We describe Cartwright and Felleisen's modular and compositional approach to effects, discuss extensions and present its implementations in Haskell. The principal ideas, in the words of Cartwright and Felleisen, are:
When designing a program we should start thinking what effect we want
to achieve rather than which monad transformer to use. Instead of
jumping straight to StateT
and so on, we ought to identify what
transformation on the world and its resources we wish to
effect. Written formally, this transformation often takes the form of an
effect handler. Our framework is designed around such handlers,
encouraging custom effects for program fragments and their composition.
Our alternative to a monad transformer stack is the single monad, for the coroutine-like communication of a client with its handler. Its type reflects possible requests, i.e., possible effects of a computation. To support arbitrary effects and their combinations, requests are values of an extensible union type, which allows adding and, notably, subtracting summands. Extending and, upon handling, shrinking of the union of possible requests is reflected in its type, yielding a type-and-effect system for Haskell. The library is lightweight, generalizing the extensible exception handling to other effects and accurately tracking them in types.
[Abstract of the follow-up paper]
We present a rational reconstruction of extensible effects, the
recently proposed alternative to monad transformers, as the confluence
of efforts to make effectful computations compose. Free monads and
then extensible effects emerge from the straightforward term
representation of an effectful computation, as more and more
boilerplate is abstracted away. The generalization process further
leads to freer monads, constructed without the Functor constraint. The
continuation exposed in freer monads can then be represented as an
efficient type-aligned data structure. The end result is the
algorithmically efficient extensible effects library, which is not
only more comprehensible but also faster than earlier implementations.
As an illustration of the new library, we show three surprisingly simple applications: non-determinism with committed choice (LogicT), catching IO exceptions in the presence of other effects, and the semi-automatic management of file handles and other resources through monadic regions.
We extensively use and promote the new sort of `laziness', which underlies the left Kan extension: instead of performing an operation, keep its operands and pretend it is done.
The ambition is for Eff
to be the only monad in Haskell. Rather
than defining new monads programmers will be defining new effects,
that is, effect interpreters.
Joint work with Amr Sabry, Cameron Swords, and Hiromi Ishii.
exteff.pdf [286K]
doi:10.1145/2503778.2503791
Extensible Effects: An Alternative to Monad Transformers
The paper published in the Proceedings of the 2013 ACM SIGPLAN symposium on Haskell, pp. 59-70. Boston, MA, USA. September 23-24, 2013.
talk.pdf [231K]
The annotated slides of the talk presented at the Haskell Symposium 2013 on September 23, 2013 in Boston.
Eff1.hs [39K]
OpenUnion51.hs [5K]
FTCQueue1.hs [2K]
The complete, self-contained implementation of extensible effects
The file Eff.hs
(importing an implementation of open unions)
defines and implements the API for extensible effects: an EFF
monad as an extensible effect. It also
implements the standard monadic effects such as exception, state, and
non-determinism. The file contains many examples and test code.
The December 2015 version introduces several specialization RULEs
and inlining directives, which notably improve the performance.
Defining a new class for each effect, such as MonadState
,
is possible but not needed at all. With monad transformers, a
class per effect is meant to hide the
ordering of transformer layers in a monad transformer stack. Effect
libraries abstract over the implementation details out of the
box. Crutches -- extra classes -- become unnecessary.
We also illustrate, for pedagogical purposes, an implementation of
the State effect as a combination of the existing but differently
interpreted Reader and Writer effects. Truly the meaning of an effect
is its interpretation.
EffDynCatch.hs [5K]
Catching of dynamic and asynchronous exceptions in the presence of other effects -- the solution of the MonadCatchIO problem. This code accompanies Section 6 of the paper.
EffRegion.hs [11K]
EffRegionTest.hs [10K]
Monadic regions through extensible effects: the code accompanying Section 7 of the paper. The test code is essentially the same as that of the original Monad Regions implementation.
Benchmarks.hs [21K]
Bench_nondet.hs [13K]
A few micro-benchmarks
Eff.hs [30K]
The now obsolete code for the original (2013) implementation.
Reader00.hs [3K]
Reader0.hs [4K]
The warm-up example: the single Reader
effect as an interaction with an authority. This is the code for Sec 3.1 of the 2013 paper. Reader00.hs
describes the single Reader Int
effect; the other file generalizes to arbitrarily typed environment monad.
