This adds support for the Analog Devices Blackfin processor architecture, and
currently supports the BF533, BF532, BF531, BF537, BF536, BF534, and BF561
(Dual Core) devices, with a variety of development platforms including those
avaliable from Analog Devices (BF533-EZKit, BF533-STAMP, BF537-STAMP,
BF561-EZKIT), and Bluetechnix!  Tinyboards.

The Blackfin architecture was jointly developed by Intel and Analog Devices
Inc.  (ADI) as the Micro Signal Architecture (MSA) core and introduced it in
December of 2000.  Since then ADI has put this core into its Blackfin
processor family of devices.  The Blackfin core has the advantages of a clean,
orthogonal,RISC-like microprocessor instruction set.  It combines a dual-MAC
(Multiply/Accumulate), state-of-the-art signal processing engine and
single-instruction, multiple-data (SIMD) multimedia capabilities into a single
instruction-set architecture.

The Blackfin architecture, including the instruction set, is described by the
ADSP-BF53x/BF56x Blackfin Processor Programming Reference
http://blackfin.uclinux.org/gf/download/frsrelease/29/2549/Blackfin_PRM.pdf

The Blackfin processor is already supported by major releases of gcc, and
there are binary and source rpms/tarballs for many architectures at:
http://blackfin.uclinux.org/gf/project/toolchain/frs There is complete
documentation, including "getting started" guides available at:
http://docs.blackfin.uclinux.org/ which provides links to the sources and
patches you will need in order to set up a cross-compiling environment for
bfin-linux-uclibc

This patch, as well as the other patches (toolchain, distribution,
uClibc) are actively supported by Analog Devices Inc, at:
http://blackfin.uclinux.org/

We have tested this on LTP, and our test plan (including pass/fails) can
be found at:
http://docs.blackfin.uclinux.org/doku.php?id=testing_the_linux_kernel

[m.kozlowski@tuxland.pl: balance parenthesis in blackfin header files]
Signed-off-by: NBryan Wu <bryan.wu@analog.com>
Signed-off-by: NMariusz Kozlowski <m.kozlowski@tuxland.pl>
Signed-off-by: NAubrey Li <aubrey.li@analog.com>
Signed-off-by: NJie Zhang <jie.zhang@analog.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Address spaces contain an allocation flag that specifies restriction on the
zone for pages placed in the mapping.  I.e.  some device may require pages
to be allocated from a DMA zone.  Block devices may not be able to use
pages from HIGHMEM.

Memory policies and the common use of page migration works only on the
highest zone.  If the address space does not allow allocation from the
highest zone then the pages in the address space are not migratable simply
because we can only allocate memory for a specified node if we allow
allocation for the highest zone on each node.
Acked-by: NHugh Dickins <hugh@veritas.com>
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>

There is no user remaining and I have never seen any use of that flag.
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>

SLAB_CTOR atomic is never used which is no surprise since I cannot imagine
that one would want to do something serious in a constructor or destructor.
 In particular given that the slab allocators run with interrupts disabled.
 Actions in constructors and destructors are by their nature very limited
and usually do not go beyond initializing variables and list operations.

(The i386 pgd ctor and dtors do take a spinlock in constructor and
destructor.....  I think that is the furthest we go at this point.)

There is no flag passed to the destructor so removing SLAB_CTOR_ATOMIC also
establishes a certain symmetry.
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>

I have never seen a use of SLAB_DEBUG_INITIAL.  It is only supported by
SLAB.

I think its purpose was to have a callback after an object has been freed
to verify that the state is the constructor state again?  The callback is
performed before each freeing of an object.

I would think that it is much easier to check the object state manually
before the free.  That also places the check near the code object
manipulation of the object.

Also the SLAB_DEBUG_INITIAL callback is only performed if the kernel was
compiled with SLAB debugging on.  If there would be code in a constructor
handling SLAB_DEBUG_INITIAL then it would have to be conditional on
SLAB_DEBUG otherwise it would just be dead code.  But there is no such code
in the kernel.  I think SLUB_DEBUG_INITIAL is too problematic to make real
use of, difficult to understand and there are easier ways to accomplish the
same effect (i.e.  add debug code before kfree).

