6.828 Fall 2004 Lab 4: Preemptive Multitasking

Handed out Wednesday, October 13
Part A due Thursday, October 21
Part B due Thursday, October 28
Part C due Thursday, November 4


In this lab you will implement preemptive multitasking among multiple simultaneously active user-mode environments. In part A you will first implement round-robin scheduling and the basic process management system calls (calls that create and destroy environments, and allocate/map memory). In part B, you will implement a Unix-like fork() function, which allows one user-mode environment to fork off other, "child" environments, which start off as virtual "clones" of the parent but can subsequently execute independently of the parent. Finally, in part C you will add support for inter-process communication (IPC), allowing different user-mode environments to communicate and synchronize with each other explicitly. You will also add support for hardware clock interrupts and preemption.

Getting Started

Download the lab 4 code from http://pdos.lcs.mit.edu/6.828/2004/labs/lab4/lab4.tar.gz, and unpack it into your 6.828 directory as before. As usual, you will need to merge our new code for this lab into your source tree, or copy the changes that you have made in the previous labs over to our tree.

Lab Requirements

This lab is divided into three parts, A, B, and C. We have allocated one week in the schedule for each part.

As before, you will need to do all of the regular exercises described in the lab and at least one challenge problem. (You do not need to do one challenge problem per part, just one for the whole lab.) Additionally, you will need to write up brief answers to the questions posed in the lab and a short (e.g., one or two paragraph) description of what you did to solve your chosen challenge problem. If you implement more than one challenge problem, you only need to describe one of them in the write-up, though of course you are welcome to do more. Place the write-up in a file called answers.txt (plain text) or answers.html (HTML format) in the top level of your lab2 directory before handing in your work.

Part A: User-level Environment Creation and Cooperative Multitasking

In the first part of this lab, you will implement some new JOS kernel system calls to allow user-level environments to create additional new environments. You will also implement cooperative round-robin scheduling, allowing the kernel to switch from one environment to another when the current environment voluntarily relinquishes the CPU (or exits). Later in part C you will implement preemptive scheduling, which allows the kernel to re-take control of the CPU from an environment after a certain time has passed even if the environment does not cooperate.

Round-Robin Scheduling

Your first task in this lab is to change the JOS kernel so that it does not always just run the environment in envs[0], but instead can alternate between multiple environments in "round-robin" fashion. Round-robin scheduling in JOS works as follows:

Exercise 1. Implement round-robin scheduling in sched_yield() as described above, and implement the crucial register state saving code in env_run(). Don't forget to modify syscall() to dispatch sys_yield() Hint: the symbol UTF defined in kern/trap.h may be useful.

Modify kern/init.c to create two (or more!) environments that all run the program user/yield.c. You should see the environments switch back and forth between each other five times before terminating, at which point the idle process runs and invokes the JOS kernel debugger. If this does not happen or the output looks wrong, then fix your code before proceeding.

Make sure you can answer these questions:
  1. In your implementation of env_run() you should have called lcr3(). But before and after the call to lcr3(), your code makes references (at least it should) to the variable e--the argument to env_run. How can this work?

    Upon loading the %cr3 register, the addressing context used by the MMU is instantly changed. The problem is that a virtual address (namely e) has meaning relative to a given address context--the address context specifies the physical address to which the virtual address map. Why can the pointer e be dereferenced both before and after the addressing switch?

Challenge! Add a less trivial scheduling policy to the kernel, such as a fixed-priority scheduler that allows each environment to be assigned a priority and ensures that higher-priority environments are always chosen in preference to lower-priority environments. If you're feeling really adventurous, try implementing a Unix-style adjustable-priority scheduler or even a lottery or stride scheduler. (Look up "lottery scheduling" and "stride scheduling" in Google.)

Write a test program or two that verifies that your scheduling algorithm is working correctly (i.e., the right environments get run in the right order). It may be easier to write these test programs once you have implemented fork() and IPC in parts B and C of this lab.

Challenge! The JOS kernel currently does not allow applications to use the x86 processor's x87 floating-point unit (FPU), MMX instructions, or Streaming SIMD Extensions (SSE). Extend the Env structure to provide a save area for the processor's floating point state, and extend the context switching code to save and restore this state properly when switching from one environment to another. The FXSAVE and FXRSTOR instructions may be useful, but note that these are not in the old i386 user's manual because they were introduced in more recent processors. Write a user-level test program that does something cool with floating-point.

System Calls for Environment Creation

Although your kernel is now capable of running and switching between multiple user-level environments, it is still limited to running environments that the kernel initially set up. You will now implement the necessary JOS system calls to allow user environments to create and start other new user environments.

Unix provides the fork() system call as its process creation primitive. Unix fork() copies the entire address space of calling process (the parent) to create a new process (the child). The only differences between the two observable from user space are their process IDs and parent process IDs (as returned by getpid and getppid). In the parent, fork() returns the child's process ID, while in the child, fork() returns 0. The two processes do not share any memory: writes to one process's memory do not appear in the other and vice versa.

In JOS we will provide a different, much more primitive set of system system calls for creating new user-mode environments. With these system calls we will be able to implement Unix-like fork() functionality entirely in user space, in addition to other types of environment creation functionality. The new system calls we will use in JOS are as follows:

This system call creates a new environment with an almost blank slate: nothing mapped in the user portion of its address space, and not runnable. The new environment will have the same register state as the parent environment at the time of the sys_env_alloc call. The only difference is that in the parent, sys_env_alloc will return the id of the newly created environment, but in the child it will return 0. (Since the child starts out marked as not runnable, sys_env_alloc will not actually return in the child until the parent has explicitly allowed this by marking the child runnable.)
Sets the user register state to new values specified by the caller in a struct Trapframe. We will not actually use this system call until lab 5, but we introduce it here because of its logical relationship to the task of setting up new user environments.
Sets the status of a specified environment to ENV_RUNNABLE or ENV_NOT_RUNNABLE. This system call is typically used to mark a new environment ready to run, once its address space and register state has been fully initialized.
Allocates a page of physical memory and maps it at a given virtual address in a given environment's address space.
Copy a page mapping (not the contents of a page!) from one environment's address space to another, leaving a memory sharing arrangement in place so that the new and the old mappings can both be used to access the same page of physical memory.
Unmap a page mapped at a given virtual address in a given environment.

For all of the system calls above that accept environment IDs, the JOS kernel supports the special convention that a value of 0 effectively means "the current environment." This convention is implemented by envid2env() in kern/env.cc.

We have provided a very primitive implementation of Unix-like fork() functionality in the test program user/dumbfork.c. This test program uses the above system calls to create and run a child environment with a copy of its own address space. The two environments then switch back and forth using sys_yield as in the previous exercise. The parent exits after 10 iterations, whereas the child exits after 20.

Exercise 2. Implement the system calls described above in kern/syscall.c. You will need to use various functions in kern/pmap.c and kern/env.cc, particularly envid2env(). For now, whenever you call envid2env(), just pass 0 in the checkperm parameter. Be sure you check for any invalid system call arguments, returning -E_INVAL in that case. Test your JOS kernel with user/dumbfork and make sure it works before proceeding.

Challenge! Add the additional system calls necessary to read all of the vital state of an existing environment as well as set it up. Then implement a user mode program that forks off a child process, runs it for a while (e.g., a few iterations of sys_yield()), then takes a complete snaphost or checkpoint of the child process, runs the child for a while longer, and finally restores the child process to the state it was in at the checkpoint and continues it from there. Thus, you are effectively "replaying" the execution of the child process from an intermediate state. Make the child process perform some interaction with the user using sys_cgetc() or readline() so that the user can view and mutate its internal state, and verify that with your checkpoint/restart functionality you can give the child process a case of selective amnesia, making it "forget" everthing that happened beyond a certain point.

