On Tue, Aug 15, 2006 at 02:03:25PM +0200, Peter Zijlstra (a.p.zijlstra@chello.nl) wrote:

On Tue, 2006-08-15 at 15:26 +0400, Evgeniy Polyakov wrote:

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On Tue, Aug 15, 2006 at 12:55:02PM +0200, Peter Zijlstra (a.p.zijlstra@chello.nl) wrote:

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Userspace can sak for next packet and pointer to the new location will
be removed.

/sak/ask/?

I'm not understanding, if you have a page A with two packets, a and b;
once you map that page into user-space that process has access to both
packets, which is a security problem. How are you going to solve this?

Yep, there is such issue.
But no one is ever going to replace socket code with zero-copy
interfaces - Linux has backward compatibility noone ever had, so
send()/recv() will be there.

The new AIO network API should be able to provide the needed userspace
changes.

Kevent based network AIO already handle it.
This changes are different.

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It is new interface which can be changed as described in previous
e-mails - copy if next chunk belongs to different socket and so on.

But if you copy you're not zero-copy anymore. If you copy every second
packet you have a 1/2 copy receive, but not zero.

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Initial user will be sniffer, which should get all packets.

Only the root user may get all packets. And most enterprise systems I've
seen don't generally run a sniffer. That is usually done by
redirecting/copying the data stream in a router and attach a second host
to analyse the data.

Nice words. And what "second host" is supposed to do with that traffic?
I expect it should not be Linux (not an "enterprise system"?). WIll it
copy to userspace or have some kind of a zero-copy?

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As described in recent threads [3] it is also possible to eliminate any 
kind of main system OOM influence on network dataflow processing, thus 
it is possible to prevent deadlock for systems, which use network as 
memory storage (swap over network, iSCSI, NBD and so on).

How? You have never stated how you will avoid getting all packets stuck
in blocked sockets.

Each socket has it's limit, so if allocator got enough memory, blocked
sockets will not affect it's behaviour.

But isn't the total capacity of the network stack much larger than any
allocator can provide?

TCP has 768kb limit on my amd64 with 1gb of ram, so I expect allocator
can handle all requests.

But the capacity of the network stack is larger than this (arbitrary)
limit. It is possible to have all 768kb worth of packets stuck on
blocked sockets.

And so what? You said, that separated from main allocator can not handle
all memory requests being done by the stack, as you can see it easily
can.

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And there is a simple task in TODO list to dynamically grow cache when
threshold of memory is in use. It is really simple task and will be
implemented as soon as I complete suggestions mentioned by Andrew Morton.

Growing will not help, the problem is you are out of memory, you cannot
grow at that point.

You do not see the point of network tree allocator.

It can live with main system OOM since it has preallocated separate
pool, which can be increased when there is a requirement for that, for
example when system is not in OOM.

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On another note, I think you misunderstand our SLAB allocator; we do not
round up to nearest order page alloc per object; SLAB is build to avoid
that and is designed to pack equal size objects into pages. The kmalloc
allocator is build on top of several SLAB allocators; each with its
specific size objects to serve.

For example, the 64 byte SLAB will serve 64 byte objects, and packs
about PAGE_SIZE/64 per page (about since there is some overhead).

So the actual internal fragmentation of the current kmalloc/SLAB
allocator is not as bad as you paint it. The biggest problem we have
with the SLAB thing is getting pages back from it. (And the horrific
complexity of the current implementation)

Ok, not SLAB, but kmaloc/SLAB.

The page-allocator does what you describe, but hardly anybody uses that
to store small objects.

Network stack uses kmalloc.

skbuff_head_cache and skbuff_fclone_cache are SLABs.

It is quite small part of the stack, isn't it?
And btw, they still suffer from SLAB design, since it is possibly to get
another smaller object right after all skbs are allocated from given page.
It is a minor thing of course, but nevertheless worh mentioning.

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Page allocator - buddy allocator, gives out memory in 1<<n pages.

SLAB allocator - uses the page allocator for backing, each SLAB issues
objects of a fixed, predetermined size, packed in pages.

kmalloc - uses a collection of SLAB allocators to issue 'variable' size
objects (see kmalloc_sizes.h - as you will see internal fragmentation
can become quite large for larger objects, but small objects do rather
well - and one could always add a frequently used size if it shows to be
beneficial).

