This patch contains the following cleanups:
- every file should include the headers containing the prototypes for
  its global functions
- make the follosing needlessly global functions static:
  - migrate_to_node()
  - do_mbind()
  - sp_alloc()
  - mpol_rebind_policy()

[akpm@linux-foundation.org: fix uninitialised var warning]
Signed-off-by: NAdrian Bunk <bunk@stusta.de>
Acked-by: NChristoph Lameter <clameter@sgi.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

This patch makes three needlessly global functions static.
Signed-off-by: NAdrian Bunk <bunk@stusta.de>
Cc: Hugh Dickins <hugh@veritas.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

When gather_surplus_pages() fails to allocate enough huge pages to satisfy
the requested reservation, it frees what it did allocate back to the buddy
allocator.  put_page() should be called instead of update_and_free_page()
to ensure that pool counters are updated as appropriate and the page's
refcount is decremented.
Signed-off-by: NAdam Litke <agl@us.ibm.com>
Acked-by: NDave Hansen <haveblue@us.ibm.com>
Cc: David Gibson <hermes@gibson.dropbear.id.au>
Cc: William Lee Irwin III <wli@holomorphy.com>
Cc: Badari Pulavarty <pbadari@us.ibm.com>
Cc: Ken Chen <kenchen@google.com>
Cc: Lee Schermerhorn <lee.schermerhorn@hp.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Anton found a problem with the hugetlb pool allocation when some nodes have
no memory (http://marc.info/?l=linux-mm&m=118133042025995&w=2).  Lee worked
on versions that tried to fix it, but none were accepted.  Christoph has
created a set of patches which allow for GFP_THISNODE allocations to fail
if the node has no memory.

Currently, alloc_fresh_huge_page() returns NULL when it is not able to
allocate a huge page on the current node, as specified by its custom
interleave variable.  The callers of this function, though, assume that a
failure in alloc_fresh_huge_page() indicates no hugepages can be allocated
on the system period.  This might not be the case, for instance, if we have
an uneven NUMA system, and we happen to try to allocate a hugepage on a
node with less memory and fail, while there is still plenty of free memory
on the other nodes.

To correct this, make alloc_fresh_huge_page() search through all online
nodes before deciding no hugepages can be allocated.  Add a helper function
for actually allocating the hugepage.  Use a new global nid iterator to
control which nid to allocate on.

Note: we expect particular semantics for __GFP_THISNODE, which are now
enforced even for memoryless nodes.  That is, there is should be no
fallback to other nodes.  Therefore, we rely on the nid passed into
alloc_pages_node() to be the nid the page comes from.  If this is
incorrect, accounting will break.

Tested on x86 !NUMA, x86 NUMA, x86_64 NUMA and ppc64 NUMA (with 2
memoryless nodes).

Before on the ppc64 box:
Trying to clear the hugetlb pool
Done.       0 free
Trying to resize the pool to 100
Node 0 HugePages_Free:     25
Node 1 HugePages_Free:     75
Node 2 HugePages_Free:      0
Node 3 HugePages_Free:      0
Done. Initially     100 free
Trying to resize the pool to 200
Node 0 HugePages_Free:     50
Node 1 HugePages_Free:    150
Node 2 HugePages_Free:      0
Node 3 HugePages_Free:      0
Done.     200 free

After:
Trying to clear the hugetlb pool
Done.       0 free
Trying to resize the pool to 100
Node 0 HugePages_Free:     50
Node 1 HugePages_Free:     50
Node 2 HugePages_Free:      0
Node 3 HugePages_Free:      0
Done. Initially     100 free
Trying to resize the pool to 200
Node 0 HugePages_Free:    100
Node 1 HugePages_Free:    100
Node 2 HugePages_Free:      0
Node 3 HugePages_Free:      0
Done.     200 free
Signed-off-by: NNishanth Aravamudan <nacc@us.ibm.com>
Acked-by: NChristoph Lameter <clameter@sgi.com>
Cc: Adam Litke <agl@us.ibm.com>
Cc: David Gibson <hermes@gibson.dropbear.id.au>
Cc: Badari Pulavarty <pbadari@us.ibm.com>
Cc: Ken Chen <kenchen@google.com>
Cc: William Lee Irwin III <wli@holomorphy.com>
Cc: Lee Schermerhorn <lee.schermerhorn@hp.com>
Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

When shrinking the size of the hugetlb pool via the nr_hugepages sysctl, we
are careful to keep enough pages around to satisfy reservations.  But the
calculation is flawed for the following scenario:

Action                          Pool Counters (Total, Free, Resv)
======                          =============
Set pool to 1 page              1 1 0
Map 1 page MAP_PRIVATE          1 1 0
Touch the page to fault it in   1 0 0
Set pool to 3 pages             3 2 0
Map 2 pages MAP_SHARED          3 2 2
Set pool to 2 pages             2 1 2 <-- Mistake, should be 3 2 2
Touch the 2 shared pages        2 0 1 <-- Program crashes here

The last touch above will terminate the process due to lack of huge pages.