Permission Checking

The system calls we introduced above are quite "dangerous" because they currently allow any environment to make modifications at any time to the address space, register state, or scheduling status of any other environment simply by specifying the target environment's ID in the system call. To achieve better protection against buggy or malicious programs, we need to limit the way these dangerous system calls can be used. For this purpose, we will adopt the following simple security convention:
An environment can only use the above system calls to manipulate environments as long as all of the environments affected are either:
  1. The environment making the system call.
  2. Immediate children of the environment making the system call.

Exercise 3. Implement the above security convention by modifying the code in envid2env() appropriately, and by changing the above system calls to pass 1 for the checkperm parameter, thereby enabling this permission check.

This completes Part A of the lab; hand it in using make handin as usual.

Part B: Copy-on-Write Fork

As mentioned earlier, Unix provides the fork() system call as its primary process creation primitive. The fork() system call copies the address space of the calling process (the parent) to create a new process (the child).

V6 Unix implemented fork() by copying the parent's entire data segment into a new memory region allocated for the child. This is essentially the same approach that our dumbfork() function above takes. The copying of the parent's address into the child is generally, by far, the most expensive part of the fork() operation.

However, a call to fork() is frequently followed almost immediately by a call to exec() in the child process, which clears out the child process's address space and loads a new program in its place. (This is what the the shell typically does, for example.) In this case, most of the laborious work done in copying the parent's entire address space into the child is useless, because the child process will probably actually touch very little of its memory before calling exec() to execute a new program.

For this reason, Later versions of Unix took advantage of more advanced virtual memory hardware to allow the parent and child to share the memory mapped into their respective address spaces until one of the processes actually modifies it. This technique is known as copy-on-write. To do this, on fork() the kernel would merely copy the address space mappings from the parent to the child instead of the actual contents of the mapped pages, and at the same time mark the now-shared pages read-only. When one of the two processes tries to write to one of these shared pages, the process takes a page fault. At this point, the Unix kernel realizes that the page was really a "virtual" or "copy-on-write" copy, and so it makes a new, private copy of the page for the faulting process. In this way, the contents of individual pages aren't actually copied until they are actually written to. This optimization makes a fork() followed by an exec() in the child much cheaper: the child will probably only need to copy one page (the current page of its stack) before it calls exec().

In the next piece of this lab, you will implement "proper" Unix-like fork() functionality with copy-on-write, as a user space library routine. Implementing fork() and copy-on-write support in user space has the benefit that the kernel remains much simpler and thus more likely to be correct. It also lets individual user-mode programs define their own semantics for fork(). A program that wants a slightly different implementation (for example, the expensive always-copy version like dumbfork(), or one in which the parent and child actually share memory afterward) can easily provide its own.

User-level page fault handling

The first thing we must do in order to allow user-mode code to implement true copy-on-write fork() functionality is to enable user-mode programs to handle their own page faults instead of depending on the kernel to do so. Copy-on-write is only one of many possible uses for user-level page fault handling, however.

Most page faults encountered during normal program execution are usually fixable. For example, most Unix kernels initially map only a single page in a new process's stack region, and allocate and map additional stack pages later "on demand" as the process's stack consumption increases and causes page faults on stack pages that are not yet mapped. A typical Unix kernel must track, for each user-mode process, the various regions of memory that are mapped and what to do when faults happen in them. For example, a fault in the stack region will typically allocate and map in a new page. A fault in the program's BSS region will typically map in a new page and also make sure it is zeroed. In systems with demand-paged executables, a fault in the text region will read the corresponding page of the binary off of disk and then map it in.

This is a lot of information for the kernel to track and get right. Instead of taking the traditional Unix approach, in JOS we push this fault handling functionality into user space, where bugs are less damaging. This design has the added benefit of allowing programs great flexibility in defining their memory regions. Not only do we get the ability to implement copy-on-write fork(), but we will use user-level page fault handling later for mapping and accessing files on a disk-based file system.

Setting the Page Fault Handler

In order to handle its own page faults, a user environment will need to register a page fault handler entrypoint with the JOS kernel. The user environment registers its page fault entrypoint via the new sys_set_pgfault_entry system call. We have added a new member to the Env structure, env_pgfault_entry, to record this information.

Exercise 4. Implement the sys_set_pgfault_entry system call. Be sure to enable permission checking when looking up the environment ID of the target environment, since this is also a "dangerous" system call.

Normal and Exception Stacks in User Environments

During normal execution, a user environment in JOS will run on the normal user stack: its ESP register starts out pointing at USTACKTOP, and the stack data it pushes resides on the page between USTACKTOP-BY2PG and USTACKTOP-1 inclusive. When a page fault occurs in user mode, however, the kernel will restart the user environment running a designated user-level page fault handler on a different stack, namely the user exception stack. In essence, we will make the JOS kernel implement automatic "stack switching" functionality on behalf of the user environment, in much the same way that the x86 processor already implements stack switching functionality on behalf of JOS when transferring from user mode to kernel mode!

The JOS user exception stack is also one page in size, and its top is defined to be at virtual address UXSTACKTOP, so the valid bytes of the user exception stack are from UXSTACKTOP-BY2PG through UXSTACKTOP-1 inclusive. While running on this exception stack, the user-level page fault handler can use JOS's regular system calls to map new pages or adjust mappings so as to fix whatever problem originally caused the page fault. Then the user-level page fault handler returns, via an assembly language stub, to the faulting code on the original stack.

Each user environment that wants to support user-level page fault handling will need to allocate memory for its own exception stack, using the sys_mem_alloc() system call introduced in part A.

Invoking the User Page Fault Handler

You will now need to change the page fault handling code in kern/trap.c to handle page faults from user mode as follows. We will call the state of the user environment at the time of the fault the trap-time state.

If there is no page fault handler registered, the JOS kernel destroys the user environment with a message as before. Otherwise, the kernel sets up a trap frame on the exception stack that looks like this:

                 	-0  <-- UXSTACKTOP
empty			-4
empty			-8
empty			-12
empty			-16
empty			-20
trap-time eip		-24
trap-time eflags   	-28
trap-time esp		-32
tf_err (error code)	-36
fault_va        	-40 <-- %esp when handler is run

The kernel then arranges for the user environment to resume execution with the page fault handler running on the exception stack with this stack frame. (You must figure out how to make this happen.) Each empty line in the frame above is simply a 32-bit word-size space on the exception stack that the kernel does not initialize, but which the fault handler in the user environment can use. The fault_va is the virtual address at which the page fault occurred.

If the user environment is already running on the user exception stack when an exception occurs, then the page fault handler itself has faulted. In this case, you should start the new stack frame just under the current tf->tf_esp rather than at UXSTACKTOP:

(...existing contents
of exception stack...)
                 	-0  <-- tf->tf_esp
empty			-4
empty			-8
empty			-12
empty			-16
empty			-20
trap-time eip		-24
trap-time eflags   	-28
trap-time esp		-32
tf_err (error code)	-36
fault_va        	-40 <-- %esp when handler is run

To test whether tf->tf_esp is already on the user exception stack, check whether it is in the range between UXSTACKTOP-BY2PG and UXSTACKTOP-1, inclusive.

Exercise 5. Implement the code in kern/trap.c required to dispatch page faults the user-mode handler. Be sure to take appropriate precautions when writing into the exception stack. (What happens if the user environment runs out of space on the exception stack?)

User-mode Page Fault Entrypoint

Next, you need to implement the assembly routine that will take care of calling the C page fault handler and resume execution at the original faulting instruction. This assembly routine is the handler that will be registered with the kernel using sys_set_pgfault_entry().

