Lab 4: Preemptive Multitasking¶
- Handed out: Thursday, February 7, 2019
- Part A due: Thursday, February 14, 2019 (1 week)
- Part B due: Thursday, February 28, 2019 (2 weeks)
- Part C due: SUNDAY, March 10, 2019 (Extended!)
In this lab you will implement preemptive multitasking among multiple simultaneously active user-mode environments.
- In part A, you will add multiprocessor support to JOS, implement round-robin scheduling, and add basic environment management system calls (calls that create and destroy environments, and allocate/map memory).
- In part B, you will implement a Unix-like
fork(), which allows a user-mode environment to create copies of itself.
- 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.
Use Git to commit your Lab 3 source, fetch the latest version of the
course repository, and then create a local branch called
on our lab4 branch,
$ cd ~/jos $ git pull Already up-to-date. $ git checkout -b lab4 origin/lab4 Branch lab4 set up to track remote branch refs/remotes/origin/lab4. Switched to a new branch "lab4" $ git merge lab3 Merge made by recursive. ... $
Lab 4 contains a number of new source files, some of which you should browse before you start:
||Kernel-private definitions for multiprocessor support|
||Code to read the multiprocessor configuration|
||Kernel code driving the local APIC unit in each processor|
||Assembly-language entry code for non-boot CPUs|
||Kernel-private definitions for spin locks, including the big kernel lock|
||Kernel code implementing spin locks|
||Code skeleton of the scheduler that you are about to implement|
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 provide a writeup that briefly answers
to the questions posed in the lab.
Place the write-up in a file called
answers-lab4.txt in the top level of your repository directory before
handing in your work.
Part A: Multiprocessor Support and Cooperative Multitasking¶
In the first part of this lab, you will first extend JOS to run on a multiprocessor system, and then 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.
We are going to make JOS support “symmetric multiprocessing” (SMP), a multiprocessor model in which all CPUs have equivalent access to system resources such as memory and I/O buses. While all CPUs are functionally identical in SMP, during the boot process they can be classified into two types: the bootstrap processor (BSP) is responsible for initializing the system and for booting the operating system; and the application processors (APs) are activated by the BSP only after the operating system is up and running. Which processor is the BSP is determined by the hardware and the BIOS. Up to this point, all your existing JOS code has been running on the BSP.
In an SMP system, each CPU has an accompanying local APIC (LAPIC) unit.
The LAPIC units are responsible for delivering interrupts throughout the
system. The LAPIC also provides its connected CPU with a unique
identifier. In this lab, we make use of the following basic
functionality of the LAPIC unit (in
- Reading the LAPIC identifier (APIC ID) to tell which CPU our code is
currently running on (see
- Sending the
STARTUPinterprocessor interrupt (IPI) from the BSP to the APs to bring up other CPUs (see
- In part C, we program LAPIC’s built-in timer to trigger clock
interrupts to support preemptive multitasking (see
A processor accesses its LAPIC using memory-mapped I/O (MMIO). In MMIO,
a portion of physical memory is hardwired to the registers of some I/O
devices, so the same load/store instructions typically used to access
memory can be used to access device registers. You’ve already seen one
IO hole at physical address
0xA0000 (we use this to write to the VGA
display buffer). The LAPIC lives in a hole starting at physical address
0xFE000000 (32MB short of 4GB), so it’s too high for us to access
using our usual direct map at
KERNBASE. The JOS virtual memory map
leaves a 4MB gap at
MMIOBASE so we have a place to map devices like
this. Since later labs introduce more MMIO regions, you’ll write a
simple function to allocate space from this region and map device memory
kern/pmap.c. To see how
this is used, look at the beginning of
kern/lapic.c. You’ll have to do the next exercise, too, before the
mmio_map_region() will run.
