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<chapter id="mm">
<?dbhtml filename="mm.html"?>
<title>Memory management</title>
<section>
<title>Virtual memory management</title>
<section>
<title>Introduction</title>
<para>Virtual memory is a special memory management technique, used by
kernel to achieve a bunch of mission critical goals. <itemizedlist>
<listitem>
Isolate each task from other tasks that are running on the system at the same time.
</listitem>
<listitem>
Allow to allocate more memory, than is actual physical memory size of the machine.
</listitem>
<listitem>
Allowing, in general, to load and execute two programs that are linked on the same address without complicated relocations.
</listitem>
</itemizedlist></para>
<para><!--
TLB shootdown ASID/ASID:PAGE/ALL.
TLB shootdown requests can come in asynchroniously
so there is a cache of TLB shootdown requests. Upon cache overflow TLB shootdown ALL is executed
<para>
Address spaces. Address space area (B+ tree). Only for uspace. Set of syscalls (shrink/extend etc).
Special address space area type - device - prohibits shrink/extend syscalls to call on it.
Address space has link to mapping tables (hierarchical - per Address space, hash - global tables).
</para>
--></para>
</section>
<section>
<title>Paging</title>
<para>Virtual memory is usually using paged memory model, where virtual
memory address space is divided into the <emphasis>pages</emphasis>
(usually having size 4096 bytes) and physical memory is divided into the
frames (same sized as a page, of course). Each page may be mapped to
some frame and then, upon memory access to the virtual address, CPU
performs <emphasis>address translation</emphasis> during the instruction
execution. Non-existing mapping generates page fault exception, calling
kernel exception handler, thus allowing kernel to manipulate rules of
memory access. Information for pages mapping is stored by kernel in the
<link linkend="page_tables">page tables</link></para>
<para>The majority of the architectures use multi-level page tables,
which means need to access physical memory several times before getting
physical address. This fact would make serios performance overhead in
virtual memory management. To avoid this <link linkend="tlb">Traslation
Lookaside Buffer (TLB)</link> is used.</para>
<para>At the moment HelenOS does not support swapping.</para>
<para>- pouzivame vypadky stranky k alokaci ramcu on-demand v ramci
as_area - na architekturach, ktere to podporuji, podporujeme non-exec
stranky</para>
</section>
<section>
<title>Address spaces</title>
<section>
<title>Address spaces and areas</title>
<para>- adresovy prostor se sklada z tzv. address space areas
usporadanych v B+stromu; tyto areas popisuji vyuzivane casti
adresoveho prostoru patrici do user address space. Kazda cast je dana
svoji bazovou adresou, velikosti a flagy (rwx/dd).</para>
<para>- uzivatelske thready maji moznost manipulovat se svym adresovym
prostorem (vytvaret/resizovat/sdilet) as_areas pomoci syscallu</para>
</section>
<section>
<title>Address Space ID (ASID)</title>
<para>- nektery hardware umoznuje rozlisit ruzne adresove prostory od
sebe (cilem je maximalizovat vyuziti TLB); dela to tak, ze s kazdou
polozkou TLB/strankovacich tabulek sdruzi identifikator adresoveho
prostoru (ASID, RID, ppc32 ???). Tyto id mivaji ruznou sirku: 8-bitu
az 24-bitu (kolik ma ppc32?)</para>
<para>- kernel tomu rozumi a sam pouziva abstrakci ASIDu (na ia64 to
je napr. cislo odvozene od RIDu, na mips32 to je ASID samotny);
existence ASIDu je nutnou podminkou pouziti _global_ page hash table
mechanismu.</para>
<para>- na vsech arch. plati, ze asidu je mnohem mene, nez teoreticky
pocet soucasne bezicich tasku ~ adresovych prostoru, takze je
implementovan mechanismus, ktery umoznuje jednomu adresovemu prostoru
ASID odebrat a pridelit ho jinemu</para>
<para>- vztah task ~ adresovy prostor: teoreticky existuje moznost, ze
je adresovy prostor sdilen vice tasky, avsak tuto moznost nepouzivame
a neni ani nijak osetrena. Tim padem plati, ze kazdy task ma vlastni
adresovy prostor</para>
</section>
</section>
<section>
<title>Virtual address translation</title>
<section id="page_tables">
<title>Page tables</title>
<para>HelenOS kernel has two different approaches to the paging
implementation: <emphasis>4 level page tables</emphasis> and
<emphasis>global hash tables</emphasis>, which are accessible via
generic paging abstraction layer. This division was caused by the
major architectural differences between different platforms.</para>
<formalpara>
<title>4-level page tables</title>
<para>4-level page tables are the generalization of the hardware
capabilities of the certain platforms. <itemizedlist>
<listitem>
ia32 uses 2-level page tables, with full hardware support.