ExtMTL.hs [9K]
A variation of Eff.hs
emulating monad transformer classes of MTL.
The framework of extensible effects indeed can define the instances
for MonadError
, MonadReader
, MonadState
, etc. These instances
require fewer annotations in the user code; on the other hand, they
are less general, enforcing a single effect layer of a particular kind.
Monad transformers have become an integral part of Haskell, with many tutorials. Rarely do the drawbacks of transformers get a mention. As a rare exception, Chapter 18 of `Real World Haskell' points out the overhead added by each monad transformer layer, an occasional need for the ungainly explicit lifting, and the difficulty of building monad transformer libraries: when adding a new transformer, one has to explicitly specify `lifting', its interaction with all other transformers. Alas, the biggest drawback of monad transformers is hardly mentioned at all.
Monad transformers have an inherent limitation: they enforce the static ordering of effect layers and hence statically fixed effect interactions. There are practically significant computations that require interleaving of effects. `Delimited Dynamic Binding' (ICFP 2006) was first to bring up this point. The `Extensible Effects' paper expanded that discussion on new examples. Section 5 describes simple and common programming patterns that are particularly problematic with monad transformers because the static ordering of effect layers is not flexible.
Bryan O'Sullivan, Don Stewart, and John Goerzen: Real World Haskell
Chapter 18. Monad transformers
< http://book.realworldhaskell.org/read/monad-transformers.html#id6594 >
An early compelling case for open unions are extensible exceptions,
which have been part of Haskell for
many years (Simon Marlow, Haskell Workshop 2006). To permit throwing
exceptions of arbitrarily many types, the thrown exception value is an
open union (see SomeException
in Control.Exception
). Raising an
exception puts -- injects -- a particular exception value into the open
union. When we handle an exception, we project: check if the
thrown exception is of a particular desired type. (Extensible
effects operate in the same manner; supporting, in addition,
the resumption of an `exception'.) Thus, at its core,
an open union should let us inject a value of any type into the union and to project a type
from the union, that is, to find find out if the union value
was previously injected with a particular specific type. These operations
are familiar from OOP as downcast and upcast.
The open-union type of exceptions, SomeException
(or the similar exn
in ML), gives no indication of possible summands -- that is,
which particular exception types may be in the union. Therefore,
neither Haskell nor ML can ensure that a program handles all
exceptions that could be raised during its execution.
To do better, the type of an open union should be annotated with the set of possible summands. The injection function will add the type
of the injected value to the union type, unless it was there already.
As always with types, the type of the open union is an approximation
for the type of the value therein. Consider the
simplest union Either Int Bool
: at run time, the union value is
either a single Int
or a single Bool
. The union type is an
approximation: we cannot generally determine at compile-time the
specific type of the value stored in the union. We
are sure however that this type is not a String
and hence attempting
to project a String
value from an Either Int Bool
union is a
compile-time error. Such type-annotated union is called a type-indexed
co-product.
The familiar data type (Either a b)
is the simplest example of typed
unions, but it is not extensible. The
constructors Left
and Right
are injections, and the projections
are realized via pattern-matching:
prj1:: Either a b -> Maybe a prj1 (Left x) = Just x prj1 _ = NothingThe type checker does not let us inject a value of a type other than
a
and b
into Either a b
, hence restricting injection
to values that participate in the union. We can only
project at types a
and b
-- Either a b
is a union of exactly
two types, and thus not extensible. Furthermore, it is not
abstract: we must know the exact structure of the union in order to
choose the proper injection, Left
or Right
. The type Either Int Bool
is different from Either Bool Int
, although they are morally the same
union.
Heeding the drawbacks of Either
, we arrive at the following
interface for open unions:
data Union (r :: [*]) -- abstract type family Member (t :: *) (r :: [*]) :: Constraint inj :: Member t r => t -> Union r prj :: Member t r => Union r -> Maybe t decomp :: Union (t ': r) -> Either (Union r) t
The union type Union r
is tagged with r
, which is meant to be a
set of summands. For the lack of type-level sets,
it is realized as a type-level list, of the kind [*]
. The injection inj
and the projection prj
ensure that the type t
to inject
or project must be a member of the set r
, as determined by the type-level
function Member t r
. The function decomp
performs the orthogonal
decomposition. It checks to see if the union value given as the
argument is a value of type t
. If so, the value is returned as Right t
. Otherwise, we learn that the received union value does
not in fact contain the value of type t
. We return the union,
adjusting its type so that it no longer has t
. The function decomp
thus decomposes the open union into two orthogonal ``spaces:''
one with the type t
and the other without. The decomposition operation,
which shrinks the type of open unions, is the crucial distinction
of our interface from the previous designs (Liang et al. 1995,
Swierstra 2008, polymorphic variants of OCaml). It is this decomposition
operation, used to `subtract' handled exceptions/effects, that
insures that all effects are handled. The constraint Member t r
may be seen as the interface between inj
and prj
on one hand
and decomp
on the other hand: for each injection or projection
at type t
there shall be a decomp
operation for the type t
.
This basic interface of open unions has several variations and
implementations. One is OpenUnion1
, presented in the Haskell 2013
extensible-effects paper. It is essentially the one described in
Appendix C of the full HList paper, published in 2004. In the version
for extensible-effects, the summands of the open union have the kind * -> *
rather than *
. This implementation uses Dynamic
as the
open union. Therefore, all operations take constant time -- like
polymorphic variants of OCaml and unlike open unions used by Liang et
al. 1995 and Swierstra 2008.
One may notice a bit of asymmetry in the above interface. The functions inj
and prj
treat the open union index r
truly as a set of types.
The operations assert that the type t
to inject or project is a member
of the set, without prescribing where exactly t
is to occur in the concrete
representation of r
. On the other hand, decomp
specifies that the type t
must be at the head of the list that represents the set of summand
types. It is unsatisfactory, although has not presented a problem
so far for extensible effects. If the problem does arise, it may
be cured with an easily-defined conversion function of the type conv :: SameSet r r' => Union r -> Union r'
, akin to an annotation.
The other solutions to the problem (based on constraint kinds, for example)
are much more heavier-weight, requiring many more annotations. Perhaps
implicit parameters may help:
e1 = if ?x then ?y else (0::Int) -- inferred: e1 :: (?x::Bool, ?y::Int) => Int f :: ((?x::Bool) => r) -> r -- explicit signature required f x = let ?x = True in x t1 = f e1 -- inferred: t1 :: (?y::Int) => IntThe inferred type of
t1
no longer contains the ?x::Bool
constraint, which thus has been subtracted. The type of the `subtraction
function' f
, the handler, only mentions the removed constraint,
saying nothing of other constraints or if there are other constraints.
Binding an implicit parameter builds the dictionary and
makes the constraint go away. One could wish the same for
proper constraints.
One may notice that the open union interface, specifically,
the function decomp
, does not check for
duplicates in the set of summands r
. This check is trivially to add --
in fact, the HList implementation of type-indexed co-products did
have such a check and so implemented true rather than disjoint unions.
In case of extensible effects, the duplicates are harmless, letting us
nest effect handlers of the same type. The dynamically closest
handler wins -- which seems appropriate: think of reset in delimited
control. There is even a test case for nested handlers in Eff.hs
.
The implementation OpenUnion1
was received with significant
controversy, often derailing discussions of extensible effects,
which work with any open union implementation. Although OpenUnion1
provides constant-time operations, it relies on Dynamic
, which
requires the Typeable
constraint all throughout
extensible-effects. Although, with few exceptions, it is a minor
annoyance, it is the annoyance still. The OverlappingInstances
extension used by OpenUnion1
was also objected to, although without
much reason since the extension does not leak from the open union
implementation to the rest of the code. To address the concerns, OpenUnion41
was developed. It has the same interface as OpenUnion1
but uses neither Typeable
nor OverlappingInstances
. It relies on
GADTs and closed type families. Alas, its operations in the
worst case take time proportional to the
number of summands in the union. The recent OpenUnion5
implementation brought back the constant-time union operations. It
realizes ``strong sums'': existentials over a finite type universe
that do not hide the type of the containing value.
We have thus seen the design space for typed open unions and a few sample implementations. Hopefully more experience will help choose an optimal implementation and introduce it into Haskell.
Typeable
. It relies on the fact that the type parameter r
, the list of possible summand types, is actually a type
universe. An integer index in that list serves well as a value-level
representation for the indexed type. Whereas OpenUnion5
relies on
closed type families, OpenUnion51
uses overlapping instances. The
latter can resolve even for unground types, but closed type families
are subject to a strict apartness condition.OpenUnion41.hs [4K]
Another implementation of open unions, with neither
Typeable
nor overlapping instances. We pay for the lack of
Typeable
with non-constant--time injections and projections.
However, the more frequent decomposition takes short constant time.
This implementation has exactly the same interface as the others,
and can be used in any extensible-effects code
as the drop-in replacement for more
controversial implementations like OpenUnion1
.
OpenUnion1.hs [3K]
The code for open unions used in the paper. It is close to the implementation in the HList paper. It relies on Typeable, and emulates closed type families using the overlapping instances extension that is limited to that module.
OpenUnion2.hs [3K]
The version of open union without any overlapping instances, directly using closed type families.
OpenUnion3.hs [3K]
Another, somewhat dual implementation, relying on universals rather than existentials
TList.hs [2K]
The old implementation of open unions, without overlapping instances or
Typeable
The type of a value determines the set of operations on it. The set may be further restricted by the computation state: e.g., although an open file handle and a closed file handle have the same type, the sets of meaningful operations on them greatly differ. Type-state is the refinement of the notion of type to characterize such restriction. Session types, in the news lately, are one manifestation of the type-state. Whereas monads represent stateful computations, type-state calls for a parameterized monad, as a poor programmer's substructural type system. Tracking type-state seems to require quite advanced type (and effect) systems, in particular, modal or substructural type systems. None of these advanced type systems are available in mainstream Haskell.
Monad transformers -- and recently, algebraic, extensible and other `effects' -- let us combine in the same program independently developed monadic computations, each with their own side-effects. The effects may need to interact in a controlled fashion. The similar combining of parameterized monads has not been addressed so far. The problem is not straightforward, as we shall see from several failed attempts. We eventually present the solution, along the lines of extensible effects. The inferred types curiously look like session types.
For clarity, this abstract uses a rather specific example: two mutable cells. The generalizations, to other effects and to truly open unions, and optimizations are all doable and have been already described. We elide them to demonstrate the problem and the solution in the purest, barest form.
It may seem that the approaches to combine monadic computations should
straightforwardly extend to type-state (parameterized monads). That turns
out not to be the case. In the hindsight, the stumbling block is easy
to see. Consider two differently effectful expressions m1
and m2
, which
we wish to combine in the same program. It is most revealing to look
at the following combination:
m12 = if test then m1 else m2where
test :: Bool
is some decision procedure. Clearly m12
performs what is called in process calculi an internal choice. An effect
may be regarded as an interaction (communication) with the context.
The expression m12
may hence interact as m1
or as m2
,
based on an internal decision.
The type of an expression should describe what interactions it may
(potentially) make. The type of m12
should therefore represent the
union of effects of m1
and m2
. In the framework of extensible effects, m1
has the type Eff r1 a
and m2
has the type Eff r2 a
; then m12
has essentially the type Eff (union r1 r2) a
. No wonder
the implementation of open unions is the crucial part of any
extensible-effect library.
Eff r a
is an ordinary monad and r
may be considered a degenerate
session type: A computation m :: Eff [Read Int, Write String] Bool
,
for example, tells that m
, if terminates, produces a boolean and can
make Read Int
or Write String
requests -- an indeterminate number
of times and in any order. Type-state reins in: a type-state computation
may require that, e.g., any Write String
request be preceded by a Read Int
request. Ordinary monads cannot express such computations,
but parameterized monads can: m1 :: EffP [Read Int] [Write String] Bool
says that computation m1
may make a Read Int
request. If it really
does that, it may then, afterwards, make a Write String
request (or yield a value). Suppose m2
has the type EffP [Writer String] [Read Int] Bool
-- it may make a Write String
request and after that, Read Int
.