There is a related flag SLAB_CTOR_VERIFY that is frequently checked to be
clear in fs inode caches.  Remove the pointless checks (they would even be
pointless without removeal of SLAB_DEBUG_INITIAL) from the fs constructors.

This is the last slab flag that SLUB did not support.  Remove the check for
unimplemented flags from SLUB.
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 provides a new macro

KMEM_CACHE(<struct>, <flags>)

to simplify slab creation. KMEM_CACHE creates a slab with the name of the
struct, with the size of the struct and with the alignment of the struct.
Additional slab flags may be specified if necessary.

Example

struct test_slab {
	int a,b,c;
	struct list_head;
} __cacheline_aligned_in_smp;

test_slab_cache = KMEM_CACHE(test_slab, SLAB_PANIC)

will create a new slab named "test_slab" of the size sizeof(struct
test_slab) and aligned to the alignment of test slab.  If it fails then we
panic.
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 was recently posted to lkml and acked by Pekka.

The flag SLAB_MUST_HWCACHE_ALIGN is

1. Never checked by SLAB at all.

2. A duplicate of SLAB_HWCACHE_ALIGN for SLUB

3. Fulfills the role of SLAB_HWCACHE_ALIGN for SLOB.

The only remaining use is in sparc64 and ppc64 and their use there
reflects some earlier role that the slab flag once may have had. If
its specified then SLAB_HWCACHE_ALIGN is also specified.

The flag is confusing, inconsistent and has no purpose.

Remove it.
Acked-by: NPekka Enberg <penberg@cs.helsinki.fi>
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>

Remove duplicate work in kill_bdev().

It currently invalidates and then truncates the bdev's mapping.
invalidate_mapping_pages() will opportunistically remove pages from the
mapping.  And truncate_inode_pages() will forcefully remove all pages.

The only thing truncate doesn't do is flush the bh lrus.  So do that
explicitly.  This avoids (very unlikely) but possible invalid lookup
results if the same bdev is quickly re-issued.

It also will prevent extreme kernel latencies which are observed when
blockdevs which have a large amount of pagecache are unmounted, by avoiding
invalidate_mapping_pages() on that path.  invalidate_mapping_pages() has no
cond_resched (it can be called under spinlock), whereas truncate_inode_pages()
has one.

[akpm@linux-foundation.org: restore nrpages==0 optimisation]
Signed-off-by: NPeter Zijlstra <a.p.zijlstra@chello.nl>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Remove the destroy_dirty_buffers argument from invalidate_bdev(), it hasn't
been used in 6 years (so akpm says).

find * -name \*.[ch] | xargs grep -l invalidate_bdev |
while read file; do
	quilt add $file;
	sed -ie 's/invalidate_bdev(\([^,]*\),[^)]*)/invalidate_bdev(\1)/g' $file;
done
Signed-off-by: NPeter Zijlstra <a.p.zijlstra@chello.nl>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

I ported this to sparc64 as per the patch below, tested on UP SunBlade1500 and
24 cpu Niagara T1000.
Signed-off-by: NDavid S. Miller <davem@davemloft.net>
Signed-off-by: NChristoph Lameter <clameter@sgi.com>
Cc: "David S. Miller" <davem@davemloft.net>
Cc: Andi Kleen <ak@suse.de>
Cc: "Luck, Tony" <tony.luck@intel.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>

On x86_64 this cuts allocation overhead for page table pages down to a
fraction (kernel compile / editing load.  TSC based measurement of times spend
in each function):

no quicklist

pte_alloc               1569048 4.3s(401ns/2.7us/179.7us)
pmd_alloc                780988 2.1s(337ns/2.7us/86.1us)
pud_alloc                780072 2.2s(424ns/2.8us/300.6us)
pgd_alloc                260022 1s(920ns/4us/263.1us)

quicklist:

pte_alloc                452436 573.4ms(8ns/1.3us/121.1us)
pmd_alloc                196204 174.5ms(7ns/889ns/46.1us)
pud_alloc                195688 172.4ms(7ns/881ns/151.3us)
pgd_alloc                 65228 9.8ms(8ns/150ns/6.1us)

pgd allocations are the most complex and there we see the most dramatic
improvement (may be we can cut down the amount of pgds cached somewhat?).  But
even the pte allocations still see a doubling of performance.