This patch corrects the calculation so that it factors in pages being used
for private mappings.  Andrew, this is a standalone fix suitable for
mainline.  It is also now corrected in my latest dynamic pool resizing
patchset which I will send out soon.
Signed-off-by: NAdam Litke <agl@us.ibm.com>
Acked-by: NKen Chen <kenchen@google.com>
Cc: David Gibson <david@gibson.dropbear.id.au>
Cc: Badari Pulavarty <pbadari@us.ibm.com>
Cc: William Lee Irwin III <wli@holomorphy.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Support for reading from hugetlbfs files.  libhugetlbfs lets application
text/data to be placed in large pages.  When we do that, oprofile doesn't
work - since libbfd tries to read from it.

This code is very similar to what do_generic_mapping_read() does, but I
can't use it since it has PAGE_CACHE_SIZE assumptions.

[akpm@linux-foundation.org: cleanups, fix leak]
[bunk@stusta.de: make hugetlbfs_read() static]
Signed-off-by: NBadari Pulavarty <pbadari@us.ibm.com>
Acked-by: NWilliam Irwin <bill.irwin@oracle.com>
Tested-by: NNishanth Aravamudan <nacc@us.ibm.com>
Cc: Nick Piggin <nickpiggin@yahoo.com.au>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

For historical reason, expanding ftruncate that increases file size on
hugetlbfs is not allowed due to pages were pre-faulted and lack of fault
handler.  Now that we have demand faulting on hugetlb since 2.6.15, there
is no reason to hold back that limitation.

This will make hugetlbfs behave more like a normal fs.  I'm writing a user
level code that uses hugetlbfs but will fall back to tmpfs if there are no
hugetlb page available in the system.  Having hugetlbfs specific ftruncate
behavior is a bit quirky and I would like to remove that artificial
limitation.

Signed-off-by: <kenchen@google.com>
Acked-by: NWiliam Irwin <wli@holomorphy.com>
Cc: Adam Litke <agl@us.ibm.com>
Cc: David Gibson <david@gibson.dropbear.id.au>
Cc: Nishanth Aravamudan <nacc@us.ibm.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

The maximum size of the huge page pool can be controlled using the overall
size of the hugetlb filesystem (via its 'size' mount option).  However in the
common case the this will not be set as the pool is traditionally fixed in
size at boot time.  In order to maintain the expected semantics, we need to
prevent the pool expanding by default.

This patch introduces a new sysctl controlling dynamic pool resizing.  When
this is enabled the pool will expand beyond its base size up to the size of
the hugetlb filesystem.  It is disabled by default.
Signed-off-by: NAdam Litke <agl@us.ibm.com>
Acked-by: NAndy Whitcroft <apw@shadowen.org>
Acked-by: NDave McCracken <dave.mccracken@oracle.com>
Cc: William Irwin <bill.irwin@oracle.com>
Cc: David Gibson <david@gibson.dropbear.id.au>
Cc: Ken Chen <kenchen@google.com>
Cc: Badari Pulavarty <pbadari@us.ibm.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Shared mappings require special handling because the huge pages needed to
fully populate the VMA must be reserved at mmap time.  If not enough pages are
available when making the reservation, allocate all of the shortfall at once
from the buddy allocator and add the pages directly to the hugetlb pool.  If
they cannot be allocated, then fail the mapping.  The page surplus is
accounted for in the same way as for private mappings; faulted surplus pages
will be freed at unmap time.  Reserved, surplus pages that have not been used
must be freed separately when their reservation has been released.
Signed-off-by: NAdam Litke <agl@us.ibm.com>
Acked-by: NAndy Whitcroft <apw@shadowen.org>
Acked-by: NDave McCracken <dave.mccracken@oracle.com>
Cc: William Irwin <bill.irwin@oracle.com>
Cc: David Gibson <david@gibson.dropbear.id.au>
Cc: Ken Chen <kenchen@google.com>
Cc: Badari Pulavarty <pbadari@us.ibm.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Because we overcommit hugepages for MAP_PRIVATE mappings, it is possible that
the hugetlb pool will be exhausted or completely reserved when a hugepage is
needed to satisfy a page fault.  Before killing the process in this situation,
try to allocate a hugepage directly from the buddy allocator.

The explicitly configured pool size becomes a low watermark.  When dynamically
grown, the allocated huge pages are accounted as a surplus over the watermark.
 As huge pages are freed on a node, surplus pages are released to the buddy
allocator so that the pool will shrink back to the watermark.

Surplus accounting also allows for friendlier explicit pool resizing.  When
shrinking a pool that is fully in-use, increase the surplus so pages will be
returned to the buddy allocator as soon as they are freed.  When growing a
pool that has a surplus, consume the surplus first and then allocate new
pages.
Signed-off-by: NAdam Litke <agl@us.ibm.com>
Signed-off-by: NMel Gorman <mel@csn.ul.ie>
Acked-by: NAndy Whitcroft <apw@shadowen.org>
Acked-by: NDave McCracken <dave.mccracken@oracle.com>
Cc: William Irwin <bill.irwin@oracle.com>
Cc: David Gibson <david@gibson.dropbear.id.au>
Cc: Ken Chen <kenchen@google.com>
Cc: Badari Pulavarty <pbadari@us.ibm.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Dynamic huge page pool resizing.

