\begin{code}
{-# OPTIONS -fno-warn-tabs #-}
-- The above warning supression flag is a temporary kludge.
-- While working on this module you are encouraged to remove it and
-- detab the module (please do the detabbing in a separate patch). See
-- http://hackage.haskell.org/trac/ghc/wiki/Commentary/CodingStyle#TabsvsSpaces
-- for details
module TcCanonical(
canonicalize,
canOccursCheck, canEq, canEvVar,
rewriteWithFunDeps,
emitFDWorkAsWanted, emitFDWorkAsDerived,
StopOrContinue (..)
) where
#include "HsVersions.h"
import BasicTypes ( IPName )
import TcErrors
import TcRnTypes
import FunDeps
import qualified TcMType as TcM
import TcType
import Type
import Kind
import TcEvidence
import Class
import TyCon
import TypeRep
import Name ( Name )
import Var
import VarEnv
import Outputable
import Control.Monad ( when, unless, zipWithM, foldM )
import MonadUtils
import Control.Applicative ( (<|>) )
import TrieMap
import VarSet
import TcSMonad
import FastString
import Data.Maybe ( isNothing )
import Pair ( pSnd )
\end{code}
%************************************************************************
%* *
%* The Canonicaliser *
%* *
%************************************************************************
Note [Canonicalization]
~~~~~~~~~~~~~~~~~~~~~~~
Canonicalization converts a flat constraint to a canonical form. It is
unary (i.e. treats individual constraints one at a time), does not do
any zonking, but lives in TcS monad because it needs to create fresh
variables (for flattening) and consult the inerts (for efficiency).
The execution plan for canonicalization is the following:
1) Decomposition of equalities happens as necessary until we reach a
variable or type family in one side. There is no decomposition step
for other forms of constraints.
2) If, when we decompose, we discover a variable on the head then we
look at inert_eqs from the current inert for a substitution for this
variable and contine decomposing. Hence we lazily apply the inert
substitution if it is needed.
3) If no more decomposition is possible, we deeply apply the substitution
from the inert_eqs and continue with flattening.
4) During flattening, we examine whether we have already flattened some
function application by looking at all the CTyFunEqs with the same
function in the inert set. The reason for deeply applying the inert
substitution at step (3) is to maximise our chances of matching an
already flattened family application in the inert.
The net result is that a constraint coming out of the canonicalization
phase cannot be rewritten any further from the inerts (but maybe /it/ can
rewrite an inert or still interact with an inert in a further phase in the
simplifier.
\begin{code}
-- Informative results of canonicalization
data StopOrContinue
= ContinueWith Ct -- Either no canonicalization happened, or if some did
-- happen, it is still safe to just keep going with this
-- work item.
| Stop -- Some canonicalization happened, extra work is now in
-- the TcS WorkList.
instance Outputable StopOrContinue where
ppr Stop = ptext (sLit "Stop")
ppr (ContinueWith w) = ptext (sLit "ContinueWith") <+> ppr w
continueWith :: Ct -> TcS StopOrContinue
continueWith = return . ContinueWith
andWhenContinue :: TcS StopOrContinue
-> (Ct -> TcS StopOrContinue)
-> TcS StopOrContinue
andWhenContinue tcs1 tcs2
= do { r <- tcs1
; case r of
Stop -> return Stop
ContinueWith ct -> tcs2 ct }
\end{code}
Note [Caching for canonicals]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Our plan with pre-canonicalization is to be able to solve a constraint really fast from existing
bindings in TcEvBinds. So one may think that the condition (isCNonCanonical) is not necessary.
However consider the following setup:
InertSet = { [W] d1 : Num t }
WorkList = { [W] d2 : Num t, [W] c : t ~ Int}
Now, we prioritize equalities, but in our concrete example (should_run/mc17.hs) the first (d2) constraint
is dealt with first, because (t ~ Int) is an equality that only later appears in the worklist since it is
pulled out from a nested implication constraint. So, let's examine what happens:
- We encounter work item (d2 : Num t)
- Nothing is yet in EvBinds, so we reach the interaction with inerts
and set:
d2 := d1
and we discard d2 from the worklist. The inert set remains unaffected.
- Now the equation ([W] c : t ~ Int) is encountered and kicks-out (d1 : Num t) from the inerts.
Then that equation gets spontaneously solved, perhaps. We end up with:
InertSet : { [G] c : t ~ Int }
WorkList : { [W] d1 : Num t}
- Now we examine (d1), we observe that there is a binding for (Num t) in the evidence binds and
we set:
d1 := d2
and end up in a loop!
Now, the constraints that get kicked out from the inert set are always Canonical, so by restricting
the use of the pre-canonicalizer to NonCanonical constraints we eliminate this danger. Moreover, for
canonical constraints we already have good caching mechanisms (effectively the interaction solver)
and we are interested in reducing things like superclasses of the same non-canonical constraint being
generated hence I don't expect us to lose a lot by introducing the (isCNonCanonical) restriction.
A similar situation can arise in TcSimplify, at the end of the solve_wanteds function, where constraints
from the inert set are returned as new work -- our substCt ensures however that if they are not rewritten
by subst, they remain canonical and hence we will not attempt to solve them from the EvBinds. If on the
other hand they did get rewritten and are now non-canonical they will still not match the EvBinds, so we
are again good.
\begin{code}
-- Top-level canonicalization
-- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
canonicalize :: Ct -> TcS StopOrContinue
canonicalize ct@(CNonCanonical { cc_id = ev, cc_flavor = fl, cc_depth = d })
= do { traceTcS "canonicalize (non-canonical)" (ppr ct)
; {-# SCC "canEvVar" #-}
canEvVar ev (classifyPredType (evVarPred ev)) d fl }
canonicalize (CDictCan { cc_id = ev, cc_depth = d
, cc_flavor = fl
, cc_class = cls
, cc_tyargs = xis })
= {-# SCC "canClass" #-}
canClass d fl ev cls xis -- Do not add any superclasses
canonicalize (CTyEqCan { cc_id = ev, cc_depth = d
, cc_flavor = fl
, cc_tyvar = tv
, cc_rhs = xi })
= {-# SCC "canEqLeafTyVarLeftRec" #-}
canEqLeafTyVarLeftRec d fl ev tv xi
canonicalize (CFunEqCan { cc_id = ev, cc_depth = d
, cc_flavor = fl
, cc_fun = fn
, cc_tyargs = xis1
, cc_rhs = xi2 })
= {-# SCC "canEqLeafFunEqLeftRec" #-}
canEqLeafFunEqLeftRec d fl ev (fn,xis1) xi2
canonicalize (CIPCan { cc_id = ev, cc_depth = d
, cc_flavor = fl
, cc_ip_nm = nm
, cc_ip_ty = xi })
= canIP d fl ev nm xi
canonicalize (CIrredEvCan { cc_id = ev, cc_flavor = fl
, cc_depth = d
, cc_ty = xi })
= canIrred d fl ev xi
canEvVar :: EvVar -> PredTree
-> SubGoalDepth -> CtFlavor -> TcS StopOrContinue
canEvVar ev pred_classifier d fl
= case pred_classifier of
ClassPred cls tys -> canClass d fl ev cls tys
`andWhenContinue` emit_superclasses
EqPred ty1 ty2 -> canEq d fl ev ty1 ty2
IPPred nm ty -> canIP d fl ev nm ty
IrredPred ev_ty -> canIrred d fl ev ev_ty
TuplePred tys -> canTuple d fl ev tys
where emit_superclasses ct@(CDictCan {cc_id = v_new
, cc_tyargs = xis_new, cc_class = cls })
-- Add superclasses of this one here, See Note [Adding superclasses].
-- But only if we are not simplifying the LHS of a rule.
= do { sctxt <- getTcSContext
; unless (simplEqsOnly sctxt) $
newSCWorkFromFlavored d v_new fl cls xis_new
; continueWith ct }
emit_superclasses _ = panic "emit_superclasses of non-class!"
-- Tuple canonicalisation
-- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
canTuple :: SubGoalDepth -- Depth
-> CtFlavor -> EvVar -> [PredType] -> TcS StopOrContinue
canTuple d fl ev tys
= do { traceTcS "can_pred" (text "TuplePred!")
; evs <- zipWithM can_pred_tup_one tys [0..]
; if (isWanted fl) then
do {_unused_fl <- setEvBind ev (EvTupleMk evs) fl
; return Stop }
else return Stop }
where
can_pred_tup_one ty n
= do { evc <- newEvVar fl ty
; let ev' = evc_the_evvar evc
; fl' <- if isGivenOrSolved fl then
setEvBind ev' (EvTupleSel ev n) fl
else return fl
; when (isNewEvVar evc) $
addToWork (canEvVar ev' (classifyPredType (evVarPred ev')) d fl')
; return ev' }
-- Implicit Parameter Canonicalization
-- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
canIP :: SubGoalDepth -- Depth
-> CtFlavor -> EvVar
-> IPName Name -> Type -> TcS StopOrContinue
-- Precondition: EvVar is implicit parameter evidence
canIP d fl v nm ty
= -- Note [Canonical implicit parameter constraints] explains why it's
-- possible in principle to not flatten, but since flattening applies
-- the inert substitution we choose to flatten anyway.
do { (xi,co) <- flatten d fl (mkIPPred nm ty)
; let no_flattening = isTcReflCo co
; if no_flattening then
let IPPred _ xi_in = classifyPredType xi
in continueWith $ CIPCan { cc_id = v, cc_flavor = fl
, cc_ip_nm = nm, cc_ip_ty = xi_in
, cc_depth = d }
else do { evc <- newEvVar fl xi
; let v_new = evc_the_evvar evc
IPPred _ ip_xi = classifyPredType xi
; fl_new <- case fl of
Wanted {} -> setEvBind v (EvCast v_new co) fl
Given {} -> setEvBind v_new (EvCast v (mkTcSymCo co)) fl
Derived {} -> return fl
; if isNewEvVar evc then
continueWith $ CIPCan { cc_id = v_new
, cc_flavor = fl_new, cc_ip_nm = nm
, cc_ip_ty = ip_xi
, cc_depth = d }
else return Stop } }
\end{code}
Note [Canonical implicit parameter constraints]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
The type in a canonical implicit parameter constraint doesn't need to
be a xi (type-function-free type) since we can defer the flattening
until checking this type for equality with another type. If we
encounter two IP constraints with the same name, they MUST have the
same type, and at that point we can generate a flattened equality
constraint between the types. (On the other hand, the types in two
class constraints for the same class MAY be equal, so they need to be
flattened in the first place to facilitate comparing them.)
\begin{code}
-- Class Canonicalization
-- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
canClass :: SubGoalDepth -- Depth
-> CtFlavor -> EvVar
-> Class -> [Type] -> TcS StopOrContinue
-- Precondition: EvVar is class evidence
-- Note: Does NOT add superclasses, but the /caller/ is responsible for adding them!
canClass d fl v cls tys
= do { -- sctx <- getTcSContext
; (xis, cos) <- flattenMany d fl tys
; let co = mkTcTyConAppCo (classTyCon cls) cos
xi = mkClassPred cls xis
; let no_flattening = all isTcReflCo cos
-- No flattening, continue with canonical
; if no_flattening then
continueWith $ CDictCan { cc_id = v, cc_flavor = fl
, cc_tyargs = xis, cc_class = cls
, cc_depth = d }
-- Flattening happened
else do { evc <- newEvVar fl xi
; let v_new = evc_the_evvar evc
; fl_new <- case fl of
Wanted {} -> setEvBind v (EvCast v_new co) fl
Given {} -> setEvBind v_new (EvCast v (mkTcSymCo co)) fl
Derived {} -> return fl
-- Continue only if flat constraint is new
; if isNewEvVar evc then
continueWith $ CDictCan { cc_id = v_new, cc_flavor = fl_new
, cc_tyargs = xis, cc_class = cls
, cc_depth = d }
else return Stop } }
\end{code}
Note [Adding superclasses]
~~~~~~~~~~~~~~~~~~~~~~~~~~
Since dictionaries are canonicalized only once in their lifetime, the
place to add their superclasses is canonicalisation (The alternative
would be to do it during constraint solving, but we'd have to be
extremely careful to not repeatedly introduced the same superclass in
our worklist). Here is what we do:
For Givens:
We add all their superclasses as Givens.
For Wanteds:
Generally speaking we want to be able to add superclasses of
wanteds for two reasons:
(1) Oportunities for improvement. Example:
class (a ~ b) => C a b
Wanted constraint is: C alpha beta
We'd like to simply have C alpha alpha. Similar
situations arise in relation to functional dependencies.
(2) To have minimal constraints to quantify over:
For instance, if our wanted constraint is (Eq a, Ord a)
we'd only like to quantify over Ord a.
To deal with (1) above we only add the superclasses of wanteds
which may lead to improvement, that is: equality superclasses or
superclasses with functional dependencies.
We deal with (2) completely independently in TcSimplify. See
Note [Minimize by SuperClasses] in TcSimplify.
Moreover, in all cases the extra improvement constraints are
Derived. Derived constraints have an identity (for now), but
we don't do anything with their evidence. For instance they
are never used to rewrite other constraints.
See also [New Wanted Superclass Work] in TcInteract.
For Deriveds:
We do nothing.
Here's an example that demonstrates why we chose to NOT add
superclasses during simplification: [Comes from ticket #4497]
class Num (RealOf t) => Normed t
type family RealOf x
Assume the generated wanted constraint is:
RealOf e ~ e, Normed e
If we were to be adding the superclasses during simplification we'd get:
Num uf, Normed e, RealOf e ~ e, RealOf e ~ uf
==>
e ~ uf, Num uf, Normed e, RealOf e ~ e
==> [Spontaneous solve]
Num uf, Normed uf, RealOf uf ~ uf
While looks exactly like our original constraint. If we add the superclass again we'd loop.
By adding superclasses definitely only once, during canonicalisation, this situation can't
happen.
\begin{code}
newSCWorkFromFlavored :: SubGoalDepth -- Depth
-> EvVar -> CtFlavor -> Class -> [Xi] -> TcS ()
-- Returns superclasses, see Note [Adding superclasses]
newSCWorkFromFlavored d ev flavor cls xis
| isDerived flavor
= return () -- Deriveds don't yield more superclasses because we will
-- add them transitively in the case of wanteds.
| Just gk <- isGiven_maybe flavor
= case gk of
GivenOrig -> do { let sc_theta = immSuperClasses cls xis
; sc_vars <- mapM (newEvVar flavor) sc_theta
; sc_cts <- zipWithM (\scv ev_trm ->
do { let sc_evvar = evc_the_evvar scv
; _unused_fl <- setEvBind sc_evvar ev_trm flavor
-- unused because it's the same
; return $
CNonCanonical { cc_id = sc_evvar
, cc_flavor = flavor
, cc_depth = d }})
sc_vars [EvSuperClass ev n | n <- [0..]]
-- Emit now, canonicalize later in a lazier fashion
; traceTcS "newSCWorkFromFlavored" $
text "Emitting superclass work:" <+> ppr sc_cts
; updWorkListTcS $ appendWorkListCt sc_cts }
GivenSolved {} -> return ()
-- Seems very dangerous to add the superclasses for dictionaries that may be
-- partially solved because we may end up with evidence loops.
| isEmptyVarSet (tyVarsOfTypes xis)
= return () -- Wanteds with no variables yield no deriveds.
-- See Note [Improvement from Ground Wanteds]
| otherwise -- Wanted case, just add those SC that can lead to improvement.
