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<chapter id="mm">
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<title>Memory management</title>
<para>In previous chapters, this book described the scheduling subsystem as
the creator of the impression that threads execute in parallel. The memory
management subsystem, on the other hand, creates the impression that there
is enough physical memory for the kernel and that userspace tasks have the
entire address space only for themselves.</para>
<section>
<title>Physical memory management</title>
<section id="zones_and_frames">
<title>Zones and frames</title>
<para>HelenOS represents continuous areas of physical memory in
structures called frame zones (abbreviated as zones). Each zone contains
information about the number of allocated and unallocated physical
memory frames as well as the physical base address of the zone and
number of frames contained in it. A zone also contains an array of frame
structures describing each frame of the zone and, in the last, but not
the least important, front, each zone is equipped with a buddy system
that faciliates effective allocation of power-of-two sized block of
frames.</para>
<para>This organization of physical memory provides good preconditions
for hot-plugging of more zones. There is also one currently unused zone
attribute: <code>flags</code>. The attribute could be used to give a
special meaning to some zones in the future.</para>
<para>The zones are linked in a doubly-linked list. This might seem a
bit ineffective because the zone list is walked everytime a frame is
allocated or deallocated. However, this does not represent a significant
performance problem as it is expected that the number of zones will be
rather low. Moreover, most architectures merge all zones into
one.</para>
<para>For each physical memory frame found in a zone, there is a frame
structure that contains number of references and data used by buddy
system.</para>
</section>
<section id="frame_allocator">
<title>Frame allocator</title>
<para>The frame allocator satisfies kernel requests to allocate
power-of-two sized blocks of physical memory. Because of zonal
organization of physical memory, the frame allocator is always working
within a context of some frame zone. In order to carry out the
allocation requests, the frame allocator is tightly integrated with the
buddy system belonging to the zone. The frame allocator is also
responsible for updating information about the number of free and busy
frames in the zone. <figure>
<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>
<title>Frame allocator scheme.</title>
</figure></para>
<formalpara>
<title>Allocation / deallocation</title>
<para>Upon allocation request via function <code>frame_alloc</code>,
the frame allocator first tries to find a zone that can satisfy the
request (i.e. has the required amount of free frames). Once a suitable
zone is found, the frame allocator uses the buddy allocator on the
zone's buddy system to perform the allocation. During deallocation,
which is triggered by a call to <code>frame_free</code>, the frame
allocator looks up the respective zone that contains the frame being
deallocated. Afterwards, it calls the buddy allocator again, this time
to take care of deallocation within the zone's buddy system.</para>
</formalpara>
</section>
<section id="buddy_allocator">
<title>Buddy allocator</title>
<para>In the buddy system, the memory is broken down into power-of-two
sized naturally aligned blocks. These blocks are organized in an array
of lists, in which the list with index i contains all unallocated blocks
of 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 the list 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 the 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 a larger block from the 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>
<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>.<figure>
<mediaobject id="buddy_alloc">
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<imagedata fileref="images/buddy_alloc.png" format="PNG" />
</imageobject>
<imageobject role="fop">
<imagedata fileref="images.vector/buddy_alloc.svg" format="SVG" />
</imageobject>
</mediaobject>
<title>Buddy system scheme.</title>
</figure></para>
<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 with an
implementation of this interface, the buddy allocator can use
specialized external functions to find a buddy for a block, split and
coalesce blocks, manipulate block order and mark blocks busy or
available.</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 a 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>
</formalpara>
</section>
</section>
<section id="slab">
<title>Slab allocator</title>
<para>The majority of memory allocation requests in the kernel is for
small, frequently used data structures. The basic idea behind the slab
allocator is that deployment of lists of preallocated, commonly used
objects. Whenever an object is to be allocated, the slab allocator takes
the first item from the list corresponding to the object type. This
avoids the overhead of allocating and dellocating commonly used types of
objects such as threads, B+tree nodes etc. Due to the fact that the
sizes of the requested and allocated object match, the slab allocator
significantly eliminates internal fragmentation. Moreover, each list can
have a constructor and a destructor, which leads to performance gains
because constructed and then destroyed objects don't need to be
reinitialized<footnote>
<para>Provided that the assumption that the destructor leaves the
object in a consistent state holds.</para>
</footnote>.</para>
<para>In the slab allocator, objects of one type are kept in continuous
areas of physical memory called slabs. Slabs can span from one to
several physical memory frames. Slabs of objects of one type are stored
in slab caches. When the allocator needs to allocate an object, it
searches available slabs. When the slab does not contain any free
obejct, a new slab is allocated and added to the cache. Contrary to
allocation, deallocated objects are returned to their respective slabs.
