Message-ID: <20121128212845.GU6434@dastard>
Date:	Thu, 29 Nov 2012 08:28:45 +1100
From:	Dave Chinner <david@...morbit.com>
To:	Glauber Costa <glommer@...allels.com>
Cc:	Chris Wilson <chris@...is-wilson.co.uk>,
	linux-kernel@...r.kernel.org, linux-fsdevel@...r.kernel.org,
	linux-mm@...ck.org, xfs@....sgi.com
Subject: Re: [PATCH 17/19] drivers: convert shrinkers to new count/scan API

On Wed, Nov 28, 2012 at 12:21:54PM +0400, Glauber Costa wrote:
> On 11/28/2012 07:17 AM, Dave Chinner wrote:
> > On Wed, Nov 28, 2012 at 01:13:11AM +0000, Chris Wilson wrote:
> >> On Wed, 28 Nov 2012 10:14:44 +1100, Dave Chinner <david@...morbit.com> wrote:
> >>> +/*
> >>> + * XXX: (dchinner) This is one of the worst cases of shrinker abuse I've seen.
> >>> + *
> >>> + * i915_gem_purge() expects a byte count to be passed, and the minimum object
> >>> + * size is PAGE_SIZE.
> >>
> >> No, purge() expects a count of pages to be freed. Each pass of the
> >> shrinker therefore tries to free a minimum of 128 pages.
> > 
> > Ah, I got the shifts mixed up. I'd been looking at way too much crap
> > already when I saw this. But the fact this can be misunderstood says
> > something about the level of documentation that the code has (i.e.
> > none).
> > 
> >>> The shrinker doesn't work on bytes - it works on
> >>> + * *objects*.
> >>
> >> And I thought you were reviewing the shrinker API to be useful where a
> >> single object may range between 4K and 4G.
> > 
> > Which requires rewriting all the algorithms to not be dependent on
> > the subsystems using a fixed size object. The shrinker control
> > function is called shrink_slab() for a reason - it was expected to
> > be used to shrink caches of fixed sized objects allocated from slab
> > memory.
> > 
> > It has no concept of the amount of memory that each object consumes,
> > just an idea of how much *IO* it takes to replace the object in
> > memory once it's been reclaimed. The DEFAULT_SEEKS is design to
> > encode the fact it generally takes 2 IOs to replace either a LRU
> > page or a filesystem slab object, and so balances the scanning based
> > on that value. i.e. the shrinker algorithms are solidly based around
> > fixed sized objects that have some relationship to the cost of
> > physical IO operations to replace them in the cache.
> 
> One nit: It shouldn't take 2IOs to replace a slab object, right?
> objects.

A random dentry in a small directory will take on IO to read the
inode, then another to read the block the dirent sits in. TO read an
inode froma cached dentry will generally take one IO to read the
inode, and another to read related, out of inode information (e.g.
attributes or extent/block maps). Sometimes it will only take on IO,
sometimes it might take 3 or, in the case of dirents, coult take
hundreds of IOs if the directory structure is large enough.

So a default of 2 seeks to replace any single dentry/inode in the
cache is a pretty good default to use.

> This
> should be the cost of allocating a new page, that can contain, multiple
> Once the page is in, a new object should be quite cheap to come up with.

It's not the cost of allocating the page (a couple of microseconds)
that is being considered - it the 3-4 orders of magnitude worse cost
of reading the object from disk (could be 20ms). The slab/page
allocation is lost in the noise compared to the time it takes to
fill the page cache page with data or a single slab object.
Essentially, slab pages with multiple objects in them are much more
expensive to replace in the cache than a page cache page....

> This is a very wild thought, but now that I am diving deep in the
> shrinker API, and seeing things like this:
> 
> if (reclaim_state) {
>     sc->nr_reclaimed += reclaim_state->reclaimed_slab;
>     reclaim_state->reclaimed_slab = 0;
> }

That's not part of the shrinker - that's part of the vmscan
code, external to the shrinker infrastructure. It's getting
information back from the slab caches behind the shrinkers, and it's
not the full picture because many shrinkers are not backed by slab
caches. It's a work around for not not having accurate feedback from
the shrink_slab() code about how many pages were freed.

Essentially, the problem is an impedance mismatch between the way
the LRUs are scanned/balanced (in pages) and slab caches are managed
(by objects). That's what needs unifying...

> I am becoming more convinced that we should have a page-based mechanism,
> like the rest of vmscan.

Been thought about and consiered before. Would you like to rewrite
the slab code?

