Operating Systems

4.6: Design issues for (demand) Paging Systems

4.6.1: Local vs Global Allocation Policies

A local PRA is one is which a victim page is chosen among the pages of the same process that requires a new page. That is the number of pages for each process is fixed. So LRU for a local policy means the page least recently used by this process. A global policy is one in which the choice of victim is made among all pages of all processes.

• Of course we can't have a purely local policy, why?
  Answer: A new process has no pages and even if we didn't apply this for the first page loaded, the process would remain with only one page.
• Perhaps wait until a process has been running a while or give the process an initial allocation based on the size of the executable.

If we apply global LRU indiscriminately with some sort of RR processor scheduling policy, and memory is somewhat over-committed, then by the time we get around to a process, all the others have run and have probably paged out this process.

If this happens each process will need to page fault at a high rate; this is called thrashing.

It is therefore important to get a good idea of how many pages a process needs, so that we can balance the local and global desires. The working set size w(t,ω) is good for this.

An approximation to the working set policy that is useful for determining how many frames a process needs (but not which pages) is the Page Fault Frequency (PFF) algorithm.

• For each process keep track of the page fault frequency, which is the number of faults divided by the number of references.
• Actually, must use a window or a weighted calculation since you are really interested in the recent page fault frequency.
• If the PFF is too high, allocate more frames to this process. Either
  1. Raise its number of frames and use a local policy; or
  2. Bar its frames from eviction (for a while) and use a global policy.
• What if there are not enough frames?
  Answer: Reduce the MPL (see next section).

As mentioned above a question arises what to do if the sum of the working set sizes exceeds the amount of physical memory available. This question is similar to the final point about PFF and brings us to consider controlling the load (or memory pressure).

4.6.2: Load Control

To reduce the overall memory pressure, we must reduce the multiprogramming level (or install more memory while the system is running, which is hardly practical). That is, we have a connection between memory management and process management. These are the suspend/resume arcs we saw way back when.

4.6.3: Page size

• Page size “must” be a multiple of the disk block size. Why?
  Answer: When copying out a page if you have a partial disk block, you must do a read/modify/write (i.e., 2 I/Os).
• Important property of I/O that we will learn later this term is that eight I/Os each 1KB takes considerably longer than one 8KB I/O
• Characteristics of a large page size.
  • Good for demand paging I/O.
    • Better to swap in/out one big page than several small pages.
    • But if page is too big you will be swapping in data that is really not local and hence might well not be used.
  • Large internal fragmentation (1/2 page size).
  • Small page table.
  • A very large page size leads to very few pages. Process will have many faults if using demand paging and the process frequently references more regions than the number of (large) frames that the process has been allocated.
  • Possibly good for user I/O.
    • If I/O done using physical addresses, then an I/O crossing a page boundary is not contiguous and hence requires multiple actual I/Os. A large page size makes it less likely that a single user I/O will span multiple pages.
    • If I/O uses virtual addresses, then page size doesn't effect this aspect of I/O. That is, the addresses are contiguous in virtual address and hence one I/O is done.
• A small page size has the opposite characteristics.

Homework: Consider a 32-bit address machine using paging with 8KB pages and 4 byte PTEs. How many bits are used for the offset and what is the size of the largest page table? Repeat the question for 128KB pages. Repeat both parts if the system also has segmentation with at most 128 segments.

4.6.4: Separate Instruction and Data (I and D) Spaces

Skipped.

4.6.5: Shared pages

Permit several processes to each have a page loaded in the same frame. Of course this can only be done if the processes are using the same program and/or data.

• Really should share segments.
• Must keep reference counts or something so that when a process terminates, pages (even dirty pages) it shares with another process are not automatically discarded.
• Similarly, a reference count would make a widely shared page (correctly) look like a poor choice for a victim.
• A good place to store the reference count would be in a structure pointed to by both PTEs. If stored in the PTEs themselves, we must keep somehow keep the count consistent between processes.

Homework: 33

4.6.6: Cleaning Policy (Paging Daemons)

Done earlier

4.6.7: Virtual Memory Interface

Skipped.

4.7: Implementation Issues

4.7.1: Operating System Involvement with Paging

1. Process creation. OS must guess at the size of the process and then allocate a page table and a region on disk to hold the pages that are not memory resident. A few pages of the process must be loaded.
2. Ready→Running transition by the scheduler. Real memory must be allocated for the page table if the table has been swapped out (which is permitted when the process is not running). Some hardware register(s) must be set to point to the page table. (There can be many page tables resident, but the hardware must be told the location of the page table for the running process--the "active" page table.
3. Page fault. Lots of work. See 4.7.2 just below.
4. Process termination. Free the page table and the disk region for swapped out pages.

4.7.2: Page Fault Handling

1. The hardware detects the fault and traps to the kernel (switches to supervisor mode and saves state).
   
   
2. Some assembly language code save more state, establishes the C-language (or another programming language) environment, and “calls” the OS.
   
   
3. The OS determines that a page fault occurred and which page was referenced.
   
   
4. If the virtual address is invalid, process A is killed. If the virtual address is valid, the OS must find a free frame. If there is no free frames, the OS selects a victim frame. Call the process owning the victim frame, process B. (If the page replacement algorithm is local, the victim is process A.)
   
   
5. The PTE of the victim page is updated to show that the page is no longer resident.
   
   
6. If the victim page is dirty, the OS schedules an I/O write to copy the frame to disk and blocks A waiting for this I/O to occur.
   
   
7. Assuming process A needed to be blocked (i.e., the victim page is dirty) the scheduler is invoked to perform a context switch.
   • Tanenbaum “forgot” some here.
   • The process selected by the scheduler (say process C) runs.
   • Perhaps C is preempted for D or perhaps C blocks and D runs and then perhaps D is blocked and E runs, etc.
   • When the I/O to write the victim frame completes, a disk interrupt occurs. Assume processes C is running at the time.
   • Hardware trap / assembly code / OS determines I/O done.
   • The scheduler marks A as ready.
   • The scheduler picks a process to run, maybe A, maybe B, maybe C, maybe another processes.
   • At some point the scheduler does pick process A to run. Recall that at this point A is still executing OS code.
   
   
8. Now the O/S has a free frame (this may be much later in wall clock time if a victim frame had to be written). The O/S schedules an I/O to read the desired page into this free frame. Process A is blocked (perhaps for the second time) and hence the process scheduler is invoked to perform a context switch.
   
   
9. Again, another process is selected by the scheduler as above and eventually a Disk interrupt occurs when the I/O completes (trap / asm / OS determines I/O done). The PTE in process A is updated to indicate that the page is in memory.
   
   
10. The O/S may need to fix up process A (e.g. reset the program counter to re-execute the instruction that caused the page fault).
    
    
11. Process A is placed on the ready list and eventually is chosen by the scheduler to run. Recall that process A is executing O/S code.
    
    
12. The OS returns to the first assembly language routine.
    
    
13. The assembly language routine restores registers, etc. and “returns” to user mode.

The user's program running as process A is unaware that all this happened (except for the time delay).

4.7.3: Instruction Backup

A cute horror story. The 68000 was so bad in this regard that early demand paging systems for the 68000, used two processors one running one instruction behind. If the first got a page fault, there wasn't always enough information to figure out what to do so the system switched to the second processor after it did the page fault. Don't worry about instruction backup. Very machine dependent and modern implementations tend to get it right. The next generation machine, 68010, provided extra information on the stack so the horrible 2-processor kludge was no longer necessary.

4.7.4: Locking (Pinning) Pages in Memory

We discussed pinning jobs already. The same (mostly I/O) considerations apply to pages.

4.7.5: Backing Store

The issue is where on disk do we put pages.

• For program text, which is presumably read only, a good choice is the file executable itself.
  
  
• What if we decide to keep the data and stack each contiguous on the backing store. Data and stack grow so we must be prepared to grow the space on disk, which leads to the same issues and problems as we saw with MVT.
  
  
• If those issues/problems are painful, we can scatter the pages on the disk.
  • That is we employ paging!
  • This is NOT demand paging.
  • Need a table to say where the backing space for each page is located.
    • This corresponds to the page table used to tell where in real memory a page is located.
    • The format of the “memory page table” is determined by the hardware since the hardware modifies/accesses it. It is machine dependent.
    • The format of the “disk page table” is decided by the OS designers and is machine independent.
    • If the format of the memory page table was flexible, then we might well keep the disk information in it as well. But often (usually?) the format is not flexible and this is not done.
  