Empty slabs are deallocated immediately while empty slab caches are not
freed until HelenOS runs short of memory.</para>
<para><figure>
<mediaobject id="slab_alloc">
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<imagedata fileref="images/slab_alloc.png" format="PNG" />
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<imageobject role="fop">
<imagedata fileref="images.vector/slab_alloc.svg" format="SVG" />
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<title>Slab allocator scheme.</title>
</figure></para>
<section>
<para>
<termdef />
<termdef />
</para>
</section>
<section>
<title>Implementation</title>
<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></listitem>
<listitem>
empty magazines are deallocated when not needed
</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 the common full
magazine list (also called a depot), that stores full magazines that
may be used by any of the CPU magazine caches to reload active CPU
magazine. This list of magazines 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>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>
</section>
<section>
<title>Address spaces</title>
<section>
<title>Address space areas</title>
<para>Each address space consists of mutually disjunctive continuous
address space areas. Address space area is precisely defined by its
base address and the number of frames/pages is contains.</para>
<para>Address space area , that define behaviour and permissions on
the particular area. <itemizedlist>
<listitem>
<emphasis>AS_AREA_READ</emphasis>
flag indicates reading permission.
</listitem>
<listitem>
<emphasis>AS_AREA_WRITE</emphasis>
flag indicates writing permission.
</listitem>
<listitem>
<emphasis>AS_AREA_EXEC</emphasis>
flag indicates code execution permission. Some architectures do not support execution persmission restriction. In this case this flag has no effect.
</listitem>
<listitem>
<emphasis>AS_AREA_DEVICE</emphasis>
marks area as mapped to the device memory.
</listitem>
</itemizedlist></para>
<para>Kernel provides possibility tasks create/expand/shrink/share its
address space via the set of syscalls.</para>
</section>
<section>
<title>Address Space ID (ASID)</title>
<para>When switching to the different task, kernel also require to
switch mappings to the different address space. In case TLB cannot
distinguish address space mappings, all mapping information in TLB
from the old address space must be flushed, which can create certain
uncessary overhead during the task switching. To avoid this, some
architectures have capability to segregate different address spaces on
hardware level introducing the address space identifier as a part of
TLB record, telling the virtual address space translation unit to
which address space this record is applicable.</para>
<para>HelenOS kernel can take advantage of this hardware supported
identifier by having an ASID abstraction which is somehow related to
the corresponding architecture identifier. I.e. on ia64 kernel ASID is
derived from RID (region identifier) and on the mips32 kernel ASID is
actually the hardware identifier. As expected, this ASID information
record is the part of <emphasis>as_t</emphasis> structure.</para>
<para>Due to the hardware limitations, hardware ASID has limited
length from 8 bits on ia64 to 24 bits on mips32, which makes it
impossible to use it as unique address space identifier for all tasks
running in the system. In such situations special ASID stealing
algoritm is used, which takes ASID from inactive task and assigns it
to the active task.</para>
<para><classname>ASID stealing algoritm here.</classname></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. Such different functionality was
caused by the major architectural differences between supported
platforms. This abstraction is implemented with help of the global
structure of pointers to basic mapping functions
<emphasis>page_mapping_operations</emphasis>. To achieve different
functionality of page tables, corresponding layer must implement
functions, declared in
<emphasis>page_mapping_operations</emphasis></para>
<formalpara>
<title>4-level page tables</title>
<para>4-level page tables are the generalization of the hardware
capabilities of several architectures.<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. ASID support is
required to use global hash tables.</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>Due to the extensive overhead during the page mapping lookup in
the page tables, all architectures has fast assotiative cache memory
built-in CPU. This memory called TLB stores recently used page table
entries.</para>
<section id="tlb_shootdown">
<title>TLB consistency. TLB shootdown algorithm.</title>
<para>Operating system is responsible for keeping TLB consistent by
invalidating the contents of TLB, whenever there is some change in
page tables. Those changes may occur when page or group of pages
were unmapped, mapping is changed or system switching active address
space to schedule a new system task (which is a batch unmap of all
address space mappings). Moreover, this invalidation operation must
be done an all system CPUs because each CPU has its own independent
TLB cache. Thus maintaining TLB consistency on SMP configuration as
not as trivial task as it looks at the first glance. Naive solution
would assume remote TLB invalidatation, which is not possible on the
most of the architectures, because of the simple fact - flushing TLB
is allowed only on the local CPU and there is no possibility to
access other CPUs' TLB caches.</para>
<para>Technique of remote invalidation of TLB entries is called "TLB
shootdown". HelenOS uses a variation of the algorithm described by
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>As the situation demands, you will want partitial invalidation
of TLB caches. In case of simple memory mapping change it is
necessary to invalidate only one or more adjacent pages. In case if
the architecture is aware of ASIDs, during the address space
switching, kernel invalidates only entries from this particular
address space. Final option of the TLB invalidation is the complete
TLB cache invalidation, which is the operation that flushes all
entries in TLB.</para>
<para>TLB shootdown is performed in two phases. First, the initiator
process sends an IPI message indicating the TLB shootdown request to
the rest of the CPUs. Then, it waits until all CPUs confirm TLB
invalidating action execution.</para>
</section>
</section>
</section>
<section>
<title>---</title>
<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>
</chapter>