In most real-world scenarios, configuring the size of the hugetlb pool
correctly is a difficult task.  If too few pages are allocated to the pool,
applications using MAP_SHARED may fail to mmap() a hugepage region and
applications using MAP_PRIVATE may receive SIGBUS.  Isolating too much memory
in the hugetlb pool means it is not available for other uses, especially those
programs not using huge pages.

The obvious answer is to let the hugetlb pool grow and shrink in response to
the runtime demand for huge pages.  The work Mel Gorman has been doing to
establish a memory zone for movable memory allocations makes dynamically
resizing the hugetlb pool reliable within the limits of that zone.  This patch
series implements dynamic pool resizing for private and shared mappings while
being careful to maintain existing semantics.  Please reply with your comments
and feedback; even just to say whether it would be a useful feature to you.
Thanks.

How it works
============

Upon depletion of the hugetlb pool, rather than reporting an error immediately,
first try and allocate the needed huge pages directly from the buddy allocator.
Care must be taken to avoid unbounded growth of the hugetlb pool, so the
hugetlb filesystem quota is used to limit overall pool size.

The real work begins when we decide there is a shortage of huge pages.  What
happens next depends on whether the pages are for a private or shared mapping.
Private mappings are straightforward.  At fault time, if alloc_huge_page()
fails, we allocate a page from the buddy allocator and increment the source
node's surplus_huge_pages counter.  When free_huge_page() is called for a page
on a node with a surplus, the page is freed directly to the buddy allocator
instead of the hugetlb pool.

Because shared mappings require all of the pages to be reserved up front, some
additional work must be done at mmap() to support them.  We determine the
reservation shortage and allocate the required number of pages all at once.
These pages are then added to the hugetlb pool and marked reserved.  Where that
is not possible the mmap() will fail.  As with private mappings, the
appropriate surplus counters are updated.  Since reserved huge pages won't
necessarily be used by the process, we can't be sure that free_huge_page() will
always be called to return surplus pages to the buddy allocator.  To prevent
the huge page pool from bloating, we must free unused surplus pages when their
reservation has ended.

Controlling it
==============

With the entire patch series applied, pool resizing is off by default so unless
specific action is taken, the semantics are unchanged.

To take advantage of the flexibility afforded by this patch series one must
tolerate a change in semantics.  To control hugetlb pool growth, the following
techniques can be employed:

 * A sysctl tunable to enable/disable the feature entirely
 * The size= mount option for hugetlbfs filesystems to limit pool size

Performance
===========

When contiguous memory is readily available, it is expected that the cost of
dynamicly resizing the pool will be small.  This series has been performance
tested with 'stream' to measure this cost.

Stream (http://www.cs.virginia.edu/stream/) was linked with libhugetlbfs to
enable remapping of the text and data/bss segments into huge pages.

Stream with small array
-----------------------
Baseline: 	nr_hugepages = 0, No libhugetlbfs segment remapping
Preallocated:	nr_hugepages = 5, Text and data/bss remapping
Dynamic:	nr_hugepages = 0, Text and data/bss remapping

				Rate (MB/s)
Function	Baseline	Preallocated	Dynamic
Copy:		4695.6266	5942.8371	5982.2287
Scale:		4451.5776	5017.1419	5658.7843
Add:		5815.8849	7927.7827	8119.3552
Triad:		5949.4144	8527.6492	8110.6903

Stream with large array
-----------------------
Baseline: 	nr_hugepages =  0, No libhugetlbfs segment remapping
Preallocated:	nr_hugepages = 67, Text and data/bss remapping
Dynamic:	nr_hugepages =  0, Text and data/bss remapping

				Rate (MB/s)
Function	Baseline	Preallocated	Dynamic
Copy:		2227.8281	2544.2732	2546.4947
Scale:		2136.3208	2430.7294	2421.2074
Add:		2773.1449	4004.0021	3999.4331
Triad:		2748.4502	3777.0109	3773.4970

* All numbers are averages taken from 10 consecutive runs with a maximum
  standard deviation of 1.3 percent noted.

This patch:

Simply move update_and_free_page() so that it can be reused later in this
patch series.  The implementation is not changed.
Signed-off-by: NAdam Litke <agl@us.ibm.com>
Acked-by: NAndy Whitcroft <apw@shadowen.org>
Acked-by: NDave McCracken <dave.mccracken@oracle.com>
Acked-by: NWilliam Irwin <bill.irwin@oracle.com>
Cc: David Gibson <david@gibson.dropbear.id.au>
Cc: Ken Chen <kenchen@google.com>
Cc: Badari Pulavarty <pbadari@us.ibm.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

This patch is to avoid panic when memory hot-add is executed with
sparsemem-vmemmap.  Current vmemmap-sparsemem code doesn't support memory
hot-add.  Vmemmap must be populated when hot-add.  This is for
2.6.23-rc2-mm2.

Todo: # Even if this patch is applied, the message "[xxxx-xxxx] potential
        offnode page_structs" is displayed. To allocate memmap on its node,
        memmap (and pgdat) must be initialized itself like chicken and
        egg relationship.

      # vmemmap_unpopulate will be necessary for followings.
         - For cancel hot-add due to error.
         - For unplug.
Signed-off-by: NYasunori Goto <y-goto@jp.fujitsu.com>
Cc: Andy Whitcroft <apw@shadowen.org>
Cc: Christoph Lameter <clameter@sgi.com>
Cc: Mel Gorman <mel@csn.ul.ie>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Now, arch dependent code around CONFIG_MEMORY_HOTREMOVE is a mess.
This patch cleans up them. This is against 2.6.23-rc6-mm1.

 - fix compile failure on ia64/ CONFIG_MEMORY_HOTPLUG && !CONFIG_MEMORY_HOTREMOVE case.
 - For !CONFIG_MEMORY_HOTREMOVE, add generic no-op remove_memory(),
   which returns -EINVAL.
 - removed remove_pages() only used in powerpc.
 - removed no-op remove_memory() in i386, sh, sparc64, x86_64.

 - only powerpc returns -ENOSYS at memory hot remove(no-op). changes it
   to return -EINVAL.

Note:
Currently, only ia64 supports CONFIG_MEMORY_HOTREMOVE. I welcome other
archs if there are requirements and testers.
Signed-off-by: NKAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

IA64 memory unplug interface.
Signed-off-by: NKAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Logic.
 - set all pages in  [start,end)  as isolated migration-type.
   by this, all free pages in the range will be not-for-use.
 - Migrate all LRU pages in the range.
 - Test all pages in the range's refcnt is zero or not.

Todo:
 - allocate migration destination page from better area.
 - confirm page_count(page)== 0 && PageReserved(page) page is safe to be freed..
 (I don't like this kind of page but..
 - Find out pages which cannot be migrated.
 - more running tests.
 - Use reclaim for unplugging other memory type area.
Signed-off-by: NKAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Signed-off-by: NYasunori Goto <y-goto@jp.fujitsu.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Implement generic chunk-of-pages isolation method by using page grouping ops.

This patch add MIGRATE_ISOLATE to MIGRATE_TYPES. By this
 - MIGRATE_TYPES increases.
 - bitmap for migratetype is enlarged.

pages of MIGRATE_ISOLATE migratetype will not be allocated even if it is free.
By this, you can isolated *freed* pages from users. How-to-free pages is not
a purpose of this patch. You may use reclaim and migrate codes to free pages.

If start_isolate_page_range(start,end) is called,
 - migratetype of the range turns to be MIGRATE_ISOLATE  if
   its type is MIGRATE_MOVABLE. (*) this check can be updated if other
   memory reclaiming works make progress.
 - MIGRATE_ISOLATE is not on migratetype fallback list.
 - All free pages and will-be-freed pages are isolated.
To check all pages in the range are isolated or not,  use test_pages_isolated(),
To cancel isolation, use undo_isolate_page_range().

Changes V6 -> V7
 - removed unnecessary #ifdef

There are HOLES_IN_ZONE handling codes...I'm glad if we can remove them..
Signed-off-by: NYasunori Goto <y-goto@jp.fujitsu.com>
Signed-off-by: NKAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

A clean up patch for "scanning memory resource [start, end)" operation.

Now, find_next_system_ram() function is used in memory hotplug, but this
interface is not easy to use and codes are complicated.

This patch adds walk_memory_resouce(start,len,arg,func) function.
The function 'func' is called per valid memory resouce range in [start,pfn).

[pbadari@us.ibm.com: Error handling in walk_memory_resource()]
Signed-off-by: NKAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Signed-off-by: NBadari Pulavarty <pbadari@us.ibm.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

The statistics patch later needs to know what order a free page is on the free
lists.  Rather than having special knowledge of page_private() when
PageBuddy() is set, this patch places out page_order() in internal.h and adds
a VM_BUG_ON to catch using it on non-PageBuddy pages.
Signed-off-by: NMel Gorman <mel@csn.ul.ie>
Signed-off-by: NChristoph Lameter <clameter@sgi.com>
Acked-by: NAndy Whitcroft <apw@shadowen.org>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

It's a short-lived allocation.

Cc: Christoph Lameter <clameter@sgi.com>
Cc: Mel Gorman <mel@csn.ul.ie>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

We touch a cacheline in the kmem_cache structure for zeroing to get the
size. However, the hot paths in slab_alloc and slab_free do not reference
any other fields in kmem_cache, so we may have to just bring in the
cacheline for this one access.

Add a new field to kmem_cache_cpu that contains the object size. That
cacheline must already be used in the hotpaths. So we save one cacheline
on every slab_alloc if we zero.

We need to update the kmem_cache_cpu object size if an aliasing operation
changes the objsize of an non debug slab.
Signed-off-by: NChristoph Lameter <clameter@sgi.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

The kmem_cache_cpu structures introduced are currently an array placed in the
kmem_cache struct. Meaning the kmem_cache_cpu structures are overwhelmingly
on the wrong node for systems with a higher amount of nodes. These are
performance critical structures since the per node information has
to be touched for every alloc and free in a slab.

In order to place the kmem_cache_cpu structure optimally we put an array
of pointers to kmem_cache_cpu structs in kmem_cache (similar to SLAB).

However, the kmem_cache_cpu structures can now be allocated in a more
intelligent way.

We would like to put per cpu structures for the same cpu but different
slab caches in cachelines together to save space and decrease the cache
footprint. However, the slab allocators itself control only allocations
per node. We set up a simple per cpu array for every processor with
100 per cpu structures which is usually enough to get them all set up right.
If we run out then we fall back to kmalloc_node. This also solves the
bootstrap problem since we do not have to use slab allocator functions
early in boot to get memory for the small per cpu structures.

Pro:
	- NUMA aware placement improves memory performance
	- All global structures in struct kmem_cache become readonly
	- Dense packing of per cpu structures reduces cacheline
	  footprint in SMP and NUMA.
	- Potential avoidance of exclusive cacheline fetches
	  on the free and alloc hotpath since multiple kmem_cache_cpu
	  structures are in one cacheline. This is particularly important
	  for the kmalloc array.

Cons:
	- Additional reference to one read only cacheline (per cpu
	  array of pointers to kmem_cache_cpu) in both slab_alloc()
	  and slab_free().

[akinobu.mita@gmail.com: fix cpu hotplug offline/online path]
Signed-off-by: NChristoph Lameter <clameter@sgi.com>
Cc: "Pekka Enberg" <penberg@cs.helsinki.fi>
Cc: Akinobu Mita <akinobu.mita@gmail.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Set c->node to -1 if we allocate from a debug slab instead for SlabDebug
which requires access the page struct cacheline.
Signed-off-by: NChristoph Lameter <clameter@sgi.com>
Tested-by: NAlexey Dobriyan <adobriyan@sw.ru>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus 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: NChristoph Lameter <clameter@sgi.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus 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: NChristoph Lameter <clameter@sgi.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

A remote free may access the same page struct that also contains the lockless
freelist for the cpu slab. If objects have a short lifetime and are freed by
a different processor then remote frees back to the slab from which we are
currently allocating are frequent. The cacheline with the page struct needs
to be repeately acquired in exclusive mode by both the allocating thread and
the freeing thread. If this is frequent enough then performance will suffer
because of cacheline bouncing.

This patchset puts the lockless_freelist pointer in its own cacheline. In
order to make that happen we introduce a per cpu structure called
kmem_cache_cpu.

Instead of keeping an array of pointers to page structs we now keep an array
to a per cpu structure that--among other things--contains the pointer to the
lockless freelist. The freeing thread can then keep possession of exclusive
access to the page struct cacheline while the allocating thread keeps its
exclusive access to the cacheline containing the per cpu structure.

This works as long as the allocating cpu is able to service its request
from the lockless freelist. If the lockless freelist runs empty then the
allocating thread needs to acquire exclusive access to the cacheline with
the page struct lock the slab.

The allocating thread will then check if new objects were freed to the per
cpu slab. If so it will keep the slab as the cpu slab and continue with the
recently remote freed objects. So the allocating thread can take a series
of just freed remote pages and dish them out again. Ideally allocations
could be just recycling objects in the same slab this way which will lead
to an ideal allocation / remote free pattern.

The number of objects that can be handled in this way is limited by the
capacity of one slab. Increasing slab size via slub_min_objects/
slub_max_order may increase the number of objects and therefore performance.

If the allocating thread runs out of objects and finds that no objects were
put back by the remote processor then it will retrieve a new slab (from the
partial lists or from the page allocator) and start with a whole
new set of objects while the remote thread may still be freeing objects to
the old cpu slab. This may then repeat until the new slab is also exhausted.
If remote freeing has freed objects in the earlier slab then that earlier
slab will now be on the partial freelist and the allocating thread will
pick that slab next for allocation. So the loop is extended. However,
both threads need to take the list_lock to make the swizzling via
the partial list happen.

It is likely that this kind of scheme will keep the objects being passed
around to a small set that can be kept in the cpu caches leading to increased
performance.

More code cleanups become possible:

- Instead of passing a cpu we can now pass a kmem_cache_cpu structure around.
  Allows reducing the number of parameters to various functions.
- Can define a new node_match() function for NUMA to encapsulate locality
  checks.

Effect on allocations:

Cachelines touched before this patch:

	Write:	page cache struct and first cacheline of object

Cachelines touched after this patch:

	Write:	kmem_cache_cpu cacheline and first cacheline of object
	Read: page cache struct (but see later patch that avoids touching
		that cacheline)

The handling when the lockless alloc list runs empty gets to be a bit more
complicated since another cacheline has now to be written to. But that is
halfway out of the hot path.

Effect on freeing:

Cachelines touched before this patch:

	Write: page_struct and first cacheline of object

Cachelines touched after this patch depending on how we free:

  Write(to cpu_slab):	kmem_cache_cpu struct and first cacheline of object
  Write(to other):	page struct and first cacheline of object

  Read(to cpu_slab):	page struct to id slab etc. (but see later patch that
  			avoids touching the page struct on free)
  Read(to other):	cpu local kmem_cache_cpu struct to verify its not
  			the cpu slab.

Summary:

Pro:
	- Distinct cachelines so that concurrent remote frees and local
	  allocs on a cpuslab can occur without cacheline bouncing.
	- Avoids potential bouncing cachelines because of neighboring
	  per cpu pointer updates in kmem_cache's cpu_slab structure since
	  it now grows to a cacheline (Therefore remove the comment
	  that talks about that concern).