  
• What if we felt disk space was too expensive and wanted to put some of these disk pages on say tape?
  Ans: We use demand paging of the disk blocks! That way "unimportant" disk blocks will migrate out to tape and are brought back in if needed.
  Since a tape read requires seconds to complete (because the request is not likely to be for the sequentially next tape block), it is crucial that we get very few disk block faults.

Homework: Assume every instruction takes 0.1 microseconds to execute providing it is memory resident. Assume a page fault takes 10 milliseconds to service providing the necessary disk block is actually on the disk. Assume a disk block fault takes 10 seconds service. So the worst case time for an instruction is 10.0100001 seconds. Finally assume the program requires that a billion instructions be executed.

1. If the program is always completely resident, how long does it take to execute?
2. If 0.1% of the instructions cause a page fault, but all the disk blocks are on the disk, how long does the program take to execute and what percentage of the time is the program waiting for a page fault to complete?
3. If 0.1% of the instructions cause a page fault and 0.1% of the page faults cause a disk block fault, how long does the program take to execute, what percentage of the time is the program waiting for a disk block fault to complete?

4.7.6: Separation of Policy and Mechanism

Skipped.

4.8: Segmentation

Up to now, the virtual address space has been contiguous.

• Among other issues this makes memory management difficult when there are more that two dynamically growing regions.
• With two regions you start them on opposite sides of the virtual space as we did before.
• Better is to have many virtual address spaces each starting at zero.
• This split up is user visible.
• Without segmentation (equivalently said with just one segment) all procedures are packed together so if one changes in size all the virtual addresses following are changed and the program must be re-linked. With multiple segments this relinking would be limited to the symbols defined or used in the modified procedure.
• Eases flexible protection and sharing (share a segment). For example, can have a shared library.

Homework: 37.

** Two Segments

Late PDP-10s and TOPS-10

• One shared text segment, that can also contain shared (normally read only) data.
• One (private) writable data segment.
• Permission bits on each segment.
• Which kind of segment is better to evict?
  • Swapping out the shared segment hurts many tasks.
  • The shared segment is read only (probably) so no writeback is needed.
• “One segment” is OS/MVT done above.

** Three Segments

Traditional (early) Unix shown at right.

• Shared text marked execute only.
• Data segment (global and static variables).
• Stack segment (automatic variables).
• In reality, since the text doesn't grow, this was sometimes treated as 2 segments by combining text and data into one segment

** Four Segments

Just kidding.

** General (not necessarily demand) Segmentation

• Permits fine grained sharing and protection. For a simple example can share the text segment in early unix.
• Visible division of program.
• Variable size segments.
• Virtual Address = (seg#, offset).
• Does not mandate how stored in memory.
  • One possibility is that the entire program must be in memory in order to run it. Use whole process swapping. Very early versions of Unix did this.
  • Can also implement demand segmentation.
  • Can combine with demand paging (done below).
• Requires a segment table with a base and limit value for each segment. Similar to a page table. Why is there no limit value in a page table?
  Ans: All pages are the same size so the limit is obvious.
• Entries are called STEs, Segment Table Entries.
• (seg#, offset) --> if (offset<limit) base+offset else error.
• Segmentation exhibits external fragmentation, just as whole program swapping. Since segments are smaller than programs (several segments make up one program), the external fragmentation is not as bad.

** Demand Segmentation

Same idea as demand paging, but applied to segments.

• If a segment is loaded, base and limit are stored in the STE and the valid bit is set in the STE.
• The STE is accessed for each memory reference (not really, TLB).
• If the segment is not loaded, the valid bit is unset. The base and limit as well as the disk address of the segment is stored in the an OS table.
• A reference to a non-loaded segment generate a segment fault (analogous to page fault).
• To load a segment, we must solve both the placement question and the replacement question (for demand paging, there is no placement question).
• I believe demand segmentation was once implemented by Burroughs, but am not sure. It is not used in modern systems.

The following table mostly from Tanenbaum compares demand paging with demand segmentation.

Consideration	Demand Paging	Demand Segmentation
Programmer aware	No	Yes
How many addr spaces	1	Many
VA size > PA size	Yes	Yes
Protect individual procedures separately	No	Yes
Accommodate elements with changing sizes	No	Yes
Ease user sharing	No	Yes
Why invented	let the VA size exceed the PA size	Sharing, Protection, independent addr spaces
Internal fragmentation	Yes	No, in principle
External fragmentation	No	Yes
Placement question	No	Yes
Replacement question	Yes	Yes