If we then combine m1
and m2
into m12
as in our running example,
what is the type of m12
? It should be something like EffP [Read Int, Write String] ?? Bool
-- telling that the computation
may do either Read Int
or Write String
. But what is the type state
afterwards? It should be Write String
if m12
actually made
the first choice, or Read Int
otherwise. It seems we need dependent sums!
This is the unexpected predicament in what at first seemed to be
a straightforward generalization of extensible effect to parameterized
monads.
ParamEff1.hs [12K]
The accompanying complete source code
In genetics, crossover is the interchange of segments between two
chromosomes of the same `type' (or, homologous).
Genetic algorithms borrow this term to mean the interchange of segments
between two structures representing sets of optimization parameters.
These data structures are called chromosomes as well, and the
parameters, or features, are also called genes.
Here is an example with two chromosomes that are lists of integers
(which we write in the explicit (::)
and []
notation).
1 :: 2 :: 3 :: [] 10 :: 20 :: []First we cut each list in two pieces
1 :: 2 :: _ and 3 :: [] 10 :: _ and 20 :: []where
_
represents the hole, which is left at the place of a cut-off
branch. Then we swap the branches, grafting the cut-off branch from
one list into the hole of the other.1 :: 2 :: 20 :: [] 10 :: 3 :: []The second example
1 :: _ :: 3 :: [] and 2 _ :: 20 :: [] and 10cuts the leaves:
2
from the first list and the leaf 10
from the second one.
The interchange gives1 :: 10 :: 3 :: [] 2 :: 20 :: []Not all cuts produce swappable branches: e.g., the leaf
1
of the first list
is not interchangeable with the branch 3::[]
of the second because they have
incompatible types. What we have just shown is a so-called cut-and splice
crossover. Other variations (with the same cut point for both
chromosomes, with two cut points, etc.) can also be easily written
with our library.Our code implements the above surgery, of cutting and grafting. First, we write the generic traversal with branch replacement, using the Scrap-your-boilerplate generic programming library (SYB). Next, we differentiate the traversal obtaining the procedure to cut a branch at an arbitrary point. Finally, we add grafting of the swapped branches cut off two data structures, obtaining the crossover. Although the branches must match in type, the types of the whole structures may differ. We need effects: the coroutine effect for differentiating the traversal, and non-determinism for choosing a cut location. We also need state to keep track of the number of cuts made to a structure.
The most interesting part is the traversal and its differentiation. The traversal function of the signature below receives a data structure, and a function that can examine a branch and possibly replace it.
newtype Updates = Updates Int -- update count traverse :: Member (State Updates) r => (forall a. (Data a) => a -> Eff r a) -> (forall a. (Data a) => a -> Eff r a)
In the latter case, that function should increment the update count,
shared among the traversals of all branches. The traversal is terminated when the
count exceeds a threshold. The traverse
itself is a wrapper over
SYB's gfoldl
, counting updates to branches.
traverse f = check_done $ \x -> f x >>= check_done traverse_children where threshold = 1 check_done go x = get >>= \case Updates n | n >= threshold -> return x _ -> go x traverse_children = gfoldl traverse_child return traverse_child builda x = liftM2 ($) builda (traverse f x)
To differentiate the traversal -- to suspend it at each
encountered branch -- we need a coroutine effect. The branches
may have different types: for example, the traversal of an integer
list comes across branches of the type Int
and [Int]
. Therefore,
a custom coroutine effect is needed, abstracting the type of a branch
and ensuring it supports the Data
interface.
data YieldD v = forall a. Data a => YieldD a (Maybe a -> v) yieldD :: (Data a, Member (State Updates) r, Member YieldD r) => a -> Eff r a yieldD x = send (inj . YieldD x) >>= \case Nothing -> return x Just x -> modify (\ (Updates n) -> Updates (n+1)) >> return xThe request
YieldD
sends to its effect handler a value and waits to
be resumed with a possibly new value, which becomes the result of yieldD x
after incrementing the update count. If the suspension was
resumed with no new value, yieldD x
just returns x
. The request YieldD
is easily turned into the description of a structure with a
cut branch. A structure of the type a
with a hole of the type b
is
represented as a
function Maybe b -> (Eff r (Cut r a))
, that waits
for a value to fill the hole with and then produces a new Cut
:data Cut r a = CDone a | Cut (YieldD (Eff r (Cut r a)))
To differentiate traverse
we merely apply it to yieldD
, obtaining
traverse_diff :: Data a => a -> Eff r (Cut r (a,Updates))
The next step, cutting a data structure at a random point, is also easy, by non-deterministically choosing one cut point among the many encountered during the whole traversal.