Exercise 6. Implement the _pgfault_entry routine in lib/pfentry.S. There is a commented outline there to help you along. You may find it useful to reread the description from the beginning of this section as well.

Three of the empty words in the stack frame created by the kernel above are for the assembly routine to use to save the caller-save registers before moving on to C code. The other two are important for the recursive fault case. Draw a stack diagram for the recursive case and carefully examine how each word gets used at each point during the execution of the fault handler in order to figure out how to use these words correctly.

Finally, you need to implement the C user library side of the user-level page fault handling mechanism.

Exercise 7. Finish set_pgfault_handler() in lib/pgfault.c.


Change kern/init.c to run user/faultread. Build your kernel and run it. You should see:
[00000000] new env 00000800
[00000000] new env 00001001
[00001001] user fault va 00000000 ip 0080003a
TRAP frame ...
[00001001] free env 00001001

Change kern/init.c to run user/faultdie. Build your kernel and run it. You should see:

[00000000] new env 00000800
[00000000] new env 00001001
i faulted at va deadbeef, err 6
[00001001] exiting gracefully
[00001001] free env 00001001

Change kern/init.c to run user/faultalloc. Build your kernel and run it. You should see:

[00000000] new env 00000800
[00000000] new env 00001001
fault deadbeef
this string was faulted in at deadbeef
fault cafebffe
fault cafec000
this string was faulted in at cafebffe
[00001001] exiting gracefully
[00001001] free env 00001001

If you see only the first "this string" line, it means you are not handling recursive page faults properly.

Change kern/init.c to run user/faultallocbad. Build your kernel and run it. You should see:

[00000000] new env 00000800
[00000000] new env 00001001
[00001001] PFM_KILL va deadbeef ip f010263d
TRAP frame ...
[00001001] free env 00001001
(Your ip may differ from ours but should begin f01.)

Make sure you understand why user/faultalloc and user/faultallocbad behave differently.

Challenge! Extend your kernel so that not only page faults, but all types of processor exceptions that code running in user space can generate, can be redirected to a user-mode exception handler. Write user-mode test programs to test user-mode handling of various exceptions such as divide-by-zero, general protection fault, and illegal opcode.

Implementing Copy-on-Write Fork

We now have the facilities to implement full copy-on-write fork functionality, entirely in user space.

We have provided a skeleton for your fork() function in lib/fork.c. Like dumbfork(), fork() creates a new environment, then scans through the parent environment's entire address space and sets up corresponding page mappings in the child. The key difference is that, while dumbfork() copied entire pages, fork() will initially only copy page mappings. Notice that the duppage() helper function in dumbfork calls sys_mem_alloc() to allocate a new page of physical memory for each page in the parent, and then calls memcpy() to copy the contents of the parent's page into the child's new page. These calls to memcpy() represent the bulk of the time dumbfork() takes to run, and so fork() attempts to "optimize away" most of this page copying by copying pages lazily only when they are actually modified.

The basic control flow for fork() is as follows:

  1. The parent installs pgfault() as the C-level page fault handler, using the set_pgfault_handler() function you implemented above.
  2. The parent calls sys_env_alloc() to allocate a child environment.
  3. For each writable or copy-on-write page in its address space below UTOP, the parent maps the page copy-on-write into the address space of the child and then remaps the page copy-on-write in its own address space.

    The exception stack is not remapped this way, however. Instead you need to allocate a fresh page in the child for the exception stack. Since the page fault handler will be doing the actual copying and the page fault handler runs on the exception stack, the exception stack cannot be made copy-on-write: who would copy it?

  4. The parent sets the user page fault entrypoint for the child to look like its own.
  5. The child is now ready to run, so the parent marks it runnable.

