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<?xml version="1.0" encoding="UTF-8"?>
<chapter id="mm">
  <?dbhtml filename="mm.html"?>
 
  <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>Every physical memory frame in a zone, is described by a structure
      that contains number of references and other data used by buddy
      system.</para>
    </section>
 
    <section id="frame_allocator">
      <indexterm>
        <primary>frame allocator</primary>
      </indexterm>
 
      <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 a particular 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="eps">
              <imagedata fileref="images.vector/frame_alloc.eps" format="EPS" />
            </imageobject>
 
            <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">
      <indexterm>
        <primary>buddy system</primary>
      </indexterm>
 
      <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 <emphasis>i</emphasis> contains
      all unallocated blocks of size
      <emphasis>2<superscript>i</superscript></emphasis>. The index
      <emphasis>i</emphasis> is called the order of block. Should there be two
      adjacent equally sized blocks in the list <emphasis>i</emphasis> (i.e.
      buddies), the buddy allocator would coalesce them and put the resulting
      block in list <emphasis>i + 1</emphasis>, provided that the resulting
      block would be naturally aligned. Similarily, when the allocator is
      asked to allocate a block of size
      <emphasis>2<superscript>i</superscript></emphasis>, it first tries to
      satisfy the request from the list with index <emphasis>i</emphasis>. If
      the request cannot be satisfied (i.e. the list <emphasis>i</emphasis> is
      empty), the buddy allocator will try to allocate and split a larger
      block from the list with index <emphasis>i + 1</emphasis>. 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">
            <imageobject role="eps">
              <imagedata fileref="images.vector/buddy_alloc.eps" format="EPS" />
            </imageobject>
 
            <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>
 
          <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">
      <indexterm>
        <primary>slab allocator</primary>
      </indexterm>
 
      <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 commonly used objects are preallocated in continuous
      areas of physical memory called slabs<footnote>
          <para>Slabs are in fact blocks of physical memory frames allocated
          from the frame allocator.</para>
        </footnote>. Whenever an object is to be allocated, the slab allocator
      returns the first available item from a suitable slab corresponding to
      the object type<footnote>
          <para>The mechanism is rather more complicated, see the next
          paragraph.</para>
        </footnote>. Due to the fact that the sizes of the requested and
      allocated object match, the slab allocator significantly reduces
      internal fragmentation.</para>
 
      <indexterm>
        <primary>slab allocator</primary>
 
        <secondary>- slab cache</secondary>
      </indexterm>
 
      <para>Slabs of one object type are organized in a structure called slab
      cache. There are ususally more slabs in the slab cache, depending on
      previous allocations. If the the slab cache runs out of available slabs,
      new slabs are allocated. In order to exploit parallelism and to avoid
      locking of shared spinlocks, slab caches can have variants of
      processor-private slabs called magazines. On each processor, there is a
      two-magazine cache. Full magazines that are not part of any
      per-processor magazine cache are stored in a global list of full
      magazines.</para>
 
      <indexterm>
        <primary>slab allocator</primary>
 
        <secondary>- magazine</secondary>
      </indexterm>
 
      <para>Each object begins its life in a slab. When it is allocated from
      there, the slab allocator calls a constructor that is registered in the
      respective slab cache. The constructor initializes and brings the object
      into a known state. The object is then used by the user. When the user
      later frees the object, the slab allocator puts it into a processor
      private <indexterm>
          <primary>slab allocator</primary>
 
          <secondary>- magazine</secondary>
        </indexterm>magazine cache, from where it can be precedently allocated
      again. Note that allocations satisfied from a magazine are already
      initialized by the constructor. When both of the processor cached
      magazines get full, the allocator will move one of the magazines to the
      list of full magazines. Similarily, when allocating from an empty
      processor magazine cache, the kernel will reload only one magazine from
      the list of full magazines. In other words, the slab allocator tries to
      keep the processor magazine cache only half-full in order to prevent
      thrashing when allocations and deallocations interleave on magazine
      boundaries. The advantage of this setup is that during most of the
      allocations, no global spinlock needs to be held.</para>
 
      <para>Should HelenOS run short of memory, it would start deallocating
      objects from magazines, calling slab cache destructor on them and
      putting them back into slabs. When a slab contanins no allocated object,
      it is immediately freed.</para>
 
      <para>
        <figure>
          <mediaobject id="slab_alloc">
            <imageobject role="eps">
              <imagedata fileref="images.vector/slab_alloc.eps" format="EPS" />
            </imageobject>
 
            <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>
 
          <title>Slab allocator scheme.</title>
        </figure>
      </para>
 
      <section>
        <title>Implementation</title>
 
        <para>The slab allocator is closely modelled after <xref
 
        linkend="Bonwick01" /> with the following exceptions:<itemizedlist>
            <listitem>
              <para>empty slabs are immediately deallocated and</para>
            </listitem>
 
            <listitem>
              <para>empty magazines are deallocated when not needed.</para>
            </listitem>
          </itemizedlist>The following features are not currently supported
        but would be easy to do: <itemizedlist>
            <listitem>cache coloring and</listitem>
 
            <listitem>dynamic magazine grow (different magazine sizes are
            already supported, but the allocation strategy would need to be
            adjusted).</listitem>
          </itemizedlist></para>
 
        <section>
          <title>Allocation/deallocation</title>
 
          <para>The following two paragraphs summarize and complete the
          description of the slab allocator operation (i.e.
          <code>slab_alloc()</code> and <code>slab_free()</code>
          functions).</para>
 
          <formalpara>
            <title>Allocation</title>
 
            <para><emphasis>Step 1.</emphasis> When an allocation request
            comes, the slab allocator checks availability of memory in the
            current magazine of the local processor magazine cache. If the
            available memory is there, the allocator just pops the object from
            magazine and returns it.</para>
 
            <para><emphasis>Step 2.</emphasis> If the current magazine in the
            processor magazine cache is empty, the allocator will attempt to
            swap it with the last magazine from the cache and return to the
            first step. If also the last magazine is empty, the algorithm will
            fall through to Step 3.</para>
 
            <para><emphasis>Step 3.</emphasis> Now the allocator is in the
            situation when both magazines in the processor magazine cache are
            empty. The allocator reloads one magazine from the shared list of
            full magazines. If the reload is successful (i.e. there are full
            magazines in the list), the algorithm continues with Step
            1.</para>
 
            <para><emphasis>Step 4.</emphasis> In this fail-safe step, an
            object is allocated from the conventional slab layer and a pointer
            to it is returned. If also the last magazine is full,</para>
          </formalpara>
 
          <formalpara>
            <title>Deallocation</title>
 
            <para><emphasis>Step 1.</emphasis> During a deallocation request,
            the slab allocator checks if the current magazine of the local
            processor magazine cache is not full. If it is, the pointer to the
            objects is just pushed into the magazine and the algorithm
            returns.</para>
 
            <para><emphasis>Step 2.</emphasis> If the current magazine is
            full, the allocator will attempt to swap it with the last magazine
            from the cache and return to the first step. If also the last
            magazine is empty, the algorithm will fall through to Step
            3.</para>
 
            <para><emphasis>Step 3.</emphasis> Now the allocator is in the
            situation when both magazines in the processor magazine cache are
            full. The allocator tries to allocate a new empty magazine and
            flush one of the full magazines to the shared list of full
            magazines. If it is successfull, the algoritm continues with Step
            1.</para>
 
            <para><emphasis>Step 4. </emphasis>In case of low memory condition
            when the allocation of empty magazine fails, the object is moved
            directly into slab. In the worst case object deallocation does not
            need to allocate any additional memory.</para>
          </formalpara>
        </section>
      </section>
    </section>
  </section>
 
  <section>
    <title>Virtual memory management</title>
 
    <para>Virtual memory is essential for an operating system because it makes
    several things possible. First, it helps to isolate tasks from each other
    by encapsulating them in their private address spaces. Second, virtual
    memory can give tasks the feeling of more memory available than is
    actually possible. And third, by using virtual memory, there might be
    multiple copies of the same program, linked to the same addresses, running
    in the system. There are at least two known mechanisms for implementing
    virtual memory: segmentation and paging. Even though some processor
    architectures supported by HelenOS<footnote>
        <para>ia32 has full-fledged segmentation.</para>
      </footnote> provide both mechanism, the kernel makes use solely of
    paging.</para>
 