random_walk :: (Member Choose r, Data a) => a -> Eff r (Cut r (a,Updates)) random_walk a = traverse_diff a >>= check where check y@CDone{} = return y check y@(Cut (YieldD x k)) = return y `mplus` (k Nothing >>= check)
Finally, crossover
swaps the branches between two data structures
cut at non-deterministically chosen points, provided the branches have
the suitable type.
crossover :: (Member Choose r, Data a, Data b) => a -> b -> Eff r (a,b) crossover x y = do tx <- random_walk x ty <- random_walk y case (tx,ty) of (Cut (YieldD x kx), Cut (YieldD y ky)) | Just x' <- cast x, Just y' <- cast y -> do (xnew,_) <- zip_up =<< kx (Just y') (ynew,_) <- zip_up =<< ky (Just x') return (xnew,ynew) _ -> mzero
As an example, crossing over [1,2,3]
with Just 10
produces three possible
results: ([10,2,3],Just 1)
, ([1,10,3],Just 2)
and ([1,2,10],Just 3)
.
There are 18 possible ways to crossover [1,2,3]
and [10,20]
, including
the two cases described earlier. The code has more examples,
including crossover between two trees.
We have presented the library for generic crossover between two arbitrary algebraic data structures. The library is a showcase for extensible effects, demonstrating how easy it is to define a custom effect for a given problem and use it alongside standard effects. We are spared annoying lifting, so typical of monad transformers. Raising and handling of effects is as simple as using side-effectful operations in ML or other impure languages. Unlike ML however, the effects of a function are seen in its type, and the type checker watches that all effects are handled in the end.
Lazy state is best understood on the delightfully puzzling example designed by Albert Y.C. Lai and popularized by Edward Kmett. In MTL, it is written thusly:
lzs = do lzs modify ((1::Int):)One may wonder what could possibly come out of this left recursion. In a strict language, surely nothing:
lzs
loops. To see what we can
get in Haskell, let's look at the type: lzs :: Monad m => StateT
[Int] m ()
. It is a computation over the integer-list state. Let s
be the final state of lzs
. The computation starts by calling itself;
upon return the state will be s
, by assumption; the state is then
modified by prepending 1
. We hence obtain the equation s = 1:s
whose solution, in Haskell, is the infinite sequence of ones. To
check*Main> take 5 . snd $ runState lzs undefined [1,1,1,1,1]
lzs
indeed ended in the state that is [1,1,..]
.
MTL, as extensible effects, has been motivated by the desire to
compose and extend effectful computations. From its inferred type lzs
:: Monad m => StateT [Int] m ()
, lzs
is indeed extensible: it is a
state computation on top of an arbitrary effectful m
. Let's take it
to be the Error monad. Although lzs
itself throws no exceptions, it
may be used as a part of a larger program that does.
lzseRun :: Either String ((), [Int]) lzseRun = runIdentity $ runErrorT $ runStateT lzs undefinedAlas,
lzseRun
goes into the infinite loop. Again, lzs
, taken as it
was, does not make any use of the underlying Error monad. Yet that
monad drastically changed the lzs
behavior. (We shall see that
extensible effects can do better.)We learned two lessons from this experiment. First, trying to be too clever with laziness invariably comes to bite back. As much as looking at the program a wrong way may send it into the infinite loop. Second, lazy state is not easily extensible, by nature. It is not a good counter-example to extensible effects, which emphasize easy extensibility.
Still, the challenge remains: Can we write something like lzs
,
an on-demand stateful computation, with extensible effects? At first
blush, the answer is no. First, the state effect provided by the Eff
library is intentionally strict, to stem space leaks. There is the
other reason, left as an exercise to the reader. (Solving it
will make the explanation below more understandable.)
If we contemplate the lazy state more, we realize how crucial it is
to delay the recursive invocation of lzs
until its state
is demanded. The modify
operation in lzs
obtains the current state
but does not demand its value. That is why lzs
does not diverge.