Cons:
	- Freeing objects now requires the reading of one additional
	  cacheline. That can be mitigated for some cases by the following
	  patches but its not possible to completely eliminate these
	  references.

	- Memory usage grows slightly.

	The size of each per cpu object is blown up from one word
	(pointing to the page_struct) to one cacheline with various data.
	So this is NR_CPUS*NR_SLABS*L1_BYTES more memory use. Lets say
	NR_SLABS is 100 and a cache line size of 128 then we have just
	increased SLAB metadata requirements by 12.8k per cpu.
	(Another later patch reduces these requirements)
Signed-off-by: NChristoph Lameter <clameter@sgi.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

This patch makes needlessly global code static.
Signed-off-by: NAdrian Bunk <bunk@stusta.de>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

This patch provides fragmentation avoidance statistics via /proc/pagetypeinfo.
 The information is collected only on request so there is no runtime overhead.
 The statistics are in three parts:

The first part prints information on the size of blocks that pages are
being grouped on and looks like

Page block order: 10
Pages per block:  1024

The second part is a more detailed version of /proc/buddyinfo and looks like

Free pages count per migrate type at order       0      1      2      3      4      5      6      7      8      9     10
Node    0, zone      DMA, type    Unmovable      0      0      0      0      0      0      0      0      0      0      0
Node    0, zone      DMA, type  Reclaimable      1      0      0      0      0      0      0      0      0      0      0
Node    0, zone      DMA, type      Movable      0      0      0      0      0      0      0      0      0      0      0
Node    0, zone      DMA, type      Reserve      0      4      4      0      0      0      0      1      0      1      0
Node    0, zone   Normal, type    Unmovable    111      8      4      4      2      3      1      0      0      0      0
Node    0, zone   Normal, type  Reclaimable    293     89      8      0      0      0      0      0      0      0      0
Node    0, zone   Normal, type      Movable      1      6     13      9      7      6      3      0      0      0      0
Node    0, zone   Normal, type      Reserve      0      0      0      0      0      0      0      0      0      0      4

The third part looks like

Number of blocks type     Unmovable  Reclaimable      Movable      Reserve
Node 0, zone      DMA            0            1            2            1
Node 0, zone   Normal            3           17           94            4

To walk the zones within a node with interrupts disabled, walk_zones_in_node()
is introduced and shared between /proc/buddyinfo, /proc/zoneinfo and
/proc/pagetypeinfo to reduce code duplication.  It seems specific to what
vmstat.c requires but could be broken out as a general utility function in
mmzone.c if there were other other potential users.
Signed-off-by: NMel Gorman <mel@csn.ul.ie>
Acked-by: NAndy Whitcroft <apw@shadowen.org>
Acked-by: NChristoph Lameter <clameter@sgi.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Currently mobility grouping works at the MAX_ORDER_NR_PAGES level.  This makes
sense for the majority of users where this is also the huge page size.
However, on platforms like ia64 where the huge page size is runtime
configurable it is desirable to group at a lower order.  On x86_64 and
occasionally on x86, the hugepage size may not always be MAX_ORDER_NR_PAGES.

This patch groups pages together based on the value of HUGETLB_PAGE_ORDER.  It
uses a compile-time constant if possible and a variable where the huge page
size is runtime configurable.

It is assumed that grouping should be done at the lowest sensible order and
that the user would not want to override this.  If this is not true,
page_block order could be forced to a variable initialised via a boot-time
kernel parameter.

One potential issue with this patch is that IA64 now parses hugepagesz with
early_param() instead of __setup().  __setup() is called after the memory
allocator has been initialised and the pageblock bitmaps already setup.  In
tests on one IA64 there did not seem to be any problem with using
early_param() and in fact may be more correct as it guarantees the parameter
is handled before the parsing of hugepages=.
Signed-off-by: NMel Gorman <mel@csn.ul.ie>
Acked-by: NAndy Whitcroft <apw@shadowen.org>
Acked-by: NChristoph Lameter <clameter@sgi.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

move_freepages_block() returns the number of blocks moved.  This value is used
to determine if a block of pages should be stolen for the exclusive use of a
migrate type or not.  However, the value returned is being used correctly.
This patch fixes the calculation to return the number of base pages that have
been moved.

This should be considered a fix to the patch
move-free-pages-between-lists-on-steal.patch

Credit to Andy Whitcroft for spotting the problem.
Signed-off-by: NMel Gorman <mel@csn.ul.ie>
Acked-by: NAndy Whitcroft <apw@shadowen.org>
Acked-by: NChristoph Lameter <clameter@sgi.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Grouping high-order atomic allocations together was intended to allow
bursty users of atomic allocations to work such as e1000 in situations
where their preallocated buffers were depleted.  This did not work in at
least one case with a wireless network adapter needing order-1 allocations
frequently.  To resolve that, the free pages used for min_free_kbytes were
moved to separate contiguous blocks with the patch
bias-the-location-of-pages-freed-for-min_free_kbytes-in-the-same-max_order_nr_pages-blocks.