1. Proven code from the IA64 arch.

	The method used here has been fine tuned for years and
	is NUMA aware. It is based on the knowledge that accesses
	to page table pages are sparse in nature. Taking a page
	off the freelists instead of allocating a zeroed pages
	allows a reduction of number of cachelines touched
	in addition to getting rid of the slab overhead. So
	performance improves. This is particularly useful if pgds
	contain standard mappings. We can save on the teardown
	and setup of such a page if we have some on the quicklists.
	This includes avoiding lists operations that are otherwise
	necessary on alloc and free to track pgds.

2. Light weight alternative to use slab to manage page size pages

	Slab overhead is significant and even page allocator use
	is pretty heavy weight. The use of a per cpu quicklist
	means that we touch only two cachelines for an allocation.
	There is no need to access the page_struct (unless arch code
	needs to fiddle around with it). So the fast past just
	means bringing in one cacheline at the beginning of the
	page. That same cacheline may then be used to store the
	page table entry. Or a second cacheline may be used
	if the page table entry is not in the first cacheline of
	the page. The current code will zero the page which means
	touching 32 cachelines (assuming 128 byte). We get down
	from 32 to 2 cachelines in the fast path.

3. x86_64 gets lightweight page table page management.

	This will allow x86_64 arch code to faster repopulate pgds
	and other page table entries. The list operations for pgds
	are reduced in the same way as for i386 to the point where
	a pgd is allocated from the page allocator and when it is
	freed back to the page allocator. A pgd can pass through
	the quicklists without having to be reinitialized.

64 Consolidation of code from multiple arches

	So far arches have their own implementation of quicklist
	management. This patch moves that feature into the core allowing
	an easier maintenance and consistent management of quicklists.

Page table pages have the characteristics that they are typically zero or in a
known state when they are freed.  This is usually the exactly same state as
needed after allocation.  So it makes sense to build a list of freed page
table pages and then consume the pages already in use first.  Those pages have
already been initialized correctly (thus no need to zero them) and are likely
already cached in such a way that the MMU can use them most effectively.  Page
table pages are used in a sparse way so zeroing them on allocation is not too
useful.

Such an implementation already exits for ia64.  Howver, that implementation
did not support constructors and destructors as needed by i386 / x86_64.  It
also only supported a single quicklist.  The implementation here has
constructor and destructor support as well as the ability for an arch to
specify how many quicklists are needed.

Quicklists are defined by an arch defining CONFIG_QUICKLIST.  If more than one
quicklist is necessary then we can define NR_QUICK for additional lists.  F.e.
 i386 needs two and thus has

config NR_QUICK
	int
	default 2

If an arch has requested quicklist support then pages can be allocated
from the quicklist (or from the page allocator if the quicklist is
empty) via:

quicklist_alloc(<quicklist-nr>, <gfpflags>, <constructor>)

Page table pages can be freed using:

quicklist_free(<quicklist-nr>, <destructor>, <page>)

Pages must have a definite state after allocation and before
they are freed. If no constructor is specified then pages
will be zeroed on allocation and must be zeroed before they are
freed.

If a constructor is used then the constructor will establish
a definite page state. F.e. the i386 and x86_64 pgd constructors
establish certain mappings.

Constructors and destructors can also be used to track the pages.
i386 and x86_64 use a list of pgds in order to be able to dynamically
update standard mappings.
Signed-off-by: NChristoph Lameter <clameter@sgi.com>
Cc: "David S. Miller" <davem@davemloft.net>
Cc: Andi Kleen <ak@suse.de>
Cc: "Luck, Tony" <tony.luck@intel.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>

If slab tracking is on then build a list of full slabs so that we can verify
the integrity of all slabs and are also able to built list of alloc/free
callers.
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>

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 patch adds PageTail(page) and PageHead(page) to check if a page is the
head or the tail of a compound page.  This is done by masking the two bits
describing the state of a compound page and then comparing them.  So one
comparision and a branch instead of two bit checks and two branches.
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

If we add a new flag so that we can distinguish between the first page and the
tail pages then we can avoid to use page->private in the first page.
page->private == page for the first page, so there is no real information in
there.

Freeing up page->private makes the use of compound pages more transparent.
They become more usable like real pages.  Right now we have to be careful f.e.
 if we are going beyond PAGE_SIZE allocations in the slab on i386 because we
can then no longer use the private field.  This is one of the issues that
cause us not to support debugging for page size slabs in SLAB.

Having page->private available for SLUB would allow more meta information in
the page struct.  I can probably avoid the 16 bit ints that I have in there
right now.

Also if page->private is available then a compound page may be equipped with
buffer heads.  This may free up the way for filesystems to support larger
blocks than page size.

We add PageTail as an alias of PageReclaim.  Compound pages cannot currently
be reclaimed.  Because of the alias one needs to check PageCompound first.

The RFC for the this approach was discussed at
http://marc.info/?t=117574302800001&r=1&w=2

[nacc@us.ibm.com: fix hugetlbfs]
Signed-off-by: NChristoph Lameter <clameter@sgi.com>
Signed-off-by: NNishanth Aravamudan <nacc@us.ibm.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Makes SLUB behave like SLAB in this area to avoid issues....

Throw a stack dump to alert people.

At some point the behavior should be switched back.  NULL is no memory as
far as I can tell and if the use asked for 0 bytes then he need to get no
memory.
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 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: NChristoph Lameter <clameter@sgi.com>
Signed-off-by: NHugh Dickins <hugh@veritas.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

i386 uses kmalloc to allocate the threadinfo structure assuming that the
allocations result in a page sized aligned allocation.  That has worked so
far because SLAB exempts page sized slabs from debugging and aligns them in
special ways that goes beyond the restrictions imposed by
KMALLOC_ARCH_MINALIGN valid for other slabs in the kmalloc array.

SLUB also works fine without debugging since page sized allocations neatly
align at page boundaries.  However, if debugging is switched on then SLUB
will extend the slab with debug information.  The resulting slab is not
longer of page size.  It will only be aligned following the requirements
imposed by KMALLOC_ARCH_MINALIGN.  As a result the threadinfo structure may
not be page aligned which makes i386 fail to boot with SLUB debug on.

Replace the calls to kmalloc with calls into the page allocator.

An alternate solution may be to create a custom slab cache where the
alignment is set to PAGE_SIZE.  That would allow slub debugging to be
applied to the threadinfo structure.
Signed-off-by: NChristoph Lameter <clameter@sgi.com>
Cc: William Lee Irwin III <wli@holomorphy.com>
Cc: Andi Kleen <ak@suse.de>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Rename file_ra_state.prev_page to prev_index and file_ra_state.offset to
prev_offset.  Also update of prev_index in do_generic_mapping_read() is now
moved close to the update of prev_offset.

[wfg@mail.ustc.edu.cn: fix it]
Signed-off-by: NJan Kara <jack@suse.cz>
Cc: Nick Piggin <nickpiggin@yahoo.com.au>
Cc: WU Fengguang <wfg@mail.ustc.edu.cn>
Signed-off-by: NFengguang Wu <wfg@mail.ustc.edu.cn>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Introduce ra.offset and store in it an offset where the previous read
ended.  This way we can detect whether reads are really sequential (and
thus we should not mark the page as accessed repeatedly) or whether they
are random and just happen to be in the same page (and the page should
really be marked accessed again).
Signed-off-by: NJan Kara <jack@suse.cz>
Acked-by: NNick Piggin <nickpiggin@yahoo.com.au>
Cc: WU Fengguang <wfg@mail.ustc.edu.cn>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Adds /proc/pid/clear_refs.  When any non-zero number is written to this file,
pte_mkold() and ClearPageReferenced() is called for each pte and its
corresponding page, respectively, in that task's VMAs.  This file is only
writable by the user who owns the task.

It is now possible to measure _approximately_ how much memory a task is using
by clearing the reference bits with

	echo 1 > /proc/pid/clear_refs

and checking the reference count for each VMA from the /proc/pid/smaps output
at a measured time interval.  For example, to observe the approximate change
in memory footprint for a task, write a script that clears the references
(echo 1 > /proc/pid/clear_refs), sleeps, and then greps for Pgs_Referenced and
extracts the size in kB.  Add the sizes for each VMA together for the total
referenced footprint.  Moments later, repeat the process and observe the
difference.