= do { let sc_rec_theta = transSuperClasses cls xis
impr_theta = filter is_improvement_pty sc_rec_theta
Wanted wloc = flavor
; sc_cts <- mapM (\pty -> do { scv <- newEvVar (Derived wloc) pty
; if isNewEvVar scv then
return [ CNonCanonical { cc_id = evc_the_evvar scv
, cc_flavor = Derived wloc
, cc_depth = d } ]
else return [] }
) impr_theta
; let sc_cts_flat = concat sc_cts
; traceTcS "newSCWorkFromFlavored" (text "Emitting superclass work:" <+> ppr sc_cts_flat)
; updWorkListTcS $ appendWorkListCt sc_cts_flat }
is_improvement_pty :: PredType -> Bool
-- Either it's an equality, or has some functional dependency
is_improvement_pty ty = go (classifyPredType ty)
where
go (EqPred {}) = True
go (ClassPred cls _tys) = not $ null fundeps
where (_,fundeps) = classTvsFds cls
go (IPPred {}) = False
go (TuplePred ts) = any is_improvement_pty ts
go (IrredPred {}) = True -- Might have equalities after reduction?
\end{code}
\begin{code}
-- Irreducibles canonicalization
-- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
canIrred :: SubGoalDepth -- Depth
-> CtFlavor -> EvVar -> TcType -> TcS StopOrContinue
-- Precondition: ty not a tuple and no other evidence form
canIrred d fl v ty
= do { traceTcS "can_pred" (text "IrredPred = " <+> ppr ty)
; (xi,co) <- flatten d fl ty -- co :: xi ~ ty
; let no_flattening = xi `eqType` ty
-- In this particular case it is not safe to
-- say 'isTcReflCo' because the new constraint may
-- be reducible!
; if no_flattening then
continueWith $ CIrredEvCan { cc_id = v, cc_flavor = fl
, cc_ty = xi, cc_depth = d }
else do
{ -- Flattening consults and applies family equations from the
-- inerts, so 'xi' may become reducible. So just recursively
-- canonicalise the resulting evidence variable
evc <- newEvVar fl xi
; let v' = evc_the_evvar evc
; fl' <- case fl of
Wanted {} -> setEvBind v (EvCast v' co) fl
Given {} -> setEvBind v' (EvCast v (mkTcSymCo co)) fl
Derived {} -> return fl
; if isNewEvVar evc then
canEvVar v' (classifyPredType (evVarPred v')) d fl'
else
return Stop }
}
\end{code}
%************************************************************************
%* *
%* Flattening (eliminating all function symbols) *
%* *
%************************************************************************
Note [Flattening]
~~~~~~~~~~~~~~~~~~~~
flatten ty ==> (xi, cc)
where
xi has no type functions
cc = Auxiliary given (equality) constraints constraining
the fresh type variables in xi. Evidence for these
is always the identity coercion, because internally the
fresh flattening skolem variables are actually identified
with the types they have been generated to stand in for.
Note that it is flatten's job to flatten *every type function it sees*.
flatten is only called on *arguments* to type functions, by canEqGiven.
Recall that in comments we use alpha[flat = ty] to represent a
flattening skolem variable alpha which has been generated to stand in
for ty.
----- Example of flattening a constraint: ------
flatten (List (F (G Int))) ==> (xi, cc)
where
xi = List alpha
cc = { G Int ~ beta[flat = G Int],
F beta ~ alpha[flat = F beta] }
Here
* alpha and beta are 'flattening skolem variables'.
* All the constraints in cc are 'given', and all their coercion terms
are the identity.
NB: Flattening Skolems only occur in canonical constraints, which
are never zonked, so we don't need to worry about zonking doing
accidental unflattening.
Note that we prefer to leave type synonyms unexpanded when possible,
so when the flattener encounters one, it first asks whether its
transitive expansion contains any type function applications. If so,
it expands the synonym and proceeds; if not, it simply returns the
unexpanded synonym.
\begin{code}
-- Flatten a bunch of types all at once.
flattenMany :: SubGoalDepth -- Depth
-> CtFlavor -> [Type] -> TcS ([Xi], [TcCoercion])
-- Coercions :: Xi ~ Type
-- Returns True iff (no flattening happened)
flattenMany d ctxt tys
= -- pprTrace "flattenMany" empty $
go tys
where go [] = return ([],[])
go (ty:tys) = do { (xi,co) <- flatten d ctxt ty
; (xis,cos) <- go tys
; return (xi:xis,co:cos) }
-- Flatten a type to get rid of type function applications, returning
-- the new type-function-free type, and a collection of new equality
-- constraints. See Note [Flattening] for more detail.
flatten :: SubGoalDepth -- Depth
-> CtFlavor -> TcType -> TcS (Xi, TcCoercion)
-- Postcondition: Coercion :: Xi ~ TcType
flatten d ctxt ty
| Just ty' <- tcView ty
= do { (xi, co) <- flatten d ctxt ty'
; return (xi,co) }
-- DV: The following is tedious to do but maybe we should return to this
-- Preserve type synonyms if possible
-- ; if no_flattening
-- then return (xi, mkTcReflCo xi,no_flattening) -- Importantly, not xi!
-- else return (xi,co,no_flattening)
-- }
flatten d ctxt v@(TyVarTy _)
= do { ieqs <- getInertEqs
; let co = liftInertEqsTy ieqs ctxt v -- co : v ~ ty
ty = pSnd (tcCoercionKind co)
; if v `eqType` ty then
return (ty,mkTcReflCo ty)
else -- NB recursive call. Why? See Note [Non-idempotent inert substitution]
-- Actually I believe that applying the substition only *twice* will suffice
do { (ty_final,co') <- flatten d ctxt ty -- co' : ty_final ~ ty
; return (ty_final,co' `mkTcTransCo` mkTcSymCo co) } }
\end{code}
Note [Non-idempotent inert substitution]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
The inert substitution is not idempotent in the broad sense. It is only idempotent in
that it cannot rewrite the RHS of other inert equalities any further. An example of such
an inert substitution is:
[Ś] g1 : ta8 ~ ta4
[W] g2 : ta4 ~ a5Fj
Observe that the wanted cannot rewrite the solved goal, despite the fact that ta4 appears on
an RHS of an equality. Now, imagine a constraint:
[W] g3: ta8 ~ Int
coming in. If we simply apply once the inert substitution we will get:
[W] g3_1: ta4 ~ Int
and because potentially ta4 is untouchable we will try to insert g3_1 in the inert set,
getting a panic since the inert only allows ONE equation per LHS type variable (as it
should).
For this reason, when we reach to flatten a type variable, we flatten it recursively,
so that we can make sure that the inert substitution /is/ fully applied.
This insufficient rewriting was the reason for #5668.
\begin{code}
flatten d ctxt (AppTy ty1 ty2)
= do { (xi1,co1) <- flatten d ctxt ty1
; (xi2,co2) <- flatten d ctxt ty2
; return (mkAppTy xi1 xi2, mkTcAppCo co1 co2) }
flatten d ctxt (FunTy ty1 ty2)
= do { (xi1,co1) <- flatten d ctxt ty1
; (xi2,co2) <- flatten d ctxt ty2
; return (mkFunTy xi1 xi2, mkTcFunCo co1 co2) }
flatten d fl (TyConApp tc tys)
-- For a normal type constructor or data family application, we just
-- recursively flatten the arguments.
| not (isSynFamilyTyCon tc)
= do { (xis,cos) <- flattenMany d fl tys
; return (mkTyConApp tc xis, mkTcTyConAppCo tc cos) }
-- Otherwise, it's a type function application, and we have to
-- flatten it away as well, and generate a new given equality constraint
-- between the application and a newly generated flattening skolem variable.
| otherwise
= ASSERT( tyConArity tc <= length tys ) -- Type functions are saturated
do { (xis, cos) <- flattenMany d fl tys
; let (xi_args, xi_rest) = splitAt (tyConArity tc) xis
-- The type function might be *over* saturated
-- in which case the remaining arguments should
-- be dealt with by AppTys
fam_ty = mkTyConApp tc xi_args
; (ret_co, rhs_xi, ct) <-
do { is_cached <- getCachedFlatEq tc xi_args fl Any
; case is_cached of
Just (rhs_xi,ret_eq) ->
do { traceTcS "is_cached!" $ ppr ret_eq
; return (ret_eq, rhs_xi, []) }
Nothing
| isGivenOrSolved fl ->
do { rhs_xi_var <- newFlattenSkolemTy fam_ty
; (fl',eqv)
<- newGivenEqVar fl fam_ty rhs_xi_var (mkTcReflCo fam_ty)
; let ct = CFunEqCan { cc_id = eqv
, cc_flavor = fl' -- Given
, cc_fun = tc
, cc_tyargs = xi_args
, cc_rhs = rhs_xi_var
, cc_depth = d }
-- Update the flat cache: just an optimisation!
; updateFlatCache eqv fl' tc xi_args rhs_xi_var WhileFlattening
; return (mkTcCoVarCo eqv, rhs_xi_var, [ct]) }
| otherwise ->
-- Derived or Wanted: make a new /unification/ flatten variable
do { rhs_xi_var <- newFlexiTcSTy (typeKind fam_ty)
; let wanted_flavor = mkWantedFlavor fl
; evc <- newEqVar wanted_flavor fam_ty rhs_xi_var
; let eqv = evc_the_evvar evc -- Not going to be cached
ct = CFunEqCan { cc_id = eqv
, cc_flavor = wanted_flavor
-- Always Wanted, not Derived
, cc_fun = tc
, cc_tyargs = xi_args
, cc_rhs = rhs_xi_var
, cc_depth = d }
-- Update the flat cache: just an optimisation!
; updateFlatCache eqv fl tc xi_args rhs_xi_var WhileFlattening
; return (mkTcCoVarCo eqv, rhs_xi_var, [ct]) } }
-- Emit the flat constraints
; updWorkListTcS $ appendWorkListEqs ct
; let (cos_args, cos_rest) = splitAt (tyConArity tc) cos
; return ( mkAppTys rhs_xi xi_rest -- NB mkAppTys: rhs_xi might not be a type variable
-- cf Trac #5655
, mkTcAppCos (mkTcSymCo ret_co `mkTcTransCo` mkTcTyConAppCo tc cos_args)
cos_rest
) }
flatten d ctxt ty@(ForAllTy {})
-- We allow for-alls when, but only when, no type function
-- applications inside the forall involve the bound type variables.
= do { let (tvs, rho) = splitForAllTys ty
; when (under_families tvs rho) $ flattenForAllErrorTcS ctxt ty
; (rho', co) <- flatten d ctxt rho
; return (mkForAllTys tvs rho', foldr mkTcForAllCo co tvs) }
where under_families tvs rho
= go (mkVarSet tvs) rho
where go _bound (TyVarTy _tv) = False
go bound (TyConApp tc tys)
| isSynFamilyTyCon tc
, (args,rest) <- splitAt (tyConArity tc) tys
= (tyVarsOfTypes args `intersectsVarSet` bound) || any (go bound) rest
| otherwise = any (go bound) tys
go bound (FunTy arg res) = go bound arg || go bound res
go bound (AppTy fun arg) = go bound fun || go bound arg
go bound (ForAllTy tv ty) = go (bound `extendVarSet` tv) ty
getCachedFlatEq :: TyCon -> [Xi] -> CtFlavor
-> FlatEqOrigin
-> TcS (Maybe (Xi, TcCoercion))
-- Returns a coercion between (TyConApp tc xi_args ~ xi) if such an inert item exists
-- But also applies the substitution to the item via calling flatten recursively
getCachedFlatEq tc xi_args fl feq_origin
= do { let pty = mkTyConApp tc xi_args
; traceTcS "getCachedFlatEq" $ ppr (mkTyConApp tc xi_args)
; flat_cache <- getTcSEvVarFlatCache
; inerts <- getTcSInerts
; case lookupFunEq pty fl (inert_funeqs inerts) of
Nothing
-> lookup_in_flat_cache pty flat_cache
res -> return res }
where lookup_in_flat_cache pty flat_cache
= case lookupTM pty flat_cache of
Just (co',(xi',fl',when_generated)) -- ev' :: (TyConApp tc xi_args) ~ xi'
| fl' `canRewrite` fl
, feq_origin `origin_matches` when_generated
-> do { traceTcS "getCachedFlatEq" $ text "success!"
; (xi'',co) <- flatten 0 fl' xi' -- co :: xi'' ~ xi'
-- The only purpose of this flattening is to apply the
-- inert substitution (since everything in the flat cache
-- by construction will have a family-free RHS.
; return $ Just (xi'', co' `mkTcTransCo` (mkTcSymCo co)) }
_ -> do { traceTcS "getCachedFlatEq" $ text "failure!" <+> pprEvVarCache flat_cache
; return Nothing }
-----------------
addToWork :: TcS StopOrContinue -> TcS ()
addToWork tcs_action = tcs_action >>= stop_or_emit
where stop_or_emit Stop = return ()
stop_or_emit (ContinueWith ct) = updWorkListTcS $
extendWorkListCt ct
canEqEvVarsCreated :: SubGoalDepth
-> [CtFlavor] -> [EvVarCreated] -> [Type] -> [Type]
-> TcS StopOrContinue
canEqEvVarsCreated _d _fl [] _ _ = return Stop
canEqEvVarsCreated d (fl:fls) (evc:evcs) (ty1:tys1) (ty2:tys2)
| isNewEvVar evc
= let do_one evc0 sy1 sy2
| isNewEvVar evc0
= canEq_ d fl (evc_the_evvar evc0) sy1 sy2
| otherwise = return ()
in do { _unused <- zipWith3M do_one evcs tys1 tys2
; canEq d fl (evc_the_evvar evc) ty1 ty2 }
| otherwise
= canEqEvVarsCreated d fls evcs tys1 tys2
canEqEvVarsCreated _ _ _ _ _ = return Stop
canEq_ :: SubGoalDepth
-> CtFlavor -> EqVar -> Type -> Type -> TcS ()
canEq_ d fl eqv ty1 ty2 = addToWork (canEq d fl eqv ty1 ty2)
canEq :: SubGoalDepth
-> CtFlavor -> EqVar -> Type -> Type -> TcS StopOrContinue
canEq _d fl eqv ty1 ty2
| eqType ty1 ty2 -- Dealing with equality here avoids
-- later spurious occurs checks for a~a
= do { when (isWanted fl) $
do { _ <- setEqBind eqv (mkTcReflCo ty1) fl; return () }
; return Stop }
-- Split up an equality between function types into two equalities.
canEq d fl eqv (FunTy s1 t1) (FunTy s2 t2)
= do { argeqv <- newEqVar fl s1 s2
; reseqv <- newEqVar fl t1 t2
; let argeqv_v = evc_the_evvar argeqv
reseqv_v = evc_the_evvar reseqv
; (fl1,fl2) <- case fl of
Wanted {} ->
do { _ <- setEqBind eqv (mkTcFunCo (mkTcCoVarCo argeqv_v) (mkTcCoVarCo reseqv_v)) fl
; return (fl,fl) }
Given {} ->
do { fl1 <- setEqBind argeqv_v (mkTcNthCo 0 (mkTcCoVarCo eqv)) fl
; fl2 <- setEqBind reseqv_v (mkTcNthCo 1 (mkTcCoVarCo eqv)) fl
; return (fl1,fl2)
}
Derived {} ->
return (fl,fl)
; canEqEvVarsCreated d [fl2,fl1] [reseqv,argeqv] [t1,s1] [t2,s2] }
-- If one side is a variable, orient and flatten,
-- WITHOUT expanding type synonyms, so that we tend to
-- substitute a ~ Age rather than a ~ Int when @type Age = Int@
canEq d fl eqv ty1@(TyVarTy {}) ty2
= canEqLeaf d fl eqv ty1 ty2
canEq d fl eqv ty1 ty2@(TyVarTy {})
= canEqLeaf d fl eqv ty1 ty2
canEq d fl eqv ty1@(TyConApp fn tys) ty2
| isSynFamilyTyCon fn, length tys == tyConArity fn
= canEqLeaf d fl eqv ty1 ty2
canEq d fl eqv ty1 ty2@(TyConApp fn tys)
| isSynFamilyTyCon fn, length tys == tyConArity fn
= canEqLeaf d fl eqv ty1 ty2
canEq d fl eqv (TyConApp tc1 tys1) (TyConApp tc2 tys2)
| isDecomposableTyCon tc1 && isDecomposableTyCon tc2
, tc1 == tc2
, length tys1 == length tys2
= -- Generate equalities for each of the corresponding arguments
do { let (kis1, tys1') = span isKind tys1
(_kis2, tys2') = span isKind tys2
; let kicos = map mkTcReflCo kis1
; argeqvs <- zipWithM (newEqVar fl) tys1' tys2'
; fls <- case fl of
Wanted {} ->
do { _ <- setEqBind eqv
(mkTcTyConAppCo tc1 (kicos ++ map (mkTcCoVarCo . evc_the_evvar) argeqvs)) fl
; return (map (\_ -> fl) argeqvs) }
Given {} ->
let do_one argeqv n = setEqBind (evc_the_evvar argeqv)
(mkTcNthCo n (mkTcCoVarCo eqv)) fl
in zipWithM do_one argeqvs [(length kicos)..]
Derived {} -> return (map (\_ -> fl) argeqvs)
; canEqEvVarsCreated d fls argeqvs tys1' tys2' }
-- See Note [Equality between type applications]
-- Note [Care with type applications] in TcUnify
canEq d fl eqv ty1 ty2
| Nothing <- tcView ty1 -- Naked applications ONLY
, Nothing <- tcView ty2 -- See Note [Naked given applications]
, Just (s1,t1) <- tcSplitAppTy_maybe ty1
, Just (s2,t2) <- tcSplitAppTy_maybe ty2
= ASSERT( not (isKind t1) && not (isKind t2) )
if isGivenOrSolved fl then
do { traceTcS "canEq/(app case)" $
text "Ommitting decomposition of given equality between: "
<+> ppr ty1 <+> text "and" <+> ppr ty2
-- We cannot decompose given applications
-- because we no longer have 'left' and 'right'
; return Stop }
else
do { evc1 <- newEqVar fl s1 s2
; evc2 <- newEqVar fl t1 t2
; let eqv1 = evc_the_evvar evc1
eqv2 = evc_the_evvar evc2
; when (isWanted fl) $
do { _ <- setEqBind eqv (mkTcAppCo (mkTcCoVarCo eqv1) (mkTcCoVarCo eqv2)) fl
; return () }
; canEqEvVarsCreated d [fl,fl] [evc1,evc2] [s1,t1] [s2,t2] }
canEq d fl eqv s1@(ForAllTy {}) s2@(ForAllTy {})
| tcIsForAllTy s1, tcIsForAllTy s2,
Wanted {} <- fl
= canEqFailure d fl eqv
| otherwise
= do { traceTcS "Ommitting decomposition of given polytype equality" (pprEq s1 s2)
; return Stop }
-- Finally expand any type synonym applications.
canEq d fl eqv ty1 ty2 | Just ty1' <- tcView ty1 = canEq d fl eqv ty1' ty2
canEq d fl eqv ty1 ty2 | Just ty2' <- tcView ty2 = canEq d fl eqv ty1 ty2'
canEq d fl eqv _ _ = canEqFailure d fl eqv
canEqFailure :: SubGoalDepth
-> CtFlavor -> EvVar -> TcS StopOrContinue
canEqFailure d fl eqv = do { emitFrozenError fl eqv d; return Stop }
\end{code}
Note [Naked given applications]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Consider:
data A a
type T a = A a
and the given equality:
[G] A a ~ T Int
We will reach the case canEq where we do a tcSplitAppTy_maybe, but if
we dont have the guards (Nothing <- tcView ty1) (Nothing <- tcView
ty2) then the given equation is going to fall through and get
completely forgotten!
What we want instead is this clause to apply only when there is no
immediate top-level synonym; if there is one it will be later on
unfolded by the later stages of canEq.
Test-case is in typecheck/should_compile/GivenTypeSynonym.hs
Note [Equality between type applications]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
If we see an equality of the form s1 t1 ~ s2 t2 we can always split
it up into s1 ~ s2 /\ t1 ~ t2, since s1 and s2 can't be type
functions (type functions use the TyConApp constructor, which never
shows up as the LHS of an AppTy). Other than type functions, types
in Haskell are always
(1) generative: a b ~ c d implies a ~ c, since different type
constructors always generate distinct types
(2) injective: a b ~ a d implies b ~ d; we never generate the
same type from different type arguments.
Note [Canonical ordering for equality constraints]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Implemented as (<+=) below:
- Type function applications always come before anything else.
- Variables always come before non-variables (other than type
function applications).
Note that we don't need to unfold type synonyms on the RHS to check
the ordering; that is, in the rules above it's OK to consider only
whether something is *syntactically* a type function application or
not. To illustrate why this is OK, suppose we have an equality of the
form 'tv ~ S a b c', where S is a type synonym which expands to a
top-level application of the type function F, something like
type S a b c = F d e
Then to canonicalize 'tv ~ S a b c' we flatten the RHS, and since S's
expansion contains type function applications the flattener will do
the expansion and then generate a skolem variable for the type
function application, so we end up with something like this:
tv ~ x
F d e ~ x
where x is the skolem variable. This is one extra equation than
absolutely necessary (we could have gotten away with just 'F d e ~ tv'
if we had noticed that S expanded to a top-level type function
application and flipped it around in the first place) but this way
keeps the code simpler.
Unlike the OutsideIn(X) draft of May 7, 2010, we do not care about the
ordering of tv ~ tv constraints. There are several reasons why we
might:
(1) In order to be able to extract a substitution that doesn't
mention untouchable variables after we are done solving, we might
prefer to put touchable variables on the left. However, in and
of itself this isn't necessary; we can always re-orient equality
constraints at the end if necessary when extracting a substitution.
(2) To ensure termination we might think it necessary to put
variables in lexicographic order. However, this isn't actually
necessary as outlined below.
While building up an inert set of canonical constraints, we maintain
the invariant that the equality constraints in the inert set form an
acyclic rewrite system when viewed as L-R rewrite rules. Moreover,
the given constraints form an idempotent substitution (i.e. none of
the variables on the LHS occur in any of the RHS's, and type functions
never show up in the RHS at all), the wanted constraints also form an
idempotent substitution, and finally the LHS of a given constraint
never shows up on the RHS of a wanted constraint. There may, however,
be a wanted LHS that shows up in a given RHS, since we do not rewrite
given constraints with wanted constraints.
Suppose we have an inert constraint set
tg_1 ~ xig_1 -- givens
tg_2 ~ xig_2
...
tw_1 ~ xiw_1 -- wanteds
tw_2 ~ xiw_2
...
where each t_i can be either a type variable or a type function
application. Now suppose we take a new canonical equality constraint,
t' ~ xi' (note among other things this means t' does not occur in xi')
and try to react it with the existing inert set. We show by induction
on the number of t_i which occur in t' ~ xi' that this process will
terminate.
There are several ways t' ~ xi' could react with an existing constraint:
TODO: finish this proof. The below was for the case where the entire
inert set is an idempotent subustitution...
(b) We could have t' = t_j for some j. Then we obtain the new
equality xi_j ~ xi'; note that neither xi_j or xi' contain t_j. We
now canonicalize the new equality, which may involve decomposing it
into several canonical equalities, and recurse on these. However,
none of the new equalities will contain t_j, so they have fewer
occurrences of the t_i than the original equation.
(a) We could have t_j occurring in xi' for some j, with t' /=
t_j. Then we substitute xi_j for t_j in xi' and continue. However,
since none of the t_i occur in xi_j, we have decreased the
number of t_i that occur in xi', since we eliminated t_j and did not
introduce any new ones.
\begin{code}
data TypeClassifier
= FskCls TcTyVar -- ^ Flatten skolem
| VarCls TcTyVar -- ^ Non-flatten-skolem variable
| FunCls TyCon [Type] -- ^ Type function, exactly saturated
| OtherCls TcType -- ^ Neither of the above
{- Useless these days!
unClassify :: TypeClassifier -> TcType
unClassify (VarCls tv) = TyVarTy tv
unClassify (FskCls tv) = TyVarTy tv
unClassify (FunCls fn tys) = TyConApp fn tys
unClassify (OtherCls ty) = ty
-}
classify :: TcType -> TypeClassifier
classify (TyVarTy tv)
| isTcTyVar tv,
FlatSkol {} <- tcTyVarDetails tv = FskCls tv
| otherwise = VarCls tv
classify (TyConApp tc tys) | isSynFamilyTyCon tc
, tyConArity tc == length tys
= FunCls tc tys
classify ty | Just ty' <- tcView ty
= case classify ty' of
OtherCls {} -> OtherCls ty
var_or_fn -> var_or_fn
| otherwise
= OtherCls ty
-- See note [Canonical ordering for equality constraints].
reOrient :: CtFlavor -> TypeClassifier -> TypeClassifier -> Bool
-- (t1 `reOrient` t2) responds True
-- iff we should flip to (t2~t1)
-- We try to say False if possible, to minimise evidence generation
--
-- Postcondition: After re-orienting, first arg is not OTherCls
reOrient _fl (OtherCls {}) (FunCls {}) = True
reOrient _fl (OtherCls {}) (FskCls {}) = True
reOrient _fl (OtherCls {}) (VarCls {}) = True
reOrient _fl (OtherCls {}) (OtherCls {}) = panic "reOrient" -- One must be Var/Fun
reOrient _fl (FunCls {}) (VarCls _tv) = False
-- But consider the following variation: isGiven fl && isMetaTyVar tv
-- See Note [No touchables as FunEq RHS] in TcSMonad
reOrient _fl (FunCls {}) _ = False -- Fun/Other on rhs
reOrient _fl (VarCls {}) (FunCls {}) = True
reOrient _fl (VarCls {}) (FskCls {}) = False
reOrient _fl (VarCls {}) (OtherCls {}) = False
reOrient _fl (VarCls tv1) (VarCls tv2)
| isMetaTyVar tv2 && not (isMetaTyVar tv1) = True
| otherwise = False
-- Just for efficiency, see CTyEqCan invariants
reOrient _fl (FskCls {}) (VarCls tv2) = isMetaTyVar tv2
-- Just for efficiency, see CTyEqCan invariants
reOrient _fl (FskCls {}) (FskCls {}) = False
reOrient _fl (FskCls {}) (FunCls {}) = True
reOrient _fl (FskCls {}) (OtherCls {}) = False
------------------
canEqLeaf :: SubGoalDepth -- Depth
-> CtFlavor -> EqVar
-> Type -> Type
-> TcS StopOrContinue
-- Canonicalizing "leaf" equality constraints which cannot be
-- decomposed further (ie one of the types is a variable or
-- saturated type function application).
-- Preconditions:
-- * one of the two arguments is variable or family applications
-- * the two types are not equal (looking through synonyms)
canEqLeaf d fl eqv s1 s2
| cls1 `re_orient` cls2
= do { traceTcS "canEqLeaf (reorienting)" $ ppr eqv <+> dcolon <+> pprEq s1 s2
; delCachedEvVar eqv fl
; evc <- newEqVar fl s2 s1
; let eqv' = evc_the_evvar evc
; fl' <- case fl of
Wanted {} -> setEqBind eqv (mkTcSymCo (mkTcCoVarCo eqv')) fl
Given {} -> setEqBind eqv' (mkTcSymCo (mkTcCoVarCo eqv)) fl
Derived {} -> return fl
; if isNewEvVar evc then
do { canEqLeafOriented d fl' eqv' s2 s1 }
else return Stop
}
| otherwise
= do { traceTcS "canEqLeaf" $ ppr (mkEqPred (s1,s2))
; canEqLeafOriented d fl eqv s1 s2 }
where
re_orient = reOrient fl
cls1 = classify s1
cls2 = classify s2
canEqLeafOriented :: SubGoalDepth -- Depth
-> CtFlavor -> EqVar
-> TcType -> TcType -> TcS StopOrContinue
-- By now s1 will either be a variable or a type family application
canEqLeafOriented d fl eqv s1 s2
| let k1 = typeKind s1
, let k2 = typeKind s2
-- Establish kind invariants for CFunEqCan and CTyEqCan
= do { are_compat <- compatKindTcS k1 k2
; can_unify <- if not are_compat
then unifyKindTcS s1 s2 k1 k2
else return False
-- If the kinds cannot be unified or are not compatible, don't fail
-- right away; instead, emit a frozen error
; if (not are_compat && not can_unify) then
canEqFailure d fl eqv
else can_eq_kinds_ok d fl eqv s1 s2 }
where can_eq_kinds_ok d fl eqv s1 s2
| Just (fn,tys1) <- splitTyConApp_maybe s1
= canEqLeafFunEqLeftRec d fl eqv (fn,tys1) s2
| Just tv <- getTyVar_maybe s1
= canEqLeafTyVarLeftRec d fl eqv tv s2
| otherwise
= pprPanic "canEqLeafOriented" $
text "Non-variable or non-family equality LHS" <+> ppr eqv <+>
dcolon <+> ppr (evVarPred eqv)
canEqLeafFunEqLeftRec :: SubGoalDepth
-> CtFlavor
-> EqVar
-> (TyCon,[TcType]) -> TcType -> TcS StopOrContinue
canEqLeafFunEqLeftRec d fl eqv (fn,tys1) ty2 -- eqv :: F tys1 ~ ty2
= do { traceTcS "canEqLeafFunEqLeftRec" $ pprEq (mkTyConApp fn tys1) ty2
; (xis1,cos1) <-
{-# SCC "flattenMany" #-}
flattenMany d fl tys1 -- Flatten type function arguments
-- cos1 :: xis1 ~ tys1
-- ; inerts <- getTcSInerts
-- ; let fam_eqs = inert_funeqs inerts
; let flat_ty = mkTyConApp fn xis1
; is_cached <- getCachedFlatEq fn xis1 fl WhenSolved
-- Lookup if we have solved this goal already
{-
; let is_cached = {-# SCC "lookupFunEq" #-}
lookupFunEq flat_ty fl fam_eqs
-}
; let no_flattening = all isTcReflCo cos1
; if no_flattening && isNothing is_cached then
canEqLeafFunEqLeft d fl eqv (fn,xis1) ty2
else do
{ let (final_co, final_ty)
| no_flattening -- Just in inerts
, Just (rhs_ty, ret_eq) <- is_cached
= (mkTcSymCo ret_eq, rhs_ty)
| Nothing <- is_cached -- Just flattening
= (mkTcTyConAppCo fn cos1, flat_ty)
| Just (rhs_ty, ret_eq) <- is_cached -- Both
= (mkTcSymCo ret_eq `mkTcTransCo` mkTcTyConAppCo fn cos1, rhs_ty)
| otherwise = panic "No flattening and not cached!"