</listitem>
<listitem>
amd64 uses 4-level page tables, also coming with full hardware support.
</listitem>
<listitem>
mips and ppc32 have 2-level tables, software simulated support.
</listitem>
</itemizedlist></para>
</formalpara>
<formalpara>
<title>Global hash tables</title>
<para>- global page hash table: existuje jen jedna v celem systemu
(vyuziva ji ia64), pozn. ia64 ma zatim vypnuty VHPT. Pouziva se
genericke hash table s oddelenymi collision chains</para>
</formalpara>
<para>Thanks to the abstract paging interface, there is possibility
left have more paging implementations, for example B-Tree page
tables.</para>
</section>
<section id="tlb">
<title>Translation Lookaside buffer</title>
<para>- TLB cachuji informace ve strankovacich tabulkach; alternativne
se lze na strankovaci tabulky (ci ruzne hw rozsireni [e.g. VHPT, ppc32
hw hash table]) divat jako na velke TLB</para>
<para>- pri modifikaci mapovani nebo odstraneni mapovani ze
strankovacich tabulek je potreba zajistit konsistenci TLB a techto
tabulek; nutne delat na vsech CPU; na to mame zjednodusenou verzi TLB
shootdown mechanismu; je to variace na algoritmus popsany zde: D.
Black et al., "Translation Lookaside Buffer Consistency: A Software
Approach," Proc. Third Int'l Conf. Architectural Support for
Programming Languages and Operating Systems, 1989, pp. 113-122.</para>
<para>- nutno poznamenat, ze existuji odlehcenejsi verze TLB shootdown
algoritmu</para>
</section>
</section>
</section>
<!-- End of VM -->
<section>
<!-- Phys mem -->
<title>Physical memory management</title>
<section id="zones_and_frames">
<title>Zones and frames</title>
<para><!--graphic fileref="images/mm2.png" /--><!--graphic fileref="images/buddy_alloc.svg" format="SVG" /--></para>
<para>On some architectures not whole physical memory is available for
conventional usage. This limitations require from kernel to maintain a
table of available and unavailable ranges of physical memory addresses.
Main idea of zones is in creating memory zone entity, that is a
continuous chunk of memory available for allocation. If some chunk is
not available, we simply do not put it in any zone.</para>
<para>Zone is also serves for informational purposes, containing
information about number of free and busy frames. Physical memory
allocation is also done inside the certain zone. Allocation of zone
frame must be organized by the <link linkend="frame_allocator">frame
allocator</link> associated with the zone.</para>
<para>Some of the architectures (mips32, ppc32) have only one zone, that
covers whole physical memory, and the others (like ia32) may have
multiple zones. Information about zones on current machine is stored in
BIOS hardware tables or can be hardcoded into kernel during compile
time.</para>
</section>
<section id="frame_allocator">
<title>Frame allocator</title>
<para><mediaobject id="frame_alloc">
<imageobject role="html">
<imagedata fileref="images/frame_alloc.png" format="PNG" />
</imageobject>
<imageobject role="fop">
<imagedata fileref="images.vector/frame_alloc.svg" format="SVG" />
</imageobject>
</mediaobject></para>
<formalpara>
<title>Overview</title>
<para>Frame allocator provides physical memory allocation for the
kernel. Because of zonal organization of physical memory, frame
allocator is always working in context of some zone, thus making
impossible to allocate a piece of memory, which lays in different
zone, which cannot happen, because two adjacent zones can be merged
into one. Frame allocator is also being responsible to update
information on the number of free/busy frames in zone. Physical memory
allocation inside one <link linkend="zones_and_frames">memory
zone</link> is being handled by an instance of <link
linkend="buddy_allocator">buddy allocator</link> tailored to allocate
blocks of physical memory frames.</para>
</formalpara>
<formalpara>
<title>Allocation / deallocation</title>
<para>Upon allocation request, frame allocator tries to find first
zone, that can satisfy the incoming request (has required amount of
free frames to allocate). During deallocation, frame allocator needs
to find zone, that contain deallocated frame. This approach could
bring up two potential problems: <itemizedlist>
<listitem>
Linear search of zones does not any good to performance, but number of zones is not expected to be high. And if yes, list of zones can be replaced with more time-efficient B-tree.