Application Processor Bootstrap¶
Before booting up APs, the BSP should first collect information about
the multiprocessor system, such as the total number of CPUs, their APIC
IDs and the MMIO address of the LAPIC unit. The
kern/mpconfig.c retrieves this information by reading the MP
configuration table that resides in the BIOS’s region of memory.
boot_aps() function (in
kern/init.c) drives the AP bootstrap
process. APs start in real mode, much like how the bootloader started in
boot_aps() copies the AP entry code
kern/mpentry.S) to a memory location that is addressable in the
real mode. Unlike with the bootloader, we have some control over where
the AP will start executing code; we copy the entry code to
MPENTRY_PADDR), but any unused, page-aligned physical address below
640KB would work.
boot_aps() activates APs one after another, by sending
STARTUP IPIs to the LAPIC unit of the corresponding AP, along with
CS:IP address at which the AP should start running its
entry code (
MPENTRY_PADDR in our case). The entry code in
kern/mpentry.S is quite similar to that of
some brief setup, it puts the AP into protected mode with paging
enabled, and then calls the C setup routine
mp_main() (also in
boot_aps() waits for the AP to signal a
CPU_STARTED flag in
cpu_status field of its
before going on to wake up the next one.
and the assembly code in
kern/mpentry.S. Make sure you understand
the control flow transfer during the bootstrap of APs. Then modify your
kern/pmap.c to avoid adding the
MPENTRY_PADDR to the free list, so that we can safely copy
and run AP bootstrap code at that physical address. Your code should
pass the updated
check_page_free_list() test (but might fail the
check_kern_pgdir() test, which we will fix soon).
kern/mpentry.S side by side with
in mind that
kern/mpentry.S is compiled and linked to run above
KERNBASE just like everything else in the kernel, what is the
purpose of macro
MPBOOTPHYS? Why is it necessary in
kern/mpentry.S but not in
boot/boot.S? In other words, what
could go wrong if it were omitted in
Hint: recall the differences between the link address and the load
address that we have discussed in Lab 1.
Per-CPU State and Initialization¶
When writing a multiprocessor OS, it is important to distinguish between
per-CPU state that is private to each processor, and global state that
the whole system shares.
kern/cpu.h defines most of the per-CPU
struct CpuInfo, which stores per-CPU variables.
cpunum() always returns the ID of the CPU that calls it, which can
be used as an index into arrays like
cpus. Alternatively, the macro
thiscpu is shorthand for the current CPU’s
Here is the per-CPU state you should be aware of:
Per-CPU kernel stack. Because multiple CPUs can trap into the kernel simultaneously, we need a separate kernel stack for each processor to prevent them from interfering with each other’s execution. The array
percpu_kstacks[NCPU][KSTKSIZE]reserves space for NCPU’s worth of kernel stacks.
In Lab 2, you mapped the physical memory that
bootstackrefers to as the BSP’s kernel stack just below
KSTACKTOP. Similarly, in this lab, you will map each CPU’s kernel stack into this region with guard pages acting as a buffer between them. CPU 0’s stack will still grow down from
KSTACKTOP; CPU 1’s stack will start
KSTKGAPbytes below the bottom of CPU 0’s stack, and so on.
inc/memlayout.hshows the mapping layout.
Per-CPU TSS and TSS descriptor. A per-CPU task state segment (TSS) is also needed in order to specify where each CPU’s kernel stack lives. The TSS for CPU i is stored in
cpus[i].cpu_ts, and the corresponding TSS descriptor is defined in the GDT entry
gdt[(GD_TSS0 >> 3) + i]. The global
tsvariable defined in
kern/trap.cwill no longer be useful.
Per-CPU current environment pointer. Since each CPU can run different user process simultaneously, we redefined the symbol
curenvto refer to
thiscpu->cpu_env), which points to the environment currently executing on the current CPU (the CPU on which the code is running).
Per-CPU system registers. All registers, including system registers, are private to a CPU. Therefore, instructions that initialize these registers, such as
lidt(), etc., must be executed once on each CPU. Functions
trap_init_percpu()are defined for this purpose.
kern/pmap.c) to map per-CPU
stacks starting at
KSTACKTOP, as shown in
size of each stack is
KSTKSIZE bytes plus
KSTKGAP bytes of
unmapped guard pages. Your code should pass the new check in
The code in
initializes the TSS and TSS descriptor for the BSP. It worked in Lab 3,
but is incorrect when running on other CPUs. Change the code so that it
can work on all CPUs. (Note: your new code should not use the global
ts variable any more.)
