{-# LANGUAGE CPP, DeriveFunctor, TypeFamilies, ScopedTypeVariables, TypeApplications,
DerivingStrategies, GeneralizedNewtypeDeriving, ScopedTypeVariables, MultiWayIf, ViewPatterns #-}
{-# OPTIONS_GHC -Wno-incomplete-record-updates -Wno-incomplete-uni-patterns #-}
-- | Type definitions for the constraint solver
module GHC.Tc.Solver.Monad (
-- The work list
WorkList(..), isEmptyWorkList, emptyWorkList,
extendWorkListNonEq, extendWorkListCt,
extendWorkListCts, extendWorkListEq,
appendWorkList,
selectNextWorkItem,
workListSize,
getWorkList, updWorkListTcS, pushLevelNoWorkList,
-- The TcS monad
TcS, runTcS, runTcSDeriveds, runTcSWithEvBinds, runTcSInerts,
failTcS, warnTcS, addErrTcS, wrapTcS,
runTcSEqualities,
nestTcS, nestImplicTcS, setEvBindsTcS,
emitImplicationTcS, emitTvImplicationTcS,
runTcPluginTcS, addUsedGRE, addUsedGREs, keepAlive,
matchGlobalInst, TcM.ClsInstResult(..),
QCInst(..),
-- Tracing etc
panicTcS, traceTcS,
traceFireTcS, bumpStepCountTcS, csTraceTcS,
wrapErrTcS, wrapWarnTcS,
resetUnificationFlag, setUnificationFlag,
-- Evidence creation and transformation
MaybeNew(..), freshGoals, isFresh, getEvExpr,
newTcEvBinds, newNoTcEvBinds,
newWantedEq, newWantedEq_SI, emitNewWantedEq,
newWanted, newWanted_SI, newWantedEvVar,
newWantedNC, newWantedEvVarNC,
newDerivedNC,
newBoundEvVarId,
unifyTyVar, reportUnifications, touchabilityTest, TouchabilityTestResult(..),
setEvBind, setWantedEq,
setWantedEvTerm, setEvBindIfWanted,
newEvVar, newGivenEvVar, newGivenEvVars,
emitNewDeriveds, emitNewDerivedEq,
checkReductionDepth,
getSolvedDicts, setSolvedDicts,
getInstEnvs, getFamInstEnvs, -- Getting the environments
getTopEnv, getGblEnv, getLclEnv,
getTcEvBindsVar, getTcLevel,
getTcEvTyCoVars, getTcEvBindsMap, setTcEvBindsMap,
tcLookupClass, tcLookupId,
-- Inerts
InertSet(..), InertCans(..), emptyInert,
updInertTcS, updInertCans, updInertDicts, updInertIrreds,
getHasGivenEqs, setInertCans,
getInertEqs, getInertCans, getInertGivens,
getInertInsols, getInnermostGivenEqLevel,
getTcSInerts, setTcSInerts,
matchableGivens, prohibitedSuperClassSolve, mightEqualLater,
getUnsolvedInerts,
removeInertCts, getