WebSVN – HelenOS-doc – Diff – /design/trunk/src/ch_memory_management.xml/

 <?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>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">
             <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">
       <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>
       <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>
       <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 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.</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="html">
               <imagedata fileref="images/slab_alloc.png" format="PNG" />
             </imageobject>
           </mediaobject>
           <title>Slab allocator scheme.</title>
         </figure></para>
       <section>
         <title>Implementation</title>
         <para>The slab allocator is closely modelled after OpenSolaris slab
         allocator by Jeff Bonwick and Jonathan Adams with the following
         exceptions:<itemizedlist>
             <listitem>
-              empty slabs are immediately deallocated
+               empty slabs are immediately deallocated
             </listitem>
             <listitem>
               <para>empty magazines are deallocated when not needed</para>
             </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 the allocation strategy would need to be adjusted)
             </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">
+      <section id="paging">
-        <title>Page tables</title>
+        <title>Paging</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
+        <section>
-        <emphasis>page_mapping_operations</emphasis></para>
+          <title>Introduction</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
-        <formalpara>
+          used.</para>
+          <para>HelenOS kernel has two different approaches to the paging
+          implementation: <emphasis>4 level page tables</emphasis> and
+          <emphasis>global hash table</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>
+          <para>Thanks to the abstract paging interface, there was a place
+          left for more paging implementations (besides already implemented
+          hieararchical page tables and hash table), for example B-Tree based
+          page tables.</para>
+        </section>
+        <section>
-          <title>4-level page tables</title>
+          <title>Hierarchical 4-level page tables</title>
-          <para>4-level page tables are the generalization of the hardware
+          <para>Hierarchical 4-level page tables are the generalization of the
+          hardware capabilities of most architectures. Each address space has
-          capabilities of several architectures.<itemizedlist>
+          its own page tables.<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>
+        </section>
+        <section>
+          <title>Global hash table</title>
+          <para>Implementation of the global hash table was encouraged by the
+          ia64 architecture support. One of the major differences between
+          global hash table and hierarchical tables is that global hash table
+          exists only once in the system and the hierarchical tables are
-        <formalpara>
+          maintained per address space.</para>
+          <para>Thus, hash table contains information about all address spaces
+          mappings in the system, so, the hash of an entry must contain
+          information of both address space pointer or id and the virtual
+          address of the page. Generic hash table implementation assumes that
+          the addresses of the pointers to the address spaces are likely to be
+          on the close addresses, so it uses least significant bits for hash;
+          also it assumes that the virtual page addresses have roughly the
+          same probability of occurring, so the least significant bits of VPN
-          <title>Global hash tables</title>
+          compose the hash index.</para>
-          <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>
+          <para>Collision chains ...</para>
-        </formalpara>
+        </section>
-        <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
+          space to schedule a new system task. Moreover, this invalidation
-          address space mappings). Moreover, this invalidation operation must
+          operation must be done an all system CPUs because each CPU has its
-          be done an all system CPUs because each CPU has its own independent
+          own independent TLB cache. Thus maintaining TLB consistency on SMP
-          TLB cache. Thus maintaining TLB consistency on SMP configuration as
+          configuration as not as trivial task as it looks on the first
-          not as trivial task as it looks at the first glance. Naive solution
+          glance. Naive solution would assume that is the CPU which wants to
-          would assume remote TLB invalidatation, which is not possible on the
+          invalidate TLB will invalidate TLB caches on other CPUs. It is not
-          most of the architectures, because of the simple fact - flushing TLB
+          possible on the most of the architectures, because of the simple
-          is allowed only on the local CPU and there is no possibility to
+          fact - flushing TLB is allowed only on the local CPU and there is no
-          access other CPUs' TLB caches.</para>
+          possibility to access other CPUs' TLB caches, thus invalidate TLB
+          remotely.</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.
 -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
+          the architecture is aware of ASIDs, when kernel needs to dump some
-          switching, kernel invalidates only entries from this particular
+          ASID to use by another task, it invalidates only entries from this
-          address space. Final option of the TLB invalidation is the complete
+          particular address space. Final option of the TLB invalidation is
-          TLB cache invalidation, which is the operation that flushes all
+          the complete TLB cache invalidation, which is the operation that
-          entries in TLB.</para>
+          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>TLB shootdown is performed in two phases.</para>
+          <formalpara>
-      <para>At the moment HelenOS does not support swapping.</para>
+            <title>Phase 1.</title>
-      <para>- pouzivame vypadky stranky k alokaci ramcu on-demand v ramci
+            <para>First, initiator locks a global TLB spinlock, then request
+            is being put to the local request cache of every other CPU in the
+            system protected by its spinlock. In case the cache is full, all
-      as_area - na architekturach, ktere to podporuji, podporujeme non-exec
+            requests in the cache are replaced by one request, indicating
+            global TLB flush. Then the initiator thread sends an IPI message
+            indicating the TLB shootdown request to the rest of the CPUs and
+            waits actively until all CPUs confirm TLB invalidating action
+            execution by setting up a special flag. After setting this flag
+            this thread is blocked on the TLB spinlock, held by the
+            initiator.</para>
+          </formalpara>
+          <formalpara>
+            <title>Phase 2.</title>
+            <para>All CPUs are waiting on the TLB spinlock to execute TLB
+            invalidation action and have indicated their intention to the
+            initiator. Initiator continues, cleaning up its TLB and releasing
+            the global TLB spinlock. After this all other CPUs gain and
+            immidiately release TLB spinlock and perform TLB invalidation
-      stranky</para>
+            actions.</para>
+          </formalpara>
+        </section>
+      </section>
     </section>
   </section>
 </chapter>

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