When you finish the above exercises, run JOS in QEMU with 4 CPUs using
make qemu CPUS=4 (or
make qemu-nox CPUS=4), you should see output like
... Physical memory: 66556K available, base = 640K, extended = 65532K check_page_alloc() succeeded! check_page() succeeded! check_kern_pgdir() succeeded! check_page_installed_pgdir() succeeded! SMP: CPU 0 found 4 CPU(s) enabled interrupts: 1 2 SMP: CPU 1 starting SMP: CPU 2 starting SMP: CPU 3 starting
Our current code spins after initializing the AP in
Before letting the AP get any further, we need to first address race
conditions when multiple CPUs run kernel code simultaneously. The
simplest way to achieve this is to use a big kernel lock. The big
kernel lock is a single global lock that is held whenever an environment
enters kernel mode, and is released when the environment returns to user
mode. In this model, environments in user mode can run concurrently on
any available CPUs, but no more than one environment can run in kernel
mode; any other environments that try to enter kernel mode are forced to
kern/spinlock.h declares the big kernel lock, namely
kernel_lock. It also provides
unlock_kernel(), shortcuts to acquire and release the lock. You
should apply the big kernel lock at four locations:
i386_init(), acquire the lock before the BSP wakes up the other CPUs.
mp_main(), acquire the lock after initializing the AP, and then call
sched_yield()to start running environments on this AP.
trap(), acquire the lock when trapped from user mode. To determine whether a trap happened in user mode or in kernel mode, check the low bits of the
env_run(), release the lock right before switching to user mode. Do not do that too early or too late, otherwise you will experience races or deadlocks.
Apply the big kernel lock as described above, by calling
unlock_kernel() at the proper locations.
How to test if your locking is correct? You can’t at this moment! But you will be able to after you implement the scheduler in the next exercise.
Question 2. It seems that using the big kernel lock guarantees that only one CPU can run the kernel code at a time. Why do we still need separate kernel stacks for each CPU? Describe a scenario in which using a shared kernel stack will go wrong, even with the protection of the big kernel lock.
Challenge! The big kernel lock is simple and easy to use. Nevertheless, it eliminates all concurrency in kernel mode. Most modern operating systems use different locks to protect different parts of their shared state, an approach called fine-grained locking. Fine-grained locking can increase performance significantly, but is more difficult to implement and error-prone. If you are brave enough, drop the big kernel lock and embrace concurrency in JOS!
It is up to you to decide the locking granularity (the amount of data that a lock protects). As a hint, you may consider using spin locks to ensure exclusive access to these shared components in the JOS kernel: 1) The page allocator; 2) The console driver; 3) The scheduler; 4) The inter-process communication (IPC) state that you will implement in the part C.
Your next task in this lab is to change the JOS kernel so that it can alternate between multiple environments in “round-robin” fashion. Round-robin scheduling in JOS works as follows:
- The function
sched_yield()in the new
kern/sched.cis responsible for selecting a new environment to run. It searches sequentially through the
envsarray in circular fashion, starting just after the previously running environment (or at the beginning of the array if there was no previously running environment), picks the first environment it finds with a status of
inc/env.h), and calls
env_run()to jump into that environment.
sched_yield()must never run the same environment on two CPUs at the same time. It can tell that an environment is currently running on some CPU (possibly the current CPU) because that environment’s status will be
- We have implemented a new system call for you,
sys_yield(), which user environments can call to invoke the kernel’s
sched_yield()function and thereby voluntarily give up the CPU to a different environment.
Implement round-robin scheduling in
described above. Don’t forget to modify
syscall() to dispatch
Make sure to invoke
kern/init.c to create three (or more!) environments that all
run the program
Run make qemu. You should see the environments switch back and forth between each other five times before terminating, like below.
Test also with several CPUS: make qemu CPUS=2.
... Hello, I am environment 00001000. Hello, I am environment 00001001. Hello, I am environment 00001002. Back in environment 00001000, iteration 0. Back in environment 00001001, iteration 0. Back in environment 00001002, iteration 0. Back in environment 00001000, iteration 1. Back in environment 00001001, iteration 1. Back in environment 00001002, iteration 1. ...
yield programs exit, there will be no runnable environment
in the system, the scheduler should invoke the JOS kernel monitor. If
any of this does not happen, then fix your code before proceeding.
If you use CPUS=1 at this point, all environments should successfully run. Setting CPUS larger than 1 at this time may result in a general protection or kernel page fault once there are no more runnable environments due to unhandled timer interrupts (which we will fix below!).
3. In your implementation of
env_run() you should have called
lcr3(). Before and after the call to
lcr3(), your code makes
references (at least it should) to the variable
e, the argument
env_run. Upon loading the
%cr3 register, the addressing
context used by the MMU is instantly changed. But a virtual address
e) has meaning relative to a given address context–the
address context specifies the physical address to which the virtual
address maps. Why can the pointer
e be dereferenced both before
and after the addressing switch?
4. Whenever the kernel switches from one environment to another, it must ensure the old environment’s registers are saved so they can be restored properly later. Why? Where does this happen?
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
fork() and IPC in parts B and C of this lab.
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
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
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
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
In the parent,
fork() returns the child’s process ID, while in the
fork() returns 0. By default, each process gets its own
private address space, and neither process’s modifications to memory are
visible to the other.
You will provide a different, more primitive set of JOS system calls for
creating new user-mode environments. With these system calls you will be
able to implement a Unix-like
fork() entirely in user space, in
addition to other styles of environment creation. The new system calls
you will write for JOS are as follows:
- This system call creates a new environment with an almost blank
slate: nothing is mapped in the user portion of its address space,
and it is not runnable. The new environment will have the same
register state as the parent environment at the time of the
sys_exoforkcall. In the parent,
sys_exoforkwill return the
envid_tof the newly created environment (or a negative error code if the environment allocation failed). In the child, however, it will return 0. (Since the child starts out marked as not runnable,
sys_exoforkwill not actually return in the child until the parent has explicitly allowed this by marking the child runnable using….)
- Sets the status of a specified environment to
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 both refer to 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 convention that a value of 0 means “the current
environment.” This convention is implemented by
We have provided a very primitive implementation of a Unix-like
fork() 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
sys_yield as in the previous exercise. The parent exits
after 10 iterations, whereas the child exits after 20.
Implement the system calls described above in
kern/syscall.c. You will need to use various functions in
now, whenever you call
envid2env(), pass 1 in the
parameter. Be sure you check for any invalid system call arguments,
-E_INVAL in that case. Test your JOS kernel with
user/dumbfork and make sure it works before proceeding.
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 environment, runs
it for a while (e.g., a few iterations of
sys_yield()), then takes a
complete snapshot or checkpoint of the child environment, runs the
child for a while longer, and finally restores the child environment to
the state it was in at the checkpoint and continues it from there. Thus,
you are effectively “replaying” the execution of the child environment
from an intermediate state. Make the child environment perform some
interaction with the user using
that the user can view and mutate its internal state, and verify that
with your checkpoint/restart you can give the child environment a case
of selective amnesia, making it “forget” everything that happened beyond
a certain point.
This completes Part A of the lab; check it using make grade and
push your work as
lab4-a. If you are trying to figure out why a
particular test case is failing, run
./grade-lab4 -v, which will show
you the output of the kernel builds and QEMU runs for each test, until a
test fails. When a test fails, the script will stop, and then you can
jos.out to see what the kernel actually printed.
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).
xv6 Unix implements
fork() by copying all data from the parent’s
pages into new pages allocated for the child. This is essentially the
same approach that
dumbfork() takes. The copying of the parent’s
address space into the child is the most expensive part of the
However, a call to
fork() is frequently followed almost immediately
by a call to
exec() in the child process, which replaces the child’s
memory with a new program. This is what the the shell typically does,
for example. In this case, the time spent copying the parent’s address
space is largely wasted, because the child process will use very little
of its memory before calling
For this reason, later versions of Unix took advantage of 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 copy the address space mappings from the
parent to the child instead of the 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, writable 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
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
In the next piece of this lab, you will implement a “proper” Unix-like
fork() with copy-on-write, as a user space library routine.
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
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¶
A user-level copy-on-write
fork() needs to know about page faults on
write-protected pages, so that’s what you’ll implement first.
Copy-on-write is only one of many possible uses for user-level page
It’s common to set up an address space so that page faults indicate when some action needs to take place. 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 addresses that are not yet mapped. A typical Unix kernel must keep track of what action to take when a page fault occurs in each region of a process’s space. For example, a fault in the stack region will typically allocate and map new page of physical memory. A fault in the program’s BSS region will typically allocate a new page, fill it with zeroes, and map it. 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.
This is a lot of information for the kernel to keep track of. Instead of taking the traditional Unix approach, you will decide what to do about each page fault in user space, where bugs are less damaging. This design has the added benefit of allowing programs great flexibility in defining their memory regions; you’ll 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_env_set_pgfault_upcall system call. We have added a new member
env_pgfault_upcall, to record this
sys_env_set_pgfault_upcall system call. Be
sure to enable permission checking when looking up the environment ID of
the target environment, since this is 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-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” on behalf of
the user environment, in much the same way that the x86 processor
already implements stack switching 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-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
Each user environment that wants to support user-level page fault
handling will need to allocate memory for its own exception stack, using
sys_page_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
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 a
struct UTrapframe from
<-- UXSTACKTOP trap-time esp trap-time eflags trap-time eip trap-time eax start of struct PushRegs trap-time ecx trap-time edx trap-time ebx trap-time esp trap-time ebp trap-time esi trap-time edi end of struct PushRegs tf_err (error code) fault_va <-- %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. The
fault_va is the virtual address that caused the page fault.
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
tf->tf_esp rather than at
UXSTACKTOP. You should
first push an empty 32-bit word, then a
To test whether
tf->tf_esp is already on the user exception stack,
check whether it is in the range between
Implement the code in
kern/trap.c required to dispatch page faults to 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
_pgfault_upcall routine in
lib/pfentry.S. The interesting part is returning to the original
point in the user code that caused the page fault. You’ll return
directly there, without going back through the kernel. The hard part is
simultaneously switching stacks and re-loading the EIP.
Finally, you need to implement the C user library side of the user-level page fault handling mechanism.
user/faultread (make run-faultread). You should see:
...  new env 00001000  user fault va 00000000 ip 0080003a TRAP frame ...  free env 00001000
user/faultdie. You should see:
...  new env 00001000 i faulted at va deadbeef, err 6  exiting gracefully  free env 00001000
user/faultalloc. You should see:
...  new env 00001000 fault deadbeef this string was faulted in at deadbeef fault cafebffe fault cafec000 this string was faulted in at cafebffe  exiting gracefully  free env 00001000
If you see only the first “this string” line, it means you are not handling recursive page faults properly.
user/faultallocbad. You should see:
...  new env 00001000  user_mem_check assertion failure for va deadbeef  free env 00001000
Make sure you understand why
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¶
You now have the kernel facilities to implement copy-on-write
entirely in user space.
We have provided a skeleton for your
fork() should create a new environment, then scan
through the parent environment’s entire address space and set up
corresponding page mappings in the child. The key difference is that,
dumbfork() copied pages,
fork() will initially only copy
fork() will copy each page only when one of the
environments tries to write it.
The basic control flow for
fork() is as follows:
The parent installs
pgfault()as the C-level page fault handler, using the
set_pgfault_handler()function you implemented above.
The parent calls
sys_exofork()to create a child environment.
For each writable or copy-on-write page in its address space below UTOP, the parent calls
duppage, which should map the page copy-on-write into the address space of the child and then remap the page copy-on-write in its own address space.
duppagesets both PTEs so that the page is not writeable, and to contain
PTE_COWin the “avail” field to distinguish copy-on-write pages from genuine read-only pages.
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?
fork()also needs to handle pages that are present, but not writable or copy-on-write.
The parent sets the user page fault entrypoint for the child to look like its own.
The child is now ready to run, so the parent marks it runnable.
Each time one of the environments writes a copy-on-write page that it hasn’t yet written, it will take a page fault. Here’s the control flow for the user page fault handler:
- The kernel propagates the page fault to
_pgfault_upcall, which calls
pgfault()checks that the fault is a write (check for
FEC_WRin the error code) and that the PTE for the page is marked
PTE_COW. If not, panic.
pgfault()allocates a new page mapped at a temporary location and copies the contents of the faulting page 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.
lib/fork.c code must consult the environment’s page
tables for several of the operations above (e.g., that the PTE for a
page is marked
PTE_COW). The kernel maps the environment’s page
UVPT exactly for this purpose. It uses a clever mapping
to make it to make it easy to lookup PTEs for user
lib/entry.S sets up
uvpd so that you can
easily lookup page-table information in
Exercise 12. Implement
Test your code with the
forktree program. It should produce the
following messages, with interspersed ‘new env’, ‘free env’, and
‘exiting gracefully’ messages. The messages may not appear in this
order, and the environment IDs may be different.