After the fork, both processes will take page faults when the code they run attempts to write to a page that hasn't been copied yet. Here's the control flow for the user page fault handler:

  1. Kernel propagates page fault to _pgfault_entry, which calls fork()'s pgfault() handler.
  2. pgfault() checks that the fault is a write (check FEC_WR) and that the PTE for the page is marked PTE_COW (copy-on-write). If not, panic.
  3. pgfault() allocates a new page mapped at a temporary location and copies the contents of the faulting page contents into it. Then the fault handler maps the new page at the appropriate address with read/write permissions, in place of the old read-only mapping.

Exercise 8. Implement fork and pgfault in user/fork.c.

Test your code with the forktree program. It should produce the following messages, with interspersed 'new env', 'free env', and 'exiting gracefully' messages:

	1001: I am ''
	1802: I am '0'
	2801: I am '00'
	3802: I am '000'
	2003: I am '1'
	5001: I am '11'
	4802: I am '10'
	6801: I am '100'
	5803: I am '110'
	3004: I am '01'
	8001: I am '011'
	7803: I am '010'
	4005: I am '001'
	6006: I am '111'
	7007: I am '101'

Challenge! Implement a shared-memory fork called sfork. This version should have the parent and child sharing all their memory pages (writes in one environment appear in the other) except for pages in the stack area, which should be treated in the usual copy-on-write manner. Modify user/forktree.c to use sfork() instead of regular fork(). Also, once you have finished implementing IPC in part C, use your sfork() to run user/pingpongs. You will have to find a new way to provide the functionality of the global env pointer.

Challenge! Your implementation of fork makes a huge number of system calls. On the x86, switching into the kernel has non-trivial cost. Augment the system call interface so that it is possible to send a batch of system calls at once. Then change fork to use this interface.

How much faster is your new fork?

You can answer this (roughly) by using analytical arguments to estimate how much of an improvement batching system calls will make to the performance of your fork: How expensive is an int 0x30 instruction? How many times do you execute int 0x30 in your fork? Is accessing the TSS stack switch also expensive? And so on...

Alternatively, you can boot your kernel on real hardware and really benchmark your code. See the RDTSC (read time-stamp counter) instruction, defined in the IA32 manual, which counts the number of clock cycles that have elapsed since the last processor reset. Bochs doesn't emulate this instruction faithfully.

This ends part B. As usual, you can grade your submission with gmake grade and hand it in with gmake handin.

Part C: Preemptive Multitasking and Inter-Process communication (IPC)

In the final part of lab 4 we will enhance the JOS kernel's support for multiple environments by allowing the kernel to preempt uncoopoerative environments and by allowing environments to pass messages to each other explicitly.

Clock Interrupts and Preemption

Modify kern/init.c to run the user/spin test program. This test program forks off a child process, which simply spins forever in a tight loop once it receives control of the CPU. Neither the parent process nor the kernel ever regains the CPU. This is obviously not an ideal situation in terms of protecting the system from bugs or malicious code in user-mode environments, because any user-mode environment can bring the whole system to a halt simply by getting into an infinite loop and never giving back the CPU. In order to allow the kernel to preempt a running environment, forcefully retake control of the CPU from it, we must extend the JOS kernel to support external hardware interrupts.

Interrupt discipline

External interrupts (i.e., device interrupts) are refered as IRQs. There are 16 possible IRQs, numbered 0 through 15. The mapping from IRQ number to IDT entry is not fixed. Pic_init in picirq.c maps IRQs 0-15 to interrupts IRQ_OFFSET through IRQ_OFFSET+15.

In kern/picirq.h, IRQ_OFFSET is defined to be decimal 32. Thus the IDT entries 32-44 correspond to the IRQs 0-15. For example, the clock interrupt is IRQ 0. Thus, IDT[32] contains the address of the clock's interrupt handler routine in the kernel. This IRQ_OFFSET is chosen so that the device interrupts do not overlap with the processor exceptions, which could obviously cause confusion. (In fact, in the early days of PCs running MS-DOS, the IRQ_OFFSET effectively was zero, which indeed caused massive confusion between handling hardware interrupts and handling processor exceptions!)