For example, using an efficient Mozilla:

	accumulated time		referenced memory
	----------------		-----------------
		 0 s				 408 kB
		 1 s				 408 kB
		 2 s				 556 kB
		 3 s				1028 kB
		 4 s				 872 kB
		 5 s				1956 kB
		 6 s				 416 kB
		 7 s				1560 kB
		 8 s				2336 kB
		 9 s				1044 kB
		10 s				 416 kB

This is a valuable tool to get an approximate measurement of the memory
footprint for a task.

Cc: Hugh Dickins <hugh@veritas.com>
Cc: Paul Mundt <lethal@linux-sh.org>
Cc: Christoph Lameter <clameter@sgi.com>
Signed-off-by: NDavid Rientjes <rientjes@google.com>
[akpm@linux-foundation.org: build fixes]
[mpm@selenic.com: rename for_each_pmd]
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

If you actually clear the bit, you need to:

+         pte_update_defer(vma->vm_mm, addr, ptep);

The reason is, when updating PTEs, the hypervisor must be notified.  Using
atomic operations to do this is fine for all hypervisors I am aware of.
However, for hypervisors which shadow page tables, if these PTE
modifications are not trapped, you need a post-modification call to fulfill
the update of the shadow page table.
Acked-by: NZachary Amsden <zach@vmware.com>
Cc: Hugh Dickins <hugh@veritas.com>
Signed-off-by: NDavid Rientjes <rientjes@google.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Add ptep_test_and_clear_{dirty,young} to i386.  They advertise that they
have it and there is at least one place where it needs to be called without
the page table lock: to clear the accessed bit on write to
/proc/pid/clear_refs.

ptep_clear_flush_{dirty,young} are updated to use the new functions.  The
overall net effect to current users of ptep_clear_flush_{dirty,young} is
that we introduce an additional branch.

Cc: Hugh Dickins <hugh@veritas.com>
Cc: Ingo Molnar <mingo@redhat.com>
Signed-off-by: NDavid Rientjes <rientjes@google.com>
Cc: Andi Kleen <ak@suse.de>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Introduce a macro for suppressing gcc from generating a warning about a
probable uninitialized state of a variable.

Example:

-	spinlock_t *ptl;
+	spinlock_t *uninitialized_var(ptl);

Not a happy solution, but those warnings are obnoxious.

- Using the usual pointlessly-set-it-to-zero approach wastes several
  bytes of text.

- Using a macro means we can (hopefully) do something else if gcc changes
  cause the `x = x' hack to stop working

- Using a macro means that people who are worried about hiding true bugs
  can easily turn it off.
Signed-off-by: NBorislav Petkov <bbpetkov@yahoo.de>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Generally we work under the assumption that memory the mem_map array is
contigious and valid out to MAX_ORDER_NR_PAGES block of pages, ie.  that if we
have validated any page within this MAX_ORDER_NR_PAGES block we need not check
any other.  This is not true when CONFIG_HOLES_IN_ZONE is set and we must
check each and every reference we make from a pfn.

Add a pfn_valid_within() helper which should be used when scanning pages
within a MAX_ORDER_NR_PAGES block when we have already checked the validility
of the block normally with pfn_valid().  This can then be optimised away when
we do not have holes within a MAX_ORDER_NR_PAGES block of pages.
Signed-off-by: NAndy Whitcroft <apw@shadowen.org>
Acked-by: NMel Gorman <mel@csn.ul.ie>
Acked-by: NBob Picco <bob.picco@hp.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Add proper prototypes in include/linux/slab.h.
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>

Architectures that don't support DMA can say so by adding a config NO_DMA
to their Kconfig file.  This will prevent compilation of some dma specific
driver code.  Also dma-mapping-broken.h isn't needed anymore on at least
s390.  This avoids compilation and linking of otherwise dead/broken code.

