On Tue, Mar 30, 2021 at 03:30:54PM -0400, Johannes Weiner wrote: > Hi Willy, > > On Mon, Mar 29, 2021 at 05:58:32PM +0100, Matthew Wilcox wrote: > > I'm going to respond to some points in detail below, but there are a > > couple of overarching themes that I want to bring out up here. > > > > Grand Vision > > ~~~~~~~~~~~~ > > > > I haven't outlined my long-term plan. Partly because it is a _very_ > > long way off, and partly because I think what I'm doing stands on its > > own. But some of the points below bear on this, so I'll do it now. > > > > Eventually, I want to make struct page optional for allocations. It's too > > small for some things (allocating page tables, for example), and overly > > large for others (allocating a 2MB page, networking page_pool). I don't > > want to change its size in the meantime; having a struct page refer to > > PAGE_SIZE bytes is something that's quite deeply baked in. > > Right, I think it's overloaded and it needs to go away from many > contexts it's used in today. > > I think it describes a real physical thing, though, and won't go away > as a concept. More on that below. I'm at least 90% with you on this, and we're just quibbling over details at this point, I think. > > In broad strokes, I think that having a Power Of Two Allocator > > with Descriptor (POTAD) is a useful foundational allocator to have. > > The specific allocator that we call the buddy allocator is very clever for > > the 1990s, but touches too many cachelines to be good with today's CPUs. > > The generalisation of the buddy allocator to the POTAD lets us allocate > > smaller quantities (eg a 512 byte block) and allocate descriptors which > > differ in size from a struct page. For an extreme example, see xfs_buf > > which is 360 bytes and is the descriptor for an allocation between 512 > > and 65536 bytes. > > I actually disagree with this rather strongly. If anything, the buddy > allocator has turned out to be a pretty poor fit for the foundational > allocator. > > On paper, it is elegant and versatile in serving essentially arbitrary > memory blocks. In practice, we mostly just need 4k and 2M chunks from > it. And it sucks at the 2M ones because of the fragmentation caused by > the ungrouped 4k blocks. That's a very Intel-centric way of looking at it. Other architectures support a multitude of page sizes, from the insane ia64 (4k, 8k, 16k, then every power of four up to 4GB) to more reasonable options like (4k, 32k, 256k, 2M, 16M, 128M). But we (in software) shouldn't constrain ourselves to thinking in terms of what the hardware currently supports. Google have data showing that for their workloads, 32kB is the goldilocks size. I'm sure for some workloads, it's much higher and for others it's lower. But for almost no workload is 4kB the right choice any more, and probably hasn't been since the late 90s. > The great thing about the slab allocator isn't just that it manages > internal fragmentation of the larger underlying blocks. It also groups > related objects by lifetime/age and reclaimability, which dramatically > mitigates the external fragmentation of the memory space. > > The buddy allocator on the other hand has no idea what you want that > 4k block for, and whether it pairs up well with the 4k block it just > handed to somebody else. But the decision it makes in that moment is > crucial for its ability to serve larger blocks later on. > > We do some mobility grouping based on how reclaimable or migratable > the memory is, but it's not the full answer. I don't think that's entirely true. The vast majority of memory in any machine is either anonymous or page cache. The problem is that right now, all anonymous and page cache allocations are order-0 (... or order-9). So the buddy allocator can't know anything useful about the pages and will often allocate one order-0 page to the page cache, then allocate its buddy to the slab cache in order to allocate the radix_tree_node to store the pointer to the page in (ok, radix tree nodes come from an order-2 cache, but it still prevents this order-9 page from being assembled). If the movable allocations suddenly start being order-3 and order-4, the unmovable, unreclaimable allocations are naturally going to group down in the lower orders, and we won't have the problem that a single dentry blocks the allocation of an entire 2MB page. The problem, for me, with the ZONE_MOVABLE stuff is that it requires sysadmin intervention to set up. I don't have a ZONE_MOVABLE on my laptop. The allocator should be automatically handling movability hints without my intervention. > A variable size allocator without object type grouping will always > have difficulties producing anything but the smallest block size after > some uptime. It's inherently flawed that way. I think our buddy allocator is flawed, to be sure, but only because it doesn't handle movable hints more aggressively. For example, at the point that a largeish block gets a single non-movable allocation, all the movable allocations within that block should be migrated out. If the offending allocation is freed quickly, it all collapses into a large, useful chunk, or if not, then it provides a sponge to soak up other non-movable allocations. > > What I haven't touched on anywhere in this, is whether a folio is the > > descriptor for all POTA or whether it's specifically the page cache > > descriptor. I like the idea of having separate descriptors for objects > > in the page cache from anonymous or other allocations. But I'm not very > > familiar with the rmap code, and that wants to do things like manipulate > > the refcount on a descriptor without knowing whether it's a file or > > anon page. Or neither (eg device driver memory mapped to userspace. > > Or vmalloc memory mapped to userspace. Or ...) > > The rmap code is all about the page type specifics, but once you get > into mmap, page reclaim, page migration, we're dealing with fully > fungible blocks of memory. > > I do like the idea of using actual language typing for the different > things struct page can be today (fs page), but with a common type to > manage the fungible block of memory backing it (allocation state, LRU > & aging state, mmap state etc.) > > New types for the former are an easier sell. We all agree that there > are too many details of the page - including the compound page > implementation detail - inside the cache library, fs code and drivers. > > It's a slightly tougher sell to say that the core VM code itself > (outside the cache library) needs a tighter abstraction for the struct > page building block and the compound page structure. At least at this > time while we're still sorting out how it all may work down the line. > Certainly, we need something to describe fungible memory blocks: > either a struct page that can be 4k and 2M compound, or a new thing > that can be backed by a 2M struct page or a 4k struct smallpage. We > don't know yet, so I would table the new abstraction type for this. > > I generally don't think we want a new type that does everything that > the overloaded struct page already does PLUS the compound > abstraction. Whatever name we pick for it, it'll always be difficult > to wrap your head around such a beast. > > IMO starting with an explicit page cache descriptor that resolves to > struct page inside core VM code (and maybe ->fault) for now makes the > most sense: it greatly mitigates the PAGE_SIZE and tail page issue > right away, and it's not in conflict with, but rather helps work > toward, replacing the fungible memory unit behind it. Right, and that's what struct folio is today. It eliminates tail pages from consideration in a lot of paths. I think it also makes sense for struct folio to be used for anonymous memory. But I think that's where it stops; it isn't for Slab, it isn't for page table pages, and it's not for ZONE_DEVICE pages. > There isn't too much overlap or generic code between cache and anon > pages such that sharing a common descriptor would be a huge win (most > overlap is at the fungible memory block level, and the physical struct > page layout of course), so I don't think we should aim for a generic > abstraction for both. They're both on the LRU list, they use a lot of the same PageFlags, they both have a mapcount and refcount, and they both have memcg_data. The only things they really use differently are mapping, index and private. And then we have to consider shmem which uses both in a pretty eldritch way. > As drivers go, I think there are slightly different requirements to > filesystems, too. For filesystems, when the VM can finally do it (and > the file range permits it), I assume we want to rather transparently > increase the unit of data transfer from 4k to 2M. Most drivers that > currently hardcode alloc_page() or PAGE_SIZE OTOH probably don't want > us to bump their allocation sizes. If you take a look at my earlier work, you'll see me using a range of sizes in the page cache, starting at 16kB and gradually increasing to (theoretically) 2MB, although the algorithm tended to top out around 256kB. Doing particularly large reads could see 512kB/1MB reads, but it was very hard to hit 2MB in practice. I wasn't too concerned at the time, but my point is that we do want to automatically tune the size of the allocation unit to the workload. An application which reads in 64kB chunks is giving us a pretty clear signal that they want to manage memory in 64kB chunks. > > It'd probably be better to have the dcache realise that its old entries > > aren't useful any more and age them out instead of relying on memory > > pressure to remove old entries, so this is probably an unnecessary > > digression. > > It's difficult to identify a universally acceptable line for > usefulness of caches other than physical memory pressure. > > The good thing about the memory pressure threshold is that you KNOW > somebody else has immediate use for the memory, and you're justified > in recycling and reallocating caches from the cold end. > > Without that, you'd either have to set an arbitrary size cutoff or an > arbitrary aging cutoff (not used in the last minute e.g.). But optimal > settings for either of those depend on the workload, and aren't very > intuitive to configure. For the dentry cache, I think there is a more useful metric, and that's length of the hash chain. If it gets too long, we're spending more time walking it than we're saving by having entries cached. Starting reclaim based on "this bucket of the dcache has twenty entries in it" would probably work quite well. > Our levels of internal fragmentation are historically low, which of > course is nice by itself. But that's also what's causing problems in > the form of external fragmentation, and why we struggle to produce 2M > blocks. It's multitudes easier to free one 2M slab page of > consecutively allocated inodes than it is to free 512 batches of > different objects with conflicting lifetimes, ages, or potentially > even reclaimability. Unf. I don't think freeing 2MB worth of _anything_ is ever going to be easy enough to rely on. My actual root filesystem: xfs_inode 143134 144460 1024 32 8 : tunables 0 0 0 : slabdata 4517 4517 0 So we'd have to be able to free 2048 of those 143k inodes, and they all have to be consecutive (and aligned). I suppose we could model that and try to work out how many we'd have to be able to free in order to get all 2048 in any page free, but I bet it's a variant of the Birthday Paradox, and we'd find it's something crazy like half of them. Without slab gaining the ability to ask users to relocate allocations, I think any memory sent to slab is never coming back. So ... even if I accept every part of your vision as the way things are going to be, I think the folio patchset I have now is a step in the right direction. I'm going to send a v6 now and hope it's not too late for this merge window. -- Linux-cachefs mailing list Linux-cachefs@xxxxxxxxxx https://listman.redhat.com/mailman/listinfo/linux-cachefs