There is no "frequently used size", kmalloc() does not know what size is
frequent and what is not. And there are other mentioned problems with
kmalloc/SLAB besides power-of-two, which prevent fragmentation problem
resolution.

Yes SLAB is a horrid thing on some points but very good at a lot of
other things. But surely there are frequently used sizes, kmalloc will
not know, but a developer with profiling tools might.

Does not scale - admin must run system under profiling, add new
entries into kmalloc_sizes.h recompile the kernel... No way.

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That allocator uses power-of-two allocation, so there is extremely
large overhead for several (and in some cases for all) usage cases
(e1000 with jumbo frames and unix sockets).

Wrong example :-), e1000 is the only driver that doesn't do high order
allocs for jumbo frames. But yes, the other drivers should be fixed,
relying on higher order allocations is unsound.

:) do you read netdev@? There are several quite long recent discussions 
where network hackers blame exactly e1000 for it's hardware problems and
ugly memory usage model.
We even think how to change struct sk_buff - Holy Grail of network code
- just to help e1000 driver (well, not exactly for e1000, but that
driver was a cause).

Have you seen the latest code? It allocates single pages and puts them
in the skb_shared_info fragments. Surely it might have been the one
pushing for these changes, but they are done. Current status.

Plese check e1000_alloc_rx_buffers() function and rx_buffer_len value.
You are wrong.

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SLAB allows to have chunks of memory from differenct CPU, so it is
impossible to create defragmentation, thus kmalloc/SLAB by design will
suffer from fragmentation.

*confused* memory is not bound to CPUs other than by NUMA, but even
there there is only a single address space. 

Each slab can have objects allocated in different CPUs, it was done to
reduce freeing algorithm. If system wants to defragment several objects
into bigger one, it must check all CPUs and find in which cache those
objects are placed, which is extremely expensive, so SLAB can not
perform defragmentation.

What you are referring to is coalescence, and yes coalescing over page
boundaries is hard in the SLAB layer, the page allocator does that.

Page boundaries is only minor part of the problem.
Objects from the same page can live in different (per-cpu) SLAB caches.

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Graphs of power-of-two vs. NTA overhead is shown on projects' homepage 
- overhead is extremely large.

Yes seen that, but as stated, hardly anybody uses the page allocator to
store small objects. However if you do, you get large internal
fragmentation but zero external fragmentation (on that allocation
level).

Truncated cat /proc/slabinfo on my machine (usual desktop):
size-32             1170   1232     32
size-128             663    780    128
size-64             4239   9558     64

Sure, point being?

size-64             4497   4602     64   59    1 : tunables  120   60
8 : slabdata     78     78      0

4497 objects used out of 4602 available, object size 64 bytes, 59
objects per slab of 1 page. .... 78 pages with active objects out of 78
pages allocated. 

It was your words quoted above that "hardly anybody uses the page
allocator to store small objects", as you can see, there are quite a few
of them allocated through kmalloc but not through kmem_cache.
 

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This is where my SROG allocator comes in, it is used to group objects by
lifetime and returns the pages to the page allocator. This makes the
whole allocator short-lived and hence cannot add (external)
fragmentation on this level. The use I have for that is that I can then
properly gauge how much memory there is available. External
fragmentation and guarantees can be difficult to reconcile.

I have no idea how fast/slow the SROG allocator is, and don't really
care since its only used as a fallback allocator; what I do care about
is determinism (in space).

However, I do have a patch that converts the whole skb layer to use the
SROG allocator, not only the payload, so I could do some test. But this
is not a serious candidate until all jumbo frame capable drivers have
been converted to skb fragments instead of high order allocations - a
Good Thing [tm].

You created SROG after my suggestion and discussion about NTA and it works 
well for it's purpose (doesn't it?), further extension could lead to creation 
of NTA (or could not).

I started with a very broken in-situ allocator (that tried to do the
same thing) in the very first patch. It was only later that I realised
the full extend of the skbuff requirements.

And no, NTA is too complex an allocator to do what I need. And more
specifically its design is quite contrary to what I have done. I
create/destroy an allocator instance per packet, you have one allocator
instance and serve multiple packets.

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SROG is a wrapper on top of alloc_pages and list of free objects,
there are "several" differencies between allocators and I do not see how
they can compete right now.

Yes allocators are build in layers, the page allocator the the basic
building block in Linux.

Ok, it's your point.
Actum est, ilicet.

-- 
	Evgeniy Polyakov