does interleave on a page level) but that
situation was also present before 2.6.13. SLABs application of
policies to individual slab objects allocated in SLAB is
certainly a performance concern due to the frequent references to
memory policies which may lead a sequence of objects to come from
one node after another. SLUB will get a slab full of objects
from one node and then will switch to the next.
F. Reduction of the size of partial slab lists
SLAB has per node partial lists. This means that over time a large
number of partial slabs may accumulate on those lists. These can
only be reused if allocator occur on specific nodes. SLUB has a global
pool of partial slabs and will consume slabs from that pool to
decrease fragmentation.
G. Tunables
SLAB has sophisticated tuning abilities for each slab cache. One can
manipulate the queue sizes in detail. However, filling the queues still
requires the uses of the spin lock to check out slabs. SLUB has a global
parameter (min_slab_order) for tuning. Increasing the minimum slab
order can decrease the locking overhead. The bigger the slab order the
less motions of pages between per CPU and partial lists occur and the
better SLUB will be scaling.
G. Slab merging
We often have slab caches with similar parameters. SLUB detects those
on boot up and merges them into the corresponding general caches. This
leads to more effective memory use. About 50% of all caches can
be eliminated through slab merging. This will also decrease
slab fragmentation because partial allocated slabs can be filled
up again. Slab merging can be switched off by specifying
slub_nomerge on boot up.
Note that merging can expose heretofore unknown bugs in the kernel
because corrupted objects may now be placed differently and corrupt
differing neighboring objects. Enable sanity checks to find those.
H. Diagnostics
The current slab diagnostics are difficult to use and require a
recompilation of the kernel. SLUB contains debugging code that
is always available (but is kept out of the hot code paths).
SLUB diagnostics can be enabled via the "slab_debug" option.
Parameters can be specified to select a single or a group of
slab caches for diagnostics. This means that the system is running
with the usual performance and it is much more likely that
race conditions can be reproduced.
I. Resiliency
If basic sanity checks are on then SLUB is capable of detecting
common error conditions and recover as best as possible to allow the
system to continue.
J. Tracing
Tracing can be enabled via the slab_debug=T,<slabcache> option
during boot. SLUB will then protocol all actions on that slabcache
and dump the object contents on free.
K. On demand DMA cache creation.
Generally DMA caches are not needed. If a kmalloc is used with
__GFP_DMA then just create this single slabcache that is needed.
For systems that have no ZONE_DMA requirement the support is
completely eliminated.
L. Performance increase
Some benchmarks have shown speed improvements on kernbench in the
range of 5-10%. The locking overhead of slub is based on the
underlying base allocation size. If we can reliably allocate
larger order pages then it is possible to increase slub
performance much further. The anti-fragmentation patches may
enable further performance increases.
Tested on:
i386 UP + SMP, x86_64 UP + SMP + NUMA emulation, IA64 NUMA + Simulator
SLUB Boot options
slub_nomerge Disable merging of slabs
slub_min_order=x Require a minimum order for slab caches. This
increases the managed chunk size and therefore
reduces meta data and locking overhead.
slub_min_objects=x Mininum objects per slab. Default is 8.
slub_max_order=x Avoid generating slabs larger than order specified.
slub_debug Enable all diagnostics for all caches
slub_debug=<options> Enable selective options for all caches
slub_debug=<o>,<cache> Enable selective options for a certain set of
caches
Available Debug options
F Double Free checking, sanity and resiliency
R Red zoning
P Object / padding poisoning
U Track last free / alloc
T Trace all allocs / frees (only use for individual slabs).
To use SLUB: Apply this patch and then select SLUB as the default slab
allocator.
[hugh@veritas.com: fix an oops-causing locking error]
[akpm@linux-foundation.org: various stupid cleanups and small fixes]
Signed-off-by: Christoph Lameter <clameter@sgi.com>
Signed-off-by: Hugh Dickins <hugh@veritas.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
This moves the definition of struct page from mm.h to its own header file
page-struct.h. This is a prereq to fix SetPageUptodate which is broken on
s390:
#define SetPageUptodate(_page)
do {
struct page *__page = (_page);
if (!test_and_set_bit(PG_uptodate, &__page->flags))
page_test_and_clear_dirty(_page);
} while (0)
_page gets used twice in this macro which can cause subtle bugs. Using
__page for the page_test_and_clear_dirty call doesn't work since it causes
yet another problem with the page_test_and_clear_dirty macro as well.
In order to avoid all these problems caused by macros it seems to be a good
idea to get rid of them and convert them to static inline functions.
Because of header file include order it's necessary to have a seperate
header file for the struct page definition.
Cc: Martin Schwidefsky <schwidefsky@de.ibm.com>
Signed-off-by: Heiko Carstens <heiko.carstens@de.ibm.com>
Cc: Roman Zippel <zippel@linux-m68k.org>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
Kirill A. Shutemov