Other architectures that include dma-mapping-broken.h are arm26, h8300,
m68k, m68knommu and v850.  If these could be converted as well we could get
rid of the header file.
Signed-off-by: NHeiko Carstens <heiko.carstens@de.ibm.com>
"John W. Linville" <linville@tuxdriver.com>
Cc: Kyle McMartin <kyle@parisc-linux.org>
Cc: <James.Bottomley@SteelEye.com>
Cc: Tejun Heo <htejun@gmail.com>
Cc: Jeff Garzik <jeff@garzik.org>
Cc: Martin Schwidefsky <schwidefsky@de.ibm.com>
Cc: <geert@linux-m68k.org>
Cc: <zippel@linux-m68k.org>
Cc: <spyro@f2s.com>
Cc: <uclinux-v850@lsi.nec.co.jp>
Cc: <ysato@users.sourceforge.jp>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Ensure pages are uptodate after returning from read_cache_page, which allows
us to cut out most of the filesystem-internal PageUptodate calls.

I didn't have a great look down the call chains, but this appears to fixes 7
possible use-before uptodate in hfs, 2 in hfsplus, 1 in jfs, a few in
ecryptfs, 1 in jffs2, and a possible cleared data overwritten with readpage in
block2mtd.  All depending on whether the filler is async and/or can return
with a !uptodate page.
Signed-off-by: NNick Piggin <npiggin@suse.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>

Minimum gcc version is 3.2 now.  However, with likely profiling, even
modern gcc versions cannot always eliminate the call.

Replace the placeholder functions with the more conventional empty static
inlines, which should be optimal for everyone.
Signed-off-by: NNick Piggin <npiggin@suse.de>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Add a proper prototype for hugetlb_get_unmapped_area() in
include/linux/hugetlb.h.
Signed-off-by: NAdrian Bunk <bunk@stusta.de>
Acked-by: NWilliam Irwin <wli@holomorphy.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Add a new mm function apply_to_page_range() which applies a given function to
every pte in a given virtual address range in a given mm structure.  This is a
generic alternative to cut-and-pasting the Linux idiomatic pagetable walking
code in every place that a sequence of PTEs must be accessed.

Although this interface is intended to be useful in a wide range of
situations, it is currently used specifically by several Xen subsystems, for
example: to ensure that pagetables have been allocated for a virtual address
range, and to construct batched special pagetable update requests to map I/O
memory (in ioremap()).

[akpm@linux-foundation.org: fix warning, unpleasantly]
Signed-off-by: NIan Pratt <ian.pratt@xensource.com>
Signed-off-by: NChristian Limpach <Christian.Limpach@cl.cam.ac.uk>
Signed-off-by: NChris Wright <chrisw@sous-sol.org>
Signed-off-by: NJeremy Fitzhardinge <jeremy@xensource.com>
Cc: Christoph Lameter <clameter@sgi.com>
Cc: Matt Mackall <mpm@waste.org>
Acked-by: NIngo Molnar <mingo@elte.hu>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

At present, the serial core always allows setserial in userspace to change the
port address, irq and base clock of any serial port.  That makes sense for
legacy ISA ports, but not for (say) embedded ns16550 compatible serial ports
at peculiar addresses.  In these cases, the kernel code configuring the ports
must know exactly where they are, and their clocking arrangements (which can
be unusual on embedded boards).  It doesn't make sense for userspace to change
these settings.

Therefore, this patch defines a UPF_FIXED_PORT flag for the uart_port
structure.  If this flag is set when the serial port is configured, any
attempts to alter the port's type, io address, irq or base clock with
setserial are ignored.

In addition this patch uses the new flag for on-chip serial ports probed in
arch/powerpc/kernel/legacy_serial.c, and for other hard-wired serial ports
probed by drivers/serial/of_serial.c.
Signed-off-by: NDavid Gibson <dwg@au1.ibm.com>
Cc: Russell King <rmk@arm.linux.org.uk>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Add support for the integrated serial ports of the MIPS RM9122 processor
and its relatives.

The patch also does some whitespace cleanup.

[akpm@linux-foundation.org: cleanups]
Signed-off-by: NThomas Koeller <thomas.koeller@baslerweb.com>
Cc: Ralf Baechle <ralf@linux-mips.org>
Cc: Russell King <rmk@arm.linux.org.uk>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Serial driver patch for the PMC-Sierra MSP71xx devices.

There are three different fixes:

1 Fix for DesignWare APB THRE errata: In brief, this is a non-standard
  16550 in that the THRE interrupt will not re-assert itself simply by
  disabling and re-enabling the THRI bit in the IER, it is only re-enabled
  if a character is actually sent out.

  It appears that the "8250-uart-backup-timer.patch" in the "mm" tree
  also fixes it so we have dropped our initial workaround.  This patch now
  needs to be applied on top of that "mm" patch.

2 Fix for Busy Detect on LCR write: The DesignWare APB UART has a feature
  which causes a new Busy Detect interrupt to be generated if it's busy
  when the LCR is written.  This fix saves the value of the LCR and
  rewrites it after clearing the interrupt.

3 Workaround for interrupt/data concurrency issue: The SoC needs to
  ensure that writes that can cause interrupts to be cleared reach the UART
  before returning from the ISR.  This fix reads a non-destructive register
  on the UART so the read transaction completion ensures the previously
  queued write transaction has also completed.
Signed-off-by: NMarc St-Jean <Marc_St-Jean@pmc-sierra.com>
Cc: Russell King <rmk@arm.linux.org.uk>
Cc: Ralf Baechle <ralf@linux-mips.org>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

PCI drivers have the new_id file in sysfs which allows new IDs to be added
at runtime.  The advantage is to avoid re-compilation of a driver that
works for a new device, but it's ID table doesn't contain the new device.
This mechanism is only meant for testing, after the driver has been tested
successfully, the ID should be added in source code so that new revisions
of the kernel automatically detect the device.

The implementation follows the PCI implementation. The interface is documented
in Documentation/pcmcia/driver.txt. Computations should be done in userspace,
so the sysfs string contains the raw structure members for matching.
Signed-off-by: NBernhard Walle <bwalle@suse.de>
Cc: Dominik Brodowski <linux@dominikbrodowski.net>
Cc: Greg KH <greg@kroah.com>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

This introduce krealloc() that reallocates memory while keeping the contents
unchanged.  The allocator avoids reallocation if the new size fits the
currently used cache.  I also added a simple non-optimized version for
mm/slob.c for compatibility.

[akpm@linux-foundation.org: fix warnings]
Acked-by: NJosef Sipek <jsipek@fsl.cs.sunysb.edu>
Acked-by: NMatt Mackall <mpm@selenic.com>
Acked-by: NChristoph Lameter <clameter@sgi.com>
Signed-off-by: NPekka Enberg <penberg@cs.helsinki.fi>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

This was broken.  It adds complexity, for no good reason.  Rather than
separate __pa() and __pa_symbol(), we should deprecate __pa_symbol(),
and preferably __pa() too - and just use "virt_to_phys()" instead, which
is more readable and has nicer semantics.

However, right now, just undo the separation, and make __pa_symbol() be
the exact same as __pa().  That fixes the bugs this patch introduced,
and we can do the fairly obvious cleanups later.

Do the new __phys_addr() function (which is now the actual workhorse for
the unified __pa()/__pa_symbol()) as a real external function, that way
all the potential issues with compile/link-time optimizations of
constant symbol addresses go away, and we can also, if we choose to, add
more sanity-checking of the argument.

Cc: Eric W. Biederman <ebiederm@xmission.com>
Cc: Vivek Goyal <vgoyal@in.ibm.com>
Cc: Andi Kleen <ak@suse.de>
Cc: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>

Signed-off-by: NRussell King <rmk+kernel@arm.linux.org.uk>

utrace removes the ptrace_message field in task_struct.  Move our use
of this field into a new member in thread_info called "syscall"
Signed-off-by: NRussell King <rmk+kernel@arm.linux.org.uk>

  CC      fs/nfs/nfsroot.o
fs/nfs/nfsroot.c:131: error: tokens causes a section type conflict
make[2]: *** [fs/nfs/nfsroot.o] Error 1

This is due to mixing const and non-const content in the same section
which halfway recent gccs absolutely hate.  Fixed by dropping the const.
Signed-off-by: NRalf Baechle <ralf@linux-mips.org>
Signed-off-by: NLinus Torvalds <torvalds@linux-foundation.org>