    <section id="paging">
      <title>VAT subsystem</title>
 
      <para>In a paged virtual memory, the entire virtual address space is
      divided into small power-of-two sized naturally aligned blocks called
      pages. The processor implements a translation mechanism, that allows the
      operating system to manage mappings between set of pages and set of
      indentically sized and identically aligned pieces of physical memory
      called frames. In a result, references to continuous virtual memory
      areas don't necessarily need to reference continuos area of physical
      memory. Supported page sizes usually range from several kilobytes to
      several megabytes. Each page that takes part in the mapping is
      associated with certain attributes that further desribe the mapping
      (e.g. access rights, dirty and accessed bits and present bit).</para>
 
      <para>When the processor accesses a page that is not present (i.e. its
      present bit is not set), the operating system is notified through a
      special exception called page fault. It is then up to the operating
      system to service the page fault. In HelenOS, some page faults are fatal
      and result in either task termination or, in the worse case, kernel
      panic<footnote>
          <para>Such a condition would be either caused by a hardware failure
          or a bug in the kernel.</para>
        </footnote>, while other page faults are used to load memory on demand
      or to notify the kernel about certain events.</para>
 
      <indexterm>
        <primary>page tables</primary>
      </indexterm>
 
      <para>The set of all page mappings is stored in a memory structure
      called page tables. Some architectures have no hardware support for page
      tables<footnote>
          <para>On mips32, TLB-only model is used and the operating system is
          responsible for managing software defined page tables.</para>
        </footnote> while other processor architectures<footnote>
          <para>Like amd64 and ia32.</para>
        </footnote> understand the whole memory format thereof. Despite all
      the possible differences in page table formats, the HelenOS VAT
      subsystem<footnote>
          <para>Virtual Address Translation subsystem.</para>
        </footnote> unifies the page table operations under one programming
      interface. For all parts of the kernel, three basic functions are
      provided:</para>
 
      <itemizedlist>
        <listitem>
          <para><code>page_mapping_insert()</code>,</para>
        </listitem>
 
        <listitem>
          <para><code>page_mapping_find()</code> and</para>
        </listitem>
 
        <listitem>
          <para><code>page_mapping_remove()</code>.</para>
        </listitem>
      </itemizedlist>
 
      <para>The <code>page_mapping_insert()</code> function is used to
      introduce a mapping for one virtual memory page belonging to a
      particular address space into the page tables. Once the mapping is in
      the page tables, it can be searched by <code>page_mapping_find()</code>
      and removed by <code>page_mapping_remove()</code>. All of these
      functions internally select the page table mechanism specific functions
      that carry out the self operation.</para>
 
      <para>There are currently two supported mechanisms: generic 4-level
      hierarchical page tables and global page hash table. Both of the
      mechanisms are generic as they cover several hardware platforms. For
      instance, the 4-level hierarchical page table mechanism is used by
      amd64, ia32, mips32 and ppc32, respectively. These architectures have
      the following page table format: 4-level, 2-level, TLB-only and hardware
      hash table, respectively. On the other hand, the global page hash table
      is used on ia64 that can be TLB-only or use a hardware hash table.
      Although only two mechanisms are currently implemented, other mechanisms
      (e.g. B+tree) can be easily added.</para>
 
      <section id="page_tables">
        <indexterm>
          <primary>page tables</primary>
 
          <secondary>- hierarchical</secondary>
        </indexterm>
 
        <title>Hierarchical 4-level page tables</title>
 
        <para>Hierarchical 4-level page tables are generalization of the
        frequently used hierarchical model of page tables. In this mechanism,
        each address space has its own page tables. To avoid confusion in
        terminology used by hardware vendors, in HelenOS, the root level page
        table is called PTL0, the two middle levels are called PTL1 and PTL2,
        and, finally, the leaf level is called PTL3. All architectures using
        this mechanism are required to use PTL0 and PTL3. However, the middle
        levels can be left out, depending on the hardware hierachy or
        structure of software-only page tables. The genericity is achieved
        through a set of macros that define transitions from one level to
        another. Unused levels are optimised out by the compiler.</para>
      </section>
 
      <section>
        <indexterm>
          <primary>page tables</primary>
 
          <secondary>- hashing</secondary>
        </indexterm>
 
        <title>Global page hash table</title>
 
        <para>Implementation of the global page hash table was encouraged by
        64-bit architectures that can have rather sparse address spaces. The
        hash table contains valid mappings only. Each entry of the hash table
        contains an address space pointer, virtual memory page number (VPN),
        physical memory frame number (PFN) and a set of flags. The pair of the
        address space pointer and the virtual memory page number is used as a
        key for the hash table. One of the major differences between the
        global page hash table and hierarchical 4-level page tables is that
        there is only a single global page hash table in the system while
        hierarchical page tables exist per address space. Thus, the global
        page hash table contains information about mappings of all address
        spaces in the system.</para>
 