> Also, if we are seeing pressure from someone requesting user pages, what
> good does it make to free, say, 35 Mb of memory, if this means we are
> freeing objects across 5k different pages, without actually releasing
> any of them? (still is TBD if this is a theoretical problem or a
> practical one). It would maybe be better to free objects that are
> moderately hot, but are on pages dominated by cold objects...

Yup, that's a problem, but now you're asking shrinker
implementations to know  in great detail the physical locality of
object and not just the temporal locality.  the node-aware LRU list
does this at a coarse level, but to do page based reclaim you need
ot track pages in SL*B that contain unreferenced objects as those
are the only ones that can be reclaimed.

If you have no pages with unreferenced objects, then you don't make
progress and you have to fall back to freeing unreferenced objects
from random pages. ANd under most workloads that aren't benchmarks,
slab object population ends up with little correlation between
physical and temporal locality. Hence this is the norm rather than
the exception..

Also, handling physical locality in this way means we'd need to tie
the shrinker deep into the SLAB/SLUB/SLOB implementation that is
being used to allocate the objects..

There have been various attempts at this sort of thing in the past.
e.g:

http://marc.info/?l=linux-mm&m=112810938004047

or for slab defragmentation:

https://lkml.org/lkml/2010/1/29/332

and more on LRUs in slab caches and general shrinker design in that
thread (definitely worth reading this, at least):

https://lkml.org/lkml/2010/2/2/499

And it's made far more complex by the fact that some shrinkers don't
necessarily free the objects they are working on. e.g. the VFS inode
cache shrinker basically hands objects to XFS, and the XFS inode
cache takes over from there (via the XFS inode cache shrinker) to
free the objects. i.e. two shrinkers act on the same structure...

> > The API change is the first step in the path to removing these built
> > in assumptions. The current API is just insane and any attempt to
> > build on it is going to be futile. 
> 
> Amen, brother!
> 
> > The way I see this developing is
> > this:
> > 
> > 	- make the shrink_slab count -> scan algorithm per node
> > 
> pages are per-node.
> 
> > 	- add information about size of objects in the cache for
> > 	  fixed size object caches.
> > 		- the shrinker now has some idea of how many objects
> > 		  need to be freed to be able to free a page of
> > 		  memory, as well as the relative penalty for
> > 		  replacing them.
> this is still guesswork, telling how many pages it should free, could
> be a better idea.
> 
> > 		- tells the shrinker the size of the cache
> > 		  in bytes so overall memory footprint of the caches
> > 		  can be taken into account
> 
> > 	- add new count and scan operations for caches that are
> > 	  based on memory used, not object counts
> > 		- allows us to use the same count/scan algorithm for
> > 		  calculating how much pressure to put on caches
> > 		  with variable size objects.
> 
> IOW, pages.

Not necessarily - if we are going to deal with multiple objects in a
page as well as multi-page objects, then pages are unable to express
the full range of possibilities. We may as well use byte counts at
this point, expecially when you consider page sizes differ on
different platforms...

> > My care factor mostly ends here, as it will allow XFS to corectly
> > balance the metadata buffer cache (variable size objects) against the
> > inode, dentry and dquot caches which are object based. The next
> > steps that I'm about to give you are based on some discussions with
> > some MM people over bottles of red wine, so take it with a grain of
> > salt...
> > 
> > 	- calculate a "pressure" value for each cache controlled by a
> > 	  shrinker so that the relative memory pressure between
> > 	  caches can be compared. This allows the shrinkers to bias
> > 	  reclaim based on where the memory pressure is being
> > 	  generated
> > 
> 
> Ok, if a cache is using a lot of memory, this would indicate it has the
> dominant workload, right?

Not necessarily. Someone might jus thave run a find across their
filesystem, and that is where all the pressure is coming from. In
this case, you don't want that memory prssure to toss out all the
other caches. I suspect that the "pressure" measure is going to need
to take into account cache hit rates to work properly...

> Should we free from it, or should we free from
> the others, so this ones gets the pages it needs?

That's the million dollar question...

> > 	- start grouping shrinkers into a heirarchy, allowing
> > 	  related shrinkers (e.g. all the caches in a memcg) to be
> > 	  shrunk according resource limits that can be placed on the
> > 	  group. i.e. memory pressure is proportioned across
> > 	  groups rather than many individual shrinkers.
> > 
> pages are already grouped like that!

But shrinkers and slab caches are not.

Besides, once you start grouping shrinkers, why should we treat the
page LRU list scanning any differently from any other cache that has
a shrinker? 

> But one thing we seem to agree is that we have quite a long road ahead

Definitely.

Cheers,

Dave.
-- 
Dave Chinner
david@...morbit.com
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