To write lzs
with extensible effects, we have to make the implicit delay
explicit. We have seen the subtlety and the fragility of the lazy state;
being explicit seems to be a good idea. We hence arrive at the following
declaration of the lazy state effect:
data LazyState s v where LGet :: LazyState s s LPut :: s -> LazyState s () Delay :: Eff '[LazyState s] a -> LazyState s aThe
LGet
and LPut
operations need no further explanations; we could have
reused the State
effect already in the Eff
library, if not for the fact
that it is strict. Delay
is the effect to delay a given stateful computation
until its result or state are demanded. The familiar operations on state
are defined in terms of sending the effect requests:lget = send LGet lput = send . LPut lmodify f = do {s <- lget; lput (f s)} onDemand :: Member (LazyState s) r => Eff '[LazyState s] v -> Eff r v onDemand = send . DelayThe argument of
onDemand
is an ``effect-closed'' computation, with
the sole LazyState
. This closedness is inherent in lazy state, as we
have seen -- and is easy to understand once we get to the LazyState
handler. However, onDemand e
itself may be used with
arbitrary other effects as we show later.
The handler for LazyState
is almost conventional:
the LGet
request is replied with the current state s
and LPut s
takes s
to be the new state. (Recall, in the Eff library, handle_relay_s
is the `deep' handler.)
runStateLazy :: s -> Eff (LazyState s ': r) a -> Eff r (a,s) runStateLazy s = handle_relay_s s (\s x -> return (x,s)) (\s req k -> case req of LGet -> k s s -- reply with s and continue in the state s LPut s -> k s () Delay m -> let ~(x,s1) = run $ runStateLazy s m in k s1 x)The interesting bit is the handling of
Delay
-- by running the
argument computation m
and continuing with its outcome. In a
non-strict language, however, runStateLazy s m
is not actually
performed unless its state s1
or the result x
are demanded. This
is exactly what we want from the lazy state. It also becomes clear
why we need to statically know all the effects of the
delayed computation: we should be able to run them all.That was the whole implementation. Edward Kmett's example looks as follows and produces the expected result:
lex3 = do onDemand lex3 lmodify ((1::Int):) ex3Run = let (x,s) = run $ runStateLazy (undefined::[Int]) lex3 in (x,take 5 s) -- ((),[1,1,1,1,1])
There is more: the lazy state computation can be combined in the same program with computations that actually do other effects, like throwing exceptions:
lex31 :: Member (LazyState [Int]) r => Eff r () lex31 = do onDemand (lex31 :: Eff '[LazyState [Int]] ()) lmodify ((1::Int):) lex5 = do lex31 x <- lget throwError ((take 5 x)::[Int]) ex5Run :: Either [Int] a ex5Run = fst . run . runStateLazy (undefined::[Int]) . runError $ lex5 -- Left [1,1,1,1,1] ex51Run :: Either [Int] (a,[Int]) ex51Run = run . runError . runStateLazy (undefined::[Int]) $ lex5 -- Left [1,1,1,1,1]Unlike the
lzseRun
earlier, both ex5Run
and ex51Run
terminate, with
the result shown in the comments. Although lex31
does not throw exceptions,
it can be used with any other effect beside LazyState
, as its signature
shows. It can indeed, regardless of the order LazyState
and the other effects
are handled. There is a slight sleigh of hand, in writing the signature
for lex31
and providing the type annotation on the recursive call. The reader
is encouraged to guess its significance.
The backward state, in comparison, is trivial: the run-on-the-mill
laziness with no delay tricks. To remind, in the ordinary (forward)
state computation, get
returns what has been most recently put
.
With the backward state, get
returns what will be put
next. Yes,
it is backwards. An example should make it clear:
st = do x <- lget lput (1::Int) lput (1::Int) y <- lget lput (2::Int) lput (3::Int) z <- lget lput (4::Int) return (x,y,z)To run it as usual, `forwards', we supply the initial state, say,
0
,
and obtain the result and the final state, from the last put
:stF :: ((Int,Int,Int),Int) stF = run $ runStateLazy (0::Int) st -- ((0,1,3),4)We used the previously described
runStateLazy
. The very same st
can be
run backwards, giving the result and the state from the first put
:stB :: ((Int,Int,Int),Int) stB = runStateBack st -- ((1,2,4),1)Each
get
gets the value from the following put
.We can play infinite games by modifying the state we have not yet obtained:
ones :: [Int] ones = snd $ runStateBack $ do s <- lget lput ((1::Int):s) take 5 ones -- [1,1,1,1,1]As the final exercise, the reader is encouraged to similarly represent the storyline of ``Primer''.