; delCachedEvVar eqv fl
; evc <- newEqVar fl final_ty ty2
; let new_eqv = evc_the_evvar evc
; fl' <- case fl of
Wanted {} -> setEqBind eqv
(mkTcSymCo final_co `mkTcTransCo` (mkTcCoVarCo new_eqv)) fl
Given {} -> setEqBind new_eqv (final_co `mkTcTransCo` (mkTcCoVarCo eqv)) fl
Derived {} -> return fl
; if isNewEvVar evc then
if isNothing is_cached then
{-# SCC "canEqLeafFunEqLeft" #-}
canEqLeafFunEqLeft d fl' new_eqv (fn,xis1) ty2
else
canEq (d+1) fl' new_eqv final_ty ty2
else return Stop
}
}
lookupFunEq :: PredType -> CtFlavor -> TypeMap Ct -> Maybe (TcType, TcCoercion)
lookupFunEq pty fl fam_eqs = lookup_funeq pty fam_eqs
where lookup_funeq pty fam_eqs
| Just ct <- lookupTM pty fam_eqs
, cc_flavor ct `canRewrite` fl
= Just (cc_rhs ct, mkTcCoVarCo (cc_id ct))
| otherwise
= Nothing
canEqLeafFunEqLeft :: SubGoalDepth -- Depth
-> CtFlavor -> EqVar -> (TyCon,[Xi])
-> TcType -> TcS StopOrContinue
-- Precondition: No more flattening is needed for the LHS
canEqLeafFunEqLeft d fl eqv (fn,xis1) s2
= {-# SCC "canEqLeafFunEqLeft" #-}
do { traceTcS "canEqLeafFunEqLeft" $ pprEq (mkTyConApp fn xis1) s2
; (xi2,co2) <-
{-# SCC "flatten" #-}
flatten d fl s2 -- co2 :: xi2 ~ s2
; let no_flattening_happened = isTcReflCo co2
; if no_flattening_happened then
continueWith $ CFunEqCan { cc_id = eqv
, cc_flavor = fl
, cc_fun = fn
, cc_tyargs = xis1
, cc_rhs = xi2
, cc_depth = d }
else do { delCachedEvVar eqv fl
; evc <-
{-# SCC "newEqVar" #-}
newEqVar fl (mkTyConApp fn xis1) xi2
; let new_eqv = evc_the_evvar evc -- F xis1 ~ xi2
new_cv = mkTcCoVarCo new_eqv
cv = mkTcCoVarCo eqv -- F xis1 ~ s2
; fl' <- case fl of
Wanted {} -> setEqBind eqv (new_cv `mkTcTransCo` co2) fl
Given {} -> setEqBind new_eqv (cv `mkTcTransCo` mkTcSymCo co2) fl
Derived {} -> return fl
; if isNewEvVar evc then
do { continueWith $
CFunEqCan { cc_id = new_eqv
, cc_flavor = fl'
, cc_fun = fn
, cc_tyargs = xis1
, cc_rhs = xi2
, cc_depth = d } }
else return Stop } }
canEqLeafTyVarLeftRec :: SubGoalDepth
-> CtFlavor -> EqVar
-> TcTyVar -> TcType -> TcS StopOrContinue
canEqLeafTyVarLeftRec d fl eqv tv s2 -- eqv :: tv ~ s2
= do { traceTcS "canEqLeafTyVarLeftRec" $ pprEq (mkTyVarTy tv) s2
; (xi1,co1) <- flatten d fl (mkTyVarTy tv) -- co1 :: xi1 ~ tv
; case isTcReflCo co1 of
True -- If reflco and variable, just go on
| Just tv' <- getTyVar_maybe xi1
-> canEqLeafTyVarLeft d fl eqv tv' s2
_ -> -- If not a variable or not refl co, must rewrite and go on
do { delCachedEvVar eqv fl
; evc <- newEqVar fl xi1 s2 -- new_ev :: xi1 ~ s2
; let new_ev = evc_the_evvar evc
; fl' <- case fl of
Wanted {} -> setEqBind eqv
(mkTcSymCo co1 `mkTcTransCo` mkTcCoVarCo new_ev) fl
Given {} -> setEqBind new_ev
(co1 `mkTcTransCo` mkTcCoVarCo eqv) fl
Derived {} -> return fl
; if isNewEvVar evc then
do { canEq d fl' new_ev xi1 s2 }
else return Stop
}
}
canEqLeafTyVarLeft :: SubGoalDepth -- Depth
-> CtFlavor -> EqVar
-> TcTyVar -> TcType -> TcS StopOrContinue
-- Precondition LHS is fully rewritten from inerts (but not RHS)
canEqLeafTyVarLeft d fl eqv tv s2 -- eqv : tv ~ s2
= do { traceTcS "canEqLeafTyVarLeft" (pprEq (mkTyVarTy tv) s2)
; (xi2, co) <- flatten d fl s2 -- Flatten RHS co : xi2 ~ s2
; let no_flattening_happened = isTcReflCo co
; traceTcS "canEqLeafTyVarLeft" (nest 2 (vcat [ text "tv =" <+> ppr tv
, text "s2 =" <+> ppr s2
, text "xi2 =" <+> ppr xi2]))
-- Flattening the RHS may reveal an identity coercion, which should
-- not be reported as occurs check error!
; let is_same_tv
| Just tv' <- getTyVar_maybe xi2, tv' == tv
= True
| otherwise = False
; if is_same_tv then
do { delCachedEvVar eqv fl
; when (isWanted fl) $
do { _ <- setEqBind eqv co fl; return () }
; return Stop }
else
do { -- Do an occurs check, and return a possibly
-- unfolded version of the RHS, if we had to
-- unfold any type synonyms to get rid of tv.
occ_check_result <- canOccursCheck fl tv xi2
; let xi2'
| Just xi2_unfolded <- occ_check_result
= xi2_unfolded
| otherwise = xi2
; if no_flattening_happened then
if isNothing occ_check_result then
canEqFailure d fl (setVarType eqv $ mkEqPred (mkTyVarTy tv, xi2'))
else
continueWith $ CTyEqCan { cc_id = eqv
, cc_flavor = fl
, cc_tyvar = tv
, cc_rhs = xi2'
, cc_depth = d }
else -- Flattening happened, in any case we have to create new variable
-- even if we report an occurs check error
do { delCachedEvVar eqv fl
; evc <- newEqVar fl (mkTyVarTy tv) xi2'
; let eqv' = evc_the_evvar evc -- eqv' : tv ~ xi2'
cv = mkTcCoVarCo eqv -- cv : tv ~ s2
cv' = mkTcCoVarCo eqv' -- cv': tv ~ xi2'
; fl' <- case fl of
Wanted {} -> setEqBind eqv (cv' `mkTcTransCo` co) fl -- tv ~ xi2' ~ s2
Given {} -> setEqBind eqv' (cv `mkTcTransCo` mkTcSymCo co) fl -- tv ~ s2 ~ xi2'
Derived {} -> return fl
; if isNewEvVar evc then
if isNothing occ_check_result then
canEqFailure d fl eqv'
else continueWith CTyEqCan { cc_id = eqv'
, cc_flavor = fl'
, cc_tyvar = tv
, cc_rhs = xi2'
, cc_depth = d }
else
return Stop } } }
-- See Note [Type synonyms and canonicalization].
-- Check whether the given variable occurs in the given type. We may
-- have needed to do some type synonym unfolding in order to get rid
-- of the variable, so we also return the unfolded version of the
-- type, which is guaranteed to be syntactically free of the given
-- type variable. If the type is already syntactically free of the
-- variable, then the same type is returned.
--
-- Precondition: the two types are not equal (looking though synonyms)
canOccursCheck :: CtFlavor -> TcTyVar -> Xi -> TcS (Maybe Xi)
canOccursCheck _gw tv xi = return (expandAway tv xi)
\end{code}
@expandAway tv xi@ expands synonyms in xi just enough to get rid of
occurrences of tv, if that is possible; otherwise, it returns Nothing.
For example, suppose we have
type F a b = [a]
Then
expandAway b (F Int b) = Just [Int]
but
expandAway a (F a Int) = Nothing
We don't promise to do the absolute minimum amount of expanding
necessary, but we try not to do expansions we don't need to. We
prefer doing inner expansions first. For example,
type F a b = (a, Int, a, [a])
type G b = Char
We have
expandAway b (F (G b)) = F Char
even though we could also expand F to get rid of b.
\begin{code}
expandAway :: TcTyVar -> Xi -> Maybe Xi
expandAway tv t@(TyVarTy tv')
| tv == tv' = Nothing
| otherwise = Just t
expandAway tv xi
| not (tv `elemVarSet` tyVarsOfType xi) = Just xi
expandAway tv (AppTy ty1 ty2)
= do { ty1' <- expandAway tv ty1
; ty2' <- expandAway tv ty2
; return (mkAppTy ty1' ty2') }
-- mkAppTy <$> expandAway tv ty1 <*> expandAway tv ty2
expandAway tv (FunTy ty1 ty2)
= do { ty1' <- expandAway tv ty1
; ty2' <- expandAway tv ty2
; return (mkFunTy ty1' ty2') }
-- mkFunTy <$> expandAway tv ty1 <*> expandAway tv ty2
expandAway tv ty@(ForAllTy {})
= let (tvs,rho) = splitForAllTys ty
tvs_knds = map tyVarKind tvs
in if tv `elemVarSet` tyVarsOfTypes tvs_knds then
-- Can't expand away the kinds unless we create
-- fresh variables which we don't want to do at this point.
Nothing
else do { rho' <- expandAway tv rho
; return (mkForAllTys tvs rho') }
-- For a type constructor application, first try expanding away the
-- offending variable from the arguments. If that doesn't work, next
-- see if the type constructor is a type synonym, and if so, expand
-- it and try again.
expandAway tv ty@(TyConApp tc tys)
= (mkTyConApp tc <$> mapM (expandAway tv) tys) <|> (tcView ty >>= expandAway tv)
\end{code}
Note [Type synonyms and canonicalization]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
We treat type synonym applications as xi types, that is, they do not
count as type function applications. However, we do need to be a bit
careful with type synonyms: like type functions they may not be
generative or injective. However, unlike type functions, they are
parametric, so there is no problem in expanding them whenever we see
them, since we do not need to know anything about their arguments in
order to expand them; this is what justifies not having to treat them
as specially as type function applications. The thing that causes
some subtleties is that we prefer to leave type synonym applications
*unexpanded* whenever possible, in order to generate better error
messages.
If we encounter an equality constraint with type synonym applications
on both sides, or a type synonym application on one side and some sort
of type application on the other, we simply must expand out the type
synonyms in order to continue decomposing the equality constraint into
primitive equality constraints. For example, suppose we have
type F a = [Int]
and we encounter the equality
F a ~ [b]
In order to continue we must expand F a into [Int], giving us the
equality
[Int] ~ [b]
which we can then decompose into the more primitive equality
constraint
Int ~ b.
However, if we encounter an equality constraint with a type synonym
application on one side and a variable on the other side, we should
NOT (necessarily) expand the type synonym, since for the purpose of
good error messages we want to leave type synonyms unexpanded as much
as possible.
However, there is a subtle point with type synonyms and the occurs
check that takes place for equality constraints of the form tv ~ xi.
As an example, suppose we have
type F a = Int
and we come across the equality constraint
a ~ F a
This should not actually fail the occurs check, since expanding out
the type synonym results in the legitimate equality constraint a ~
Int. We must actually do this expansion, because unifying a with F a
will lead the type checker into infinite loops later. Put another
way, canonical equality constraints should never *syntactically*
contain the LHS variable in the RHS type. However, we don't always
need to expand type synonyms when doing an occurs check; for example,
the constraint
a ~ F b
is obviously fine no matter what F expands to. And in this case we
would rather unify a with F b (rather than F b's expansion) in order
to get better error messages later.
So, when doing an occurs check with a type synonym application on the
RHS, we use some heuristics to find an expansion of the RHS which does
not contain the variable from the LHS. In particular, given
a ~ F t1 ... tn
we first try expanding each of the ti to types which no longer contain
a. If this turns out to be impossible, we next try expanding F
itself, and so on.
%************************************************************************
%* *
%* Functional dependencies, instantiation of equations
%* *
%************************************************************************
When we spot an equality arising from a functional dependency,
we now use that equality (a "wanted") to rewrite the work-item
constraint right away. This avoids two dangers
Danger 1: If we send the original constraint on down the pipeline
it may react with an instance declaration, and in delicate
situations (when a Given overlaps with an instance) that
may produce new insoluble goals: see Trac #4952
Danger 2: If we don't rewrite the constraint, it may re-react
with the same thing later, and produce the same equality
again --> termination worries.
To achieve this required some refactoring of FunDeps.lhs (nicer
now!).
\begin{code}
rewriteWithFunDeps :: [Equation]
-> [Xi]
-> WantedLoc
-> TcS (Maybe ([Xi], [TcCoercion], [(EvVar,WantedLoc)]))
-- Not quite a WantedEvVar unfortunately
-- Because our intention could be to make
-- it derived at the end of the day
-- NB: The flavor of the returned EvVars will be decided by the caller
-- Post: returns no trivial equalities (identities) and all EvVars returned are fresh
rewriteWithFunDeps eqn_pred_locs xis wloc
= do { fd_ev_poss <- mapM (instFunDepEqn wloc) eqn_pred_locs
; let fd_ev_pos :: [(Int,(EqVar,WantedLoc))]
fd_ev_pos = concat fd_ev_poss
(rewritten_xis, cos) = unzip (rewriteDictParams fd_ev_pos xis)
; if null fd_ev_pos then return Nothing
else return (Just (rewritten_xis, cos, map snd fd_ev_pos)) }
instFunDepEqn :: WantedLoc -> Equation -> TcS [(Int,(EvVar,WantedLoc))]
-- Post: Returns the position index as well as the corresponding FunDep equality
instFunDepEqn wl (FDEqn { fd_qtvs = qtvs, fd_eqs = eqs
, fd_pred1 = d1, fd_pred2 = d2 })
= do { let tvs = varSetElems qtvs
; tvs' <- mapM instFlexiTcS tvs -- IA0_TODO: we might need to do kind substitution
; let subst = zipTopTvSubst tvs (mkTyVarTys tvs')
; foldM (do_one subst) [] eqs }
where
do_one subst ievs (FDEq { fd_pos = i, fd_ty_left = ty1, fd_ty_right = ty2 })
= let sty1 = Type.substTy subst ty1
sty2 = Type.substTy subst ty2
in if eqType sty1 sty2 then return ievs -- Return no trivial equalities
else do { eqv <- newEqVar (Derived wl) sty1 sty2 -- Create derived or cached by deriveds
; let wl' = push_ctx wl
; if isNewEvVar eqv then
return $ (i,(evc_the_evvar eqv,wl')):ievs
else -- We are eventually going to emit FD work back in the work list so
-- it is important that we only return the /freshly created/ and not
-- some existing equality!
return ievs }
push_ctx :: WantedLoc -> WantedLoc
push_ctx loc = pushErrCtxt FunDepOrigin (False, mkEqnMsg d1 d2) loc
mkEqnMsg :: (TcPredType, SDoc)
-> (TcPredType, SDoc) -> TidyEnv -> TcM (TidyEnv, SDoc)
mkEqnMsg (pred1,from1) (pred2,from2) tidy_env
= do { zpred1 <- TcM.zonkTcPredType pred1
; zpred2 <- TcM.zonkTcPredType pred2
; let { tpred1 = tidyType tidy_env zpred1
; tpred2 = tidyType tidy_env zpred2 }
; let msg = vcat [ptext (sLit "When using functional dependencies to combine"),
nest 2 (sep [ppr tpred1 <> comma, nest 2 from1]),
nest 2 (sep [ppr tpred2 <> comma, nest 2 from2])]
; return (tidy_env, msg) }
rewriteDictParams :: [(Int,(EqVar,WantedLoc))] -- A set of coercions : (pos, ty' ~ ty)
-> [Type] -- A sequence of types: tys
-> [(Type, TcCoercion)] -- Returns: [(ty', co : ty' ~ ty)]
rewriteDictParams param_eqs tys
= zipWith do_one tys [0..]
where
do_one :: Type -> Int -> (Type, TcCoercion)
do_one ty n = case lookup n param_eqs of
Just wev -> (get_fst_ty wev, mkTcCoVarCo (fst wev))
Nothing -> (ty, mkTcReflCo ty) -- Identity
get_fst_ty (wev,_wloc)
| Just (ty1, _) <- getEqPredTys_maybe (evVarPred wev )
= ty1
| otherwise
= panic "rewriteDictParams: non equality fundep!?"
emitFDWork :: Bool
-> [(EvVar,WantedLoc)]
-> SubGoalDepth -> TcS ()
emitFDWork as_wanted evlocs d
= updWorkListTcS $ appendWorkListEqs fd_cts
where fd_cts = map mk_fd_ct evlocs
mk_fl wl = if as_wanted then (Wanted wl) else (Derived wl)
mk_fd_ct (v,wl) = CNonCanonical { cc_id = v
, cc_flavor = mk_fl wl
, cc_depth = d }
emitFDWorkAsDerived, emitFDWorkAsWanted :: [(EvVar,WantedLoc)]
-> SubGoalDepth
-> TcS ()
emitFDWorkAsDerived = emitFDWork False
emitFDWorkAsWanted = emitFDWork True
\end{code}