In JOS, we make a key simplification compared to V6 Unix. External device interrupts are always disabled when in the kernel and always enabled when in user space. External interrupts are controlled by the FL_IF flag bit of the %eflags register (see inc/mmu.h). When this bit is set, external interrupts are enabled. While the bit can be modified in several ways, because of our simplification, we will handle it solely through the process of saving and restoring %eflags register as we enter and leave user mode.

You will have to ensure that the FL_IF flag is set in user processes when they run so that when an interrupt arrives, it gets passed through to the processor and handled by your interrupt code. Otherwise, interrupts are masked, or ignored until interrupts are re-enabled. Interrupts are masked by default after processor reset, and so far we have simply never gotten around to enabling them.

Exercise 9. Modify kern/trapentry.S and kern/trap.c to initialize the appropriate entries in the IDT and provide handlers for IRQs 0 through 15. Then modify the code in env_alloc() and sys_set_trapframe() to ensure that user environments are always run with interrupts enabled.

The processor never pushes an error code or checks the Descriptor Privilege Level (DPL) of the IDT entry when invoking a hardware interrupt handler. You might want to re-read section 9.2 of the 80386 Reference Manual, or section 5.8 of the IA-32 Intel Architecture Software Developer's Manual, Volume 3, at this time.

After doing this exercise, if you run your kernel with any test program that runs for a non-trivial length of time (e.g., dumbfork), you should see a kernel panic shortly into the program's execution. This is because our code has set up the clock hardware to generate clock interrupts, and interrupts are now enabled in the processor, but JOS isn't yet handling them.

Handling Clock Interrupts

In the user/spin program, after the child environment was first run, it just spun in a loop, and the kernel never got control back. We need to program the hardware to generate clock interrupts periodically, which will force control back to the kernel where we can switch control to a different user environment.

The calls to pic_init and kclock_init (from i386_init in init.c), which we have written for you, set up the clock and the interrupt controller to generate interrupts. You now need to write the code to handle these interrupts.

Exercise 10. Modify the kernel's trap() function so that it calls sched_yield() to find and run a different environment whenever a clock interrupt takes place.

You should now be able to get the user/spin test to work: the parent process should fork off the child, sys_yield() to it a couple times but in each case regain control of the CPU after one time slice, and finally kill the child process and terminate gracefully.

Make sure you can answer the following questions:

  1. How many instruction of user code are executed between each interrupt?
  2. How many instructions of kernel code are executed to handle the interrupt?
    Hint: use the vb command mentioned earlier.

Inter-Process communication (IPC)

(Technically in JOS this is "inter-environment communication" or "IEC", but everyone else calls it IPC, so we'll use the standard term.)

We've been focusing on the isolation aspects of the operating system, the ways it provides the illusion that each program has a machine all to itself. Another important service of an operating system is to allow programs to communicate with each other when they want to. It can be quite powerful to let programs interact with other programs. The Unix pipe model is the canonical example.

There are many models for interprocess communication. Even today there are still debates about which models are better for various reasons. We won't get into that debate. Instead, we'll implement a simple IPC mechanism and then try it out.


You will implement a few additional JOS kernel system calls that collectively provide a simple interprocess communication mechanism. You will implement two system calls, sys_ipc_recv and sys_ipc_can_send. Then you will implement two library wrappers ipc_recv and ipc_send.

The "messages" that user environments can send to each other using JOS's IPC mechanism consist of two components: a single 32-bit value, and optionally a single page mapping. Allowing environments to pass page mappings in messages provides an efficient way to transfer more data than will fit into a single 32-bit integer, and also allows environments to set up shared memory arrangements easily.

Sending and Receiving Messages

To receive a message, an environment calls sys_ipc_recv. This system call deschedules the current environment and does not run it again until a message has been received. When an environment is waiting to receive a message, any other environment can send it a message - not just a particular environment, and not just environments that have a parent/child arrangement with the receiving environment. In other words, the permission checking that you implemented in Part A will not apply to IPC, because the IPC system calls are carefully designed so as to be "safe": an environment cannot cause another environment to malfunction simply by sending it messages (unless the target environment is also buggy).