1000: I am '' 1001: I am '0' 2000: I am '00' 2001: I am '000' 1002: I am '1' 3000: I am '11' 3001: I am '10' 4000: I am '100' 1003: I am '01' 5000: I am '010' 4001: I am '011' 2002: I am '110' 1004: I am '001' 1005: I am '111' 1006: I am '101'
Implement a shared-memory
version should have the parent and child share all their memory pages
(so 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
user/forktree.c to use
sfork() instead of regular
fork(). Also, once you have finished implementing IPC in part C, use
sfork() to run
user/pingpongs. You will have to find a new
way to provide the functionality of the global
Your implementation of
fork makes a huge number of system
calls. On the x86, switching into the kernel using interrupts 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
to use this interface.
How much faster is your new
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
instruction? How many times do you execute
int 0x30 in your
fork? Is accessing the
TSS stack switch also expensive? And so
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. QEMU
doesn’t emulate this instruction faithfully (it can either count the
number of virtual instructions executed or use the host TSC, neither of
which reflects the number of cycles a real CPU would require).
This ends part B. As usual, you can grade your submission with make
grade and push your commit tagged with
Part C: Preemptive Multitasking and Inter-Process communication (IPC)¶
In the final part of lab 4 you will modify the kernel to preempt uncooperative environments and to allow environments to pass messages to each other explicitly.
Clock Interrupts and Preemption¶
user/spin test program. This test program forks off a child
environment, which simply spins forever in a tight loop once it receives
control of the CPU. Neither the parent environment 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 retaking control of the CPU from it, we must
extend the JOS kernel to support external hardware interrupts from the
External interrupts (i.e., device interrupts) are referred to as IRQs.
There are 16 possible IRQs, numbered 0 through 15. The mapping from IRQ
number to IDT entry is not fixed.
picirq.c maps IRQs
0-15 to IDT entries
IRQ_OFFSET is defined to be decimal 32. Thus the
IDT entries 32-47 correspond to the IRQs 0-15. For example, the clock
interrupt is IRQ 0. Thus,
IDT[IRQ_OFFSET+0] (i.e., IDT) 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
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 xv6 Unix. External
device interrupts are always disabled when in the kernel (and, like
xv6, enabled when in user space). External interrupts are controlled by
FL_IF flag bit of the
%eflags register (see
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
register as we enter and leave user mode.
You will have to ensure that the
FL_IF flag is set in user
environments 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. We masked interrupts with the very first instruction of the
bootloader, and so far we have never gotten around to re-enabling them.
initialize the appropriate entries in the IDT and provide handlers for
IRQs 0 through 15. Then modify the code in
kern/env.c to ensure that user environments are always run with
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.,
spin), you should
see the kernel print trap frames for hardware interrupts. While
interrupts are now enabled in the processor, JOS isn’t yet handling
them, so you should see it misattribute each interrupt to the currently
running user environment and destroy it. Eventually it should run out of
environments to destroy and drop into the monitor.
Handling Clock Interrupts¶
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
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.
Modify the kernel’s
trap_dispatch() function so that it
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
environment 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 environment and terminate gracefully.
This is a great time to do some regression testing. Make sure that you
haven’t broken any earlier part of that lab that used to work (e.g.
forktree) by enabling interrupts. Also, try running with multiple
make CPUS=2 target. You should also be able to pass
stresssched now. Run make grade to see for sure. You should now get
a total score of 65/75 points on this lab.
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 best. We won’t get into that debate. Instead, we’ll implement a simple IPC mechanism and then try it out.
IPC in JOS¶
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_try_send. Then you will implement two library wrappers
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 de-schedules 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
To try to send a value, an environment calls
both the receiver’s environment id and the value to be sent. If the
named environment is actually receiving (it has called
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
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
Similarly, a library function
ipc_send will take care of repeatedly
sys_ipc_try_send until the send succeeds.
When an environment calls
sys_ipc_recv with a valid
UTOP), 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_try_send with a valid
UTOP), it means the sender wants to send the page currently
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
structure to the permissions of the page received, or zero if no page
kern/syscall.c. Read the comments on both before implementing them,
since they have to work together. 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.
Then implement the
ipc_send functions in
user/primes functions to test your IPC
user/primes will generate for each prime number a new
environment until JOS runs out of environments. You might find it
interesting to read
user/primes.c to see all the forking and IPC
going on behind the scenes.
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! One of the most impressive examples 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(x+x^3).
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.
This ends part C. Make sure you pass all of the make grade tests and
don’t forget to write up your answers to the questions in
Before handing in, use git status and git diff to examine your changes and don’t forget to git add answers-lab4.txt. When you’re ready, commit your changes with git commit -am ‘my solutions to lab 4’, then make handin and follow the directions.