The implementation of runStateBack
is unremarkable. Like all
laziness tricks, it is best understood through Attribute Grammars. The
overall state is represented with two attributes: the inherited getAttr
and the synthesized putAttr
. The get
operation just
takes what is in getAttr
. All the action is happens at the put s
node of the expression tree:
this.getAttr = child.putAttr this.putAttr = sAt the root node,
putAttr
becomes getAttr
, tying the knot. The
accompanying source code implements this algorithm in the standard
way: the inherited attribute as the argument (i.e., the `environment')
and the synthesized as the result of the handler. The code shows
another implementation, exploring Haskell's laziness to make putAttr
also technically inherited, to accumulate the sequence of
updates. This implementation is compatible with deep handlers, and
lets us play with different notions of `backwardness'.In conclusion, we have implemented the effect of delaying an effectful computation until its main- or side-results are needed. We have demonstrated the lazy state as well as Attribute-Grammar--like computations with extensible effects. Our implementation makes the crucial delay points explicit. Furthermore, although the delayed computation is not effect-extensible, the overall computation is.
Edward Kmett: Comment on the r/haskell thread ``Are extensible effects a complete replacement for mtl?''
June 2015.
Nondet.hs [3K]
Exceptions.hs [2K]
Extensible.hs [12K]
TList.hs [2K]
The first modern implementation of extensible effects (February 2012). The first two files are the warm-up examples, leading to the full implementation.
Cartwright and Felleisen's paper appeared just a bit prior to Liang, Hudak and Jones' ``Monad Transformers and Modular interpreters'' that introduced monad transformers. In fact, the monad transformers paper mentions Cartwright and Felleisen's direct approach in footnote 1. Perhaps because Cartwright and Felleisen demonstrate their approach in an obscure dialect of Scheme, their work did not receive nearly as much attention as it vastly deserves.
We implement the enhanced EDLS in Haskell and add delimited control. To be precise, we implement the Base interpreter (whose sole operations are Loop and Error) and the following extensions: CBV lambda-calculus, Arithmetic, Storage, Control. The extensions can be added to the Base interpreter in any order and in any combination.
Our implementation has the following advantages over EDLS:
Our main departure from EDLS is is the removal of the `central authority'. There is no semantic `admin' function. Rather, admin is part of the source language and can be used at any place in the code. The `central authority' of EDLS must be an extensible function, requiring meta-language facilities to implement (such as quite non-standard Scheme modules). We do not have central authority. Rather, we have bureaucracy: each specific effect handler interprets its own effects as it can, throwing the rest `upstairs' for higher-level bureaucrats to deal with. Extensibility arises automatically.
We take the meaning of a program to be the union of Values and (unfulfilled) Actions. If the meaning of the program is a (non-bottom) value, the program is terminating. If the meaning of the program is an Action -- the program finished with an error, such as an action to access a non-existing storage cell, or shift without reset, or a user-raised error.
EDLS says, at the very top of p. 3, that the handle in the effect message ``is roughly a conventional continuation.'' Because the admin of EDLS is `outside' of the program (at the point of infinity, so to speak), its continuation indeed appears undelimited. By making our `admin' part of the source language, we recognize the handle in the effect message for what it is: a delimited continuation.
ExtensibleDS.hs [17K]
The complete code in Haskell98 (plus pattern guards), including several examples.
This code is not written in idiomatic Haskell and does not take
advantage of types at all. The ubiquitous projections from the
universal domain tantamount to ``dynamic typing.'' The code is
intentionally written to be close to the EDLS paper, emphasizing
denotational semantics, whose domains are untyped. One can certainly
do better, for example, employ user-defined datatypes for tagged
values, avoiding the ugly string-tagged values.
Having an Effect: The modern reconstruction of EDLS
oleg-at-okmij.org
Your comments, problem reports, questions are very welcome!
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