To try to send a value, an environment calls sys_ipc_can_send with both the receiver's environment id and the value to be sent. If the named environment is actually receiving (it has called sys_ipc_recv and not gotten a value yet), then the send delivers the message and returns 0. Otherwise the send returns -E_IPC_NOT_RECV to indicate that the target environment is not currently expecting to receive a value.

A library function ipc_recv in user space will take care of calling sys_ipc_recv and then looking up the information about the received values in the current environment's struct Env.

Similarly, a library function ipc_send will take care of repeatedly calling sys_ipc_can_send until the send succeeds.

Transferring Pages

When an environment calls sys_ipc_recv with a nonzero dstva parameter, the environment is stating that it is willing to receive a page mapping. If the sender sends a page, then that page should be mapped at dstva in the receiver's address space. If the receiver already had a page mapped at dstva, then that previous page is unmapped.

When an environment calls sys_ipc_can_send with a nonzero srcva, it means the sender wants to send the page currently mapped at srcva to the receiver, with permissions perm. After a successful IPC, the sender keeps its original mapping for the page at srcva in its address space, but the receiver also obtains a mapping for this same physical page at the dstva originally specified by the receiver, in the receiver's address space. As a result this page becomes shared between the sender and receiver.

If either the sender or the receiver does not indicate that a page should be transferred, then no page is transferred. After any IPC the kernel sets the new field env_ipc_perm in the receiver's Env structure to the permissions of the page received, or zero if no page was received.

Implementing IPC

Exercise 11. Implement sys_ipc_recv and sys_ipc_can_send in kern/syscall.c. When you call envid2env in these routines, you should set the checkperm flag to 0, meaning that any environment is allowed to send IPC messages to any other environment, and the kernel does no special permission checking other than verifying that the target envid is valid.

Use the user/pingpong and user/primes functions to test your IPC mechanism. You might find it interesting to read user/primes.c to see all the forking and IPC going on behind the scenes.

Challenge! The ipc_send function is not very fair. Run three copies of user/fairness and you will see this problem. The first two copies are both trying to send to the third copy, but only one of them will ever succeed. Make the IPC fair, so that each copy has approximately equal chance of succeeding.

Challenge! Why does ipc_send have to loop? Change the system call interface so it doesn't have to. Make sure you can handle multiple environments trying to send to one environment at the same time.

Challenge! The prime sieve is only one neat use of message passing between a large number of concurrent programs. Read C. A. R. Hoare, ``Communicating Sequential Processes,'' Communications of the ACM 21(8) (August 1978), 666-667, and implement the matrix multiplication example.

Challenge! Probably the most impressive example of the power of message passing is Doug McIlroy's power series calculator, described in M. Douglas McIlroy, ``Squinting at Power Series,'' Software--Practice and Experience, 20(7) (July 1990), 661-683. Implement his power series calculator and compute the power series for sin(1+x^2).

Challenge! Make JOS's IPC mechanism more efficient by applying some of the techniques from Liedtke's paper, "Improving IPC by Kernel Design", or any other tricks you may think of. Feel free to modify the kernel's system call API for this purpose, as long as your code is backwards compatible with what our grading scripts expect. -->

Challenge! Generalize the JOS IPC interface so it is more like L4's, supporting more complex message formats.

This ends part C. As usual, you can grade your submission with gmake grade and hand it in with gmake handin. If you are trying to figure out why a particular test case is failing, run sh grade.sh -v, which will show you the output of the kernel builds and Bochs runs for each test, until a test fails. When a test fails, the script will stop, and then you can inspect bochs.out to see what the kernel actually printed.

Version: $Revision: 1.10 $. Last modified: $Date: 2004/11/04 20:13:37 $