It is felt that keeping the free pages in the same contiguous blocks should
be sufficient for bursty short-lived high-order atomic allocations to
succeed, maybe even with the e1000.  Even if there is a failure, increasing
the value of min_free_kbytes will free pages as contiguous bloks in
contrast to the standard buddy allocator which makes no attempt to keep the
minimum number of free pages contiguous.

This patch backs out grouping high order atomic allocations together to
determine if it is really needed or not.  If a new report comes in about
high-order atomic allocations failing, the feature can be reintroduced to
determine if it fixes the problem or not.  As a side-effect, this patch
reduces by 1 the number of bits required to track the mobility type of
pages within a MAX_ORDER_NR_PAGES block.
Signed-off-by: NMel Gorman <mel@csn.ul.ie>
Acked-by: NAndy Whitcroft <apw@shadowen.org>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Grouping pages by mobility can be disabled at compile-time. This was
considered undesirable by a number of people. However, in the current stack of
patches, it is not a simple case of just dropping the configurable patch as it
would cause merge conflicts.  This patch backs out the configuration option.
Signed-off-by: NMel Gorman <mel@csn.ul.ie>
Acked-by: NAndy Whitcroft <apw@shadowen.org>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

The standard buddy allocator always favours the smallest block of pages.
The effect of this is that the pages free to satisfy min_free_kbytes tends
to be preserved since boot time at the same location of memory ffor a very
long time and as a contiguous block.  When an administrator sets the
reserve at 16384 at boot time, it tends to be the same MAX_ORDER blocks
that remain free.  This allows the occasional high atomic allocation to
succeed up until the point the blocks are split.  In practice, it is
difficult to split these blocks but when they do split, the benefit of
having min_free_kbytes for contiguous blocks disappears.  Additionally,
increasing min_free_kbytes once the system has been running for some time
has no guarantee of creating contiguous blocks.

On the other hand, CONFIG_PAGE_GROUP_BY_MOBILITY favours splitting large
blocks when there are no free pages of the appropriate type available.  A
side-effect of this is that all blocks in memory tends to be used up and
the contiguous free blocks from boot time are not preserved like in the
vanilla allocator.  This can cause a problem if a new caller is unwilling
to reclaim or does not reclaim for long enough.

A failure scenario was found for a wireless network device allocating
order-1 atomic allocations but the allocations were not intense or frequent
enough for a whole block of pages to be preserved for MIGRATE_HIGHALLOC.
This was reproduced on a desktop by booting with mem=256mb, forcing the
driver to allocate at order-1, running a bittorrent client (downloading a
debian ISO) and building a kernel with -j2.

This patch addresses the problem on the desktop machine booted with
mem=256mb.  It works by setting aside a reserve of MAX_ORDER_NR_PAGES
blocks, the number of which depends on the value of min_free_kbytes.  These
blocks are only fallen back to when there is no other free pages.  Then the
smallest possible page is used just like the normal buddy allocator instead
of the largest possible page to preserve contiguous pages The pages in free
lists in the reserve blocks are never taken for another migrate type.  The
results is that even if min_free_kbytes is set to a low value, contiguous
blocks will be preserved in the MIGRATE_RESERVE blocks.

This works better than the vanilla allocator because if min_free_kbytes is
increased, a new reserve block will be chosen based on the location of
reclaimable pages and the block will free up as contiguous pages.  In the
vanilla allocator, no effort is made to target a block of pages to free as
contiguous pages and min_free_kbytes pages are scattered randomly.

This effect has been observed on the test machine.  min_free_kbytes was set
initially low but it was kept as a contiguous free block within
MIGRATE_RESERVE.  min_free_kbytes was then set to a higher value and over a
period of time, the free blocks were within the reserve and coalescing.
How long it takes to free up depends on how quickly LRU is rotating.
Amusingly, this means that more activity will free the blocks faster.

This mechanism potentially replaces MIGRATE_HIGHALLOC as it may be more
effective than grouping contiguous free pages together.  It all depends on
whether the number of active atomic high allocations exceeds
min_free_kbytes or not.  If the number of active allocations exceeds
min_free_kbytes, it's worth it but maybe in that situation, min_free_kbytes
should be set higher.  Once there are no more reports of allocation
failures, a patch will be submitted that backs out MIGRATE_HIGHALLOC and
see if the reports stay missing.

Credit to Mariusz Kozlowski for discovering the problem, describing the
failure scenario and testing patches and scenarios.

[akpm@linux-foundation.org: cleanups]
Signed-off-by: NMel Gorman <mel@csn.ul.ie>
Acked-by: NAndy Whitcroft <apw@shadowen.org>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

There are problems in the use of SPARSEMEM and pageblock flags that causes
problems on ia64.

The first part of the problem is that units are incorrect in
SECTION_BLOCKFLAGS_BITS computation.  This results in a map_section's
section_mem_map being treated as part of a bitmap which isn't good.  This
was evident with an invalid virtual address when mem_init attempted to free
bootmem pages while relinquishing control from the bootmem allocator.

The second part of the problem occurs because the pageblock flags bitmap is
be located with the mem_section.  The SECTIONS_PER_ROOT computation using
sizeof (mem_section) may not be a power of 2 depending on the size of the
bitmap.  This renders masks and other such things not power of 2 base.
This issue was seen with SPARSEMEM_EXTREME on ia64.  This patch moves the
bitmap outside of mem_section and uses a pointer instead in the
mem_section.  The bitmaps are allocated when the section is being
initialised.