</listitem>
<listitem>
Quickly find out if zone contains required number of frames to allocate and if this chunk of memory is properly aligned. This issue is perfectly solved bu the buddy allocator.
</listitem>
</itemizedlist></para>
</formalpara>
</section>
<section id="buddy_allocator">
<title>Buddy allocator</title>
<section>
<title>Overview</title>
<para><mediaobject id="buddy_alloc">
<imageobject role="html">
<imagedata fileref="images/buddy_alloc.png" format="PNG" />
</imageobject>
<imageobject role="fop">
<imagedata fileref="images.vector/buddy_alloc.svg" format="SVG" />
</imageobject>
</mediaobject></para>
<para>In buddy allocator, memory is broken down into power-of-two
sized naturally aligned blocks. These blocks are organized in an array
of lists in which list with index i contains all unallocated blocks of
the size <mathphrase>2<superscript>i</superscript></mathphrase>. The
index i is called the order of block. Should there be two adjacent
equally sized blocks in list <mathphrase>i</mathphrase> (i.e.
buddies), the buddy allocator would coalesce them and put the
resulting block in list <mathphrase>i + 1</mathphrase>, provided that
the resulting block would be naturally aligned. Similarily, when the
allocator is asked to allocate a block of size
<mathphrase>2<superscript>i</superscript></mathphrase>, it first tries
to satisfy the request from list with index i. If the request cannot
be satisfied (i.e. the list i is empty), the buddy allocator will try
to allocate and split larger block from list with index i + 1. Both of
these algorithms are recursive. The recursion ends either when there
are no blocks to coalesce in the former case or when there are no
blocks that can be split in the latter case.</para>
<!--graphic fileref="images/mm1.png" format="EPS" /-->
<para>This approach greatly reduces external fragmentation of memory
and helps in allocating bigger continuous blocks of memory aligned to
their size. On the other hand, the buddy allocator suffers increased
internal fragmentation of memory and is not suitable for general
kernel allocations. This purpose is better addressed by the <link
linkend="slab">slab allocator</link>.</para>
</section>
<section>
<title>Implementation</title>
<para>The buddy allocator is, in fact, an abstract framework wich can
be easily specialized to serve one particular task. It knows nothing
about the nature of memory it helps to allocate. In order to beat the
lack of this knowledge, the buddy allocator exports an interface that
each of its clients is required to implement. When supplied an
implementation of this interface, the buddy allocator can use
specialized external functions to find buddy for a block, split and
coalesce blocks, manipulate block order and mark blocks busy or
available. For precize documentation of this interface, refer to
<emphasis>"HelenOS Generic Kernel Reference Manual"</emphasis>.</para>
<formalpara>
<title>Data organization</title>
<para>Each entity allocable by the buddy allocator is required to
contain space for storing block order number and a link variable
used to interconnect blocks within the same order.</para>
<para>Whatever entities are allocated by the buddy allocator, the
first entity within a block is used to represent the entire block.