        <para>The global page hash table mechanism uses the generic hash table
        type as described in the chapter dedicated to <link
        linkend="hashtables">data structures</link> earlier in this
        book.</para>
      </section>
    </section>
  </section>
 
  <section id="tlb">
    <indexterm>
      <primary>TLB</primary>
    </indexterm>
 
    <title>Translation Lookaside buffer</title>
 
    <para>Due to the extensive overhead of several extra memory accesses
    during page table lookup that are necessary on every instruction, modern
    architectures deploy fast assotiative cache of recelntly used page
    mappings. This cache is called TLB - Translation Lookaside Buffer - and is
    present on every processor in the system. As it has been already pointed
    out, TLB is the only page translation mechanism for some
    architectures.</para>
 
    <section id="tlb_shootdown">
      <indexterm>
        <primary>TLB</primary>
 
        <secondary>- TLB shootdown</secondary>
      </indexterm>
 
      <title>TLB consistency</title>
 
      <para>The operating system is responsible for keeping TLB consistent
      with the page tables. Whenever mappings are modified or purged from the
      page tables, or when an address space identifier is reused, the kernel
 
 
      needs to invalidate the respective contents of TLB. Some TLB types
 
      support partial invalidation of their content (e.g. ranges of pages or
      address spaces) while other types can be invalidated only entirely. The
      invalidation must be done on all processors for there is one TLB per
 
      processor. Maintaining TLB consistency on multiprocessor configurations
      is not as trivial as it might look from the first glance.</para>
 
      <para>The remote TLB invalidation is called TLB shootdown. HelenOS uses
      a simplified variant of the algorithm described in <xref
 
 
 
      linkend="Black89" />. </para>
 
 
 
 
 
 
 
 
 
 
      <para>TLB shootdown is performed in three phases.</para>
 
      <formalpara>
        <title>Phase 1.</title>
 
        <para>The initiator clears its TLB flag and locks the global TLB
        spinlock. The request is then enqueued into all other processors' TLB
        shootdown message queues. When the TLB shootdown message queue is full
        on any processor, the queue is purged and a single request to
        invalidate the entire TLB is stored there. Once all the TLB shootdown
        messages were dispatched, the initiator sends all other processors an
        interrupt to notify them about the incoming TLB shootdown message. It
        then spins until all processors accept the interrupt and clear their
        TLB flags.</para>
      </formalpara>
 
      <formalpara>
        <title>Phase 2.</title>
 
        <para>Except for the initiator, all other processors are spining on
        the TLB spinlock. The initiator is now free to modify the page tables
        and purge its own TLB. The initiator then unlocks the global TLB
        spinlock and sets its TLB flag.</para>
      </formalpara>
 
      <formalpara>
        <title>Phase 3.</title>
 
        <para>When the spinlock is unlocked by the initiator, other processors
        are sequentially granted the spinlock. However, once they manage to
        lock it, they immediately release it. Each processor invalidates its
        TLB according to messages found in its TLB shootdown message queue. In
        the end, each processor sets its TLB flag and resumes its previous
        operation.</para>
      </formalpara>
    </section>
 
    <section>
      <title>Address spaces</title>
 
      <section>
        <indexterm>
          <primary>address space</primary>
 
          <secondary>- area</secondary>
        </indexterm>
 
        <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>
        <indexterm>
          <primary>address space</primary>
 
          <secondary>- ASID</secondary>
        </indexterm>
 
        <title>Address Space ID (ASID)</title>
 
        <para>Every task in the operating system has it's own view of the
        virtual memory. When performing context switch between different
        tasks, the kernel must switch the address space mapping as well. As
        modern processors perform very aggressive caching of virtual mappings,
        flushing the complete TLB on every context switch would be very
        inefficient. To avoid such performance penalty, some architectures
        introduce an address space identifier, which allows storing several
        different mappings inside TLB.</para>
 
        <para>HelenOS kernel can take advantage of this hardware support by
        having an ASID abstraction. 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>
 
        <indexterm>
          <primary>address space</primary>
 
          <secondary>- ASID stealing</secondary>
        </indexterm>
 
        <para>
          <classname>ASID stealing algoritm here.</classname>
        </para>
      </section>
    </section>
  </section>
</chapter>
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