Note that sparse_early_usemap_alloc() does not use alloc_remap() like
sparse_early_mem_map_alloc().  The allocation required for the bitmap on
x86, the only architecture that uses alloc_remap is typically smaller than
a cache line.  alloc_remap() pads out allocations to the cache size which
would be a needless waste.

Credit to Bob Picco for identifying the original problem and effecting a
fix for the SECTION_BLOCKFLAGS_BITS calculation.  Credit to Andy Whitcroft
for devising the best way of allocating the bitmaps only when required for
the section.

[wli@holomorphy.com: warning fix]
Signed-off-by: NBob Picco <bob.picco@hp.com>
Signed-off-by: NAndy Whitcroft <apw@shadowen.org>
Signed-off-by: NMel Gorman <mel@csn.ul.ie>
Cc: "Luck, Tony" <tony.luck@intel.com>
Signed-off-by: NWilliam Irwin <bill.irwin@oracle.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

MIGRATE_RECLAIMABLE allocations tend to be very bursty in nature like when
updatedb starts.  It is likely this will occur in situations where MAX_ORDER
blocks of pages are not free.  This means that updatedb can scatter
MIGRATE_RECLAIMABLE pages throughout the address space.  This patch is more
agressive about stealing blocks of pages for MIGRATE_RECLAIMABLE.
Signed-off-by: NMel Gorman <mel@csn.ul.ie>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

This patch chooses blocks with lower PFNs when placing kernel allocations.
This is particularly important during fallback in low memory situations to
stop unmovable pages being placed throughout the entire address space.
Signed-off-by: NMel Gorman <mel@csn.ul.ie>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Grouping pages by mobility can only successfully operate when there are more
MAX_ORDER_NR_PAGES areas than mobility types.  When there are insufficient
areas, fallbacks cannot be avoided.  This has noticeable performance impacts
on machines with small amounts of memory in comparison to MAX_ORDER_NR_PAGES.
For example, on IA64 with a configuration including huge pages spans 1GiB with
MAX_ORDER_NR_PAGES so would need at least 4GiB of RAM before grouping pages by
mobility would be useful.  In comparison, an x86 would need 16MB.

This patch checks the size of vm_total_pages in build_all_zonelists(). If
there are not enough areas,  mobility is effectivly disabled by considering
all allocations as the same type (UNMOVABLE).  This is achived via a
__read_mostly flag.

With this patch, performance is comparable to disabling grouping pages
by mobility at compile-time on a test machine with insufficient memory.
With this patch, it is reasonable to get rid of grouping pages by mobility
a compile-time option.
Signed-off-by: NMel Gorman <mel@csn.ul.ie>
Acked-by: NAndy Whitcroft <apw@shadowen.org>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

In rare cases, the kernel needs to allocate a high-order block of pages
without sleeping.  For example, this is the case with e1000 cards configured
to use jumbo frames.  Migrating or reclaiming pages in this situation is not
an option.

This patch groups these allocations together as much as possible by adding a
new MIGRATE_TYPE.  The MIGRATE_HIGHATOMIC type are exactly what they sound
like.  Care is taken that pages of other migrate types do not use the same
blocks as high-order atomic allocations.
Signed-off-by: NMel Gorman <mel@csn.ul.ie>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

This patch marks a number of allocations that are either short-lived such as
network buffers or are reclaimable such as inode allocations.  When something
like updatedb is called, long-lived and unmovable kernel allocations tend to
be spread throughout the address space which increases fragmentation.

This patch groups these allocations together as much as possible by adding a
new MIGRATE_TYPE.  The MIGRATE_RECLAIMABLE type is for allocations that can be
reclaimed on demand, but not moved.  i.e.  they can be migrated by deleting
them and re-reading the information from elsewhere.
Signed-off-by: NMel Gorman <mel@csn.ul.ie>
Cc: Andy Whitcroft <apw@shadowen.org>
Cc: Christoph Lameter <clameter@sgi.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

When a fallback occurs, there will be free pages for one allocation type
stored on the list for another.  When a large steal occurs, this patch will
move all the free pages within one list to the other.

[y-goto@jp.fujitsu.com: fix BUG_ON check at move_freepages()]
[apw@shadowen.org: Move to using pfn_valid_within()]
Signed-off-by: NMel Gorman <mel@csn.ul.ie>
Cc: Christoph Lameter <clameter@engr.sgi.com>
Signed-off-by: NYasunori Goto <y-goto@jp.fujitsu.com>
Cc: Bjorn Helgaas <bjorn.helgaas@hp.com>
Signed-off-by: NAndy Whitcroft <andyw@uk.ibm.com>
Cc: Bob Picco <bob.picco@hp.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Per-cpu pages can accidentally cause fragmentation because they are free, but
pinned pages in an otherwise contiguous block.  When this patch is applied,
the per-cpu caches are drained after the direct-reclaim is entered if the
requested order is greater than 0.  It simply reuses the code used by suspend
and hotplug.
Signed-off-by: NMel Gorman <mel@csn.ul.ie>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>