The first entity keeps the order of the whole block. Other entities
within the block are assigned the magic value
<constant>BUDDY_INNER_BLOCK</constant>. This is especially important
for effective identification of buddies in one-dimensional array
because the entity that represents a potential buddy cannot be
associated with <constant>BUDDY_INNER_BLOCK</constant> (i.e. if it
is associated with <constant>BUDDY_INNER_BLOCK</constant> then it is
not a buddy).</para>
<para>Buddy allocator always uses first frame to represent frame
block. This frame contains <varname>buddy_order</varname> variable
to provide information about the block size it actually represents (
<mathphrase>2<superscript>buddy_order</superscript></mathphrase>
frames block). Other frames in block have this value set to magic
<constant>BUDDY_INNER_BLOCK</constant> that is much greater than
buddy <varname>max_order</varname> value.</para>
<para>Each <varname>frame_t</varname> also contains pointer member
to hold frame structure in the linked list inside one order.</para>
</formalpara>
<formalpara>
<title>Allocation algorithm</title>
<para>Upon <mathphrase>2<superscript>i</superscript></mathphrase>
frames block allocation request, allocator checks if there are any
blocks available at the order list <varname>i</varname>. If yes,
removes block from order list and returns its address. If no,
recursively allocates
<mathphrase>2<superscript>i+1</superscript></mathphrase> frame
block, splits it into two
<mathphrase>2<superscript>i</superscript></mathphrase> frame blocks.
Then adds one of the blocks to the <varname>i</varname> order list
and returns address of another.</para>
</formalpara>
<formalpara>
<title>Deallocation algorithm</title>
<para>Check if block has so called buddy (another free
<mathphrase>2<superscript>i</superscript></mathphrase> frame block
that can be linked with freed block into the
<mathphrase>2<superscript>i+1</superscript></mathphrase> block).
Technically, buddy is a odd/even block for even/odd block
respectively. Plus we can put an extra requirement, that resulting
block must be aligned to its size. This requirement guarantees
natural block alignment for the blocks coming out the allocation
system.</para>
<para>Using direct pointer arithmetics,
<varname>frame_t::ref_count</varname> and
<varname>frame_t::buddy_order</varname> variables, finding buddy is
done at constant time.</para>
</formalpara>
</section>
</section>
<section id="slab">
<title>Slab allocator</title>
<section>
<title>Overview</title>
<para><termdef><glossterm>Slab</glossterm> represents a contiguous
piece of memory, usually made of several physically contiguous
pages.</termdef> <termdef><glossterm>Slab cache</glossterm> consists
of one or more slabs.</termdef></para>
<para>The majority of memory allocation requests in the kernel are for
small, frequently used data structures. For this purpose the slab
allocator is a perfect solution. The basic idea behind the slab
allocator is to have lists of commonly used objects available packed
into pages. This avoids the overhead of allocating and destroying
commonly used types of objects such threads, virtual memory structures
etc. Also due to the exact allocated size matching, slab allocation
completely eliminates internal fragmentation issue.</para>
</section>
<section>
<title>Implementation</title>
<para><mediaobject id="slab_alloc">
<imageobject role="html">
<imagedata fileref="images/slab_alloc.png" format="PNG" />
</imageobject>
<imageobject role="fop">
<imagedata fileref="images.vector/slab_alloc.svg" format="SVG" />
</imageobject>
</mediaobject></para>
<para>The SLAB allocator is closely modelled after <ulink
url="http://www.usenix.org/events/usenix01/full_papers/bonwick/bonwick_html/">
OpenSolaris SLAB allocator by Jeff Bonwick and Jonathan Adams </ulink>
with the following exceptions: <itemizedlist>
<listitem>
empty SLABS are deallocated immediately (in Linux they are kept in linked list, in Solaris ???)
</listitem>
<listitem>
empty magazines are deallocated when not needed (in Solaris they are held in linked list in slab cache)
</listitem>
</itemizedlist> Following features are not currently supported but
would be easy to do: <itemizedlist>
<listitem>
- cache coloring
</listitem>
<listitem>
- dynamic magazine grow (different magazine sizes are already supported, but we would need to adjust allocation strategy)
</listitem>
</itemizedlist></para>
<section>
<title>Magazine layer</title>
<para>Due to the extensive bottleneck on SMP architures, caused by
global SLAB locking mechanism, making processing of all slab
allocation requests serialized, a new layer was introduced to the
classic slab allocator design. Slab allocator was extended to
support per-CPU caches 'magazines' to achieve good SMP scaling.
<termdef>Slab SMP perfromance bottleneck was resolved by introducing
a per-CPU caching scheme called as <glossterm>magazine
layer</glossterm></termdef>.</para>
<para>Magazine is a N-element cache of objects, so each magazine can
satisfy N allocations. Magazine behaves like a automatic weapon
magazine (LIFO, stack), so the allocation/deallocation become simple
push/pop pointer operation. Trick is that CPU does not access global
slab allocator data during the allocation from its magazine, thus
making possible parallel allocations between CPUs.</para>
<para>Implementation also requires adding another feature as the
CPU-bound magazine is actually a pair of magazines to avoid
thrashing when during allocation/deallocatiion of 1 item at the
magazine size boundary. LIFO order is enforced, which should avoid
fragmentation as much as possible.</para>
<para>Another important entity of magazine layer is a full magazine
depot, that stores full magazines which are used by any of the CPU
magazine caches to reload active CPU magazine. Magazine depot can be
pre-filled with full magazines during initialization, but in current
implementation it is filled during object deallocation, when CPU
magazine becomes full.</para>
<para>Slab allocator control structures are allocated from special
slabs, that are marked by special flag, indicating that it should
not be used for slab magazine layer. This is done to avoid possible
infinite recursions and deadlock during conventional slab allocaiton
requests.</para>
</section>
<section>
<title>Allocation/deallocation</title>
<para>Every cache contains list of full slabs and list of partialy
full slabs. Empty slabs are immediately freed (thrashing will be
avoided because of magazines).</para>
<para>The SLAB allocator allocates lots of space and does not free
it. When frame allocator fails to allocate the frame, it calls
slab_reclaim(). It tries 'light reclaim' first, then brutal reclaim.
The light reclaim releases slabs from cpu-shared magazine-list,
until at least 1 slab is deallocated in each cache (this algorithm
should probably change). The brutal reclaim removes all cached
objects, even from CPU-bound magazines.</para>
<formalpara>
<title>Allocation</title>
<para><emphasis>Step 1.</emphasis> When it comes to the allocation
request, slab allocator first of all checks availability of memory
in local CPU-bound magazine. If it is there, we would just "pop"
the CPU magazine and return the pointer to object.</para>
<para><emphasis>Step 2.</emphasis> If the CPU-bound magazine is
empty, allocator will attempt to reload magazin, swapping it with
second CPU magazine and returns to the first step.</para>
<para><emphasis>Step 3.</emphasis> Now we are in the situation
when both CPU-bound magazines are empty, which makes allocator to
access shared full-magazines depot to reload CPU-bound magazines.
If reload is succesful (meaning there are full magazines in depot)
algoritm continues at Step 1.</para>
<para><emphasis>Step 4.</emphasis> Final step of the allocation.
In this step object is allocated from the conventional slab layer
and pointer is returned.</para>
</formalpara>
<formalpara>
<title>Deallocation</title>
<para><emphasis>Step 1.</emphasis> During deallocation request,
slab allocator will check if the local CPU-bound magazine is not
full. In this case we will just push the pointer to this
magazine.</para>
<para><emphasis>Step 2.</emphasis> If the CPU-bound magazine is
full, allocator will attempt to reload magazin, swapping it with
second CPU magazine and returns to the first step.</para>
<para><emphasis>Step 3.</emphasis> Now we are in the situation
when both CPU-bound magazines are full, which makes allocator to
access shared full-magazines depot to put one of the magazines to
the depot and creating new empty magazine. Algoritm continues at
Step 1.</para>
</formalpara>
</section>
</section>
</section>
<!-- End of Physmem -->
</section>
<section>
<title>Memory sharing</title>
<para>Not implemented yet(?)</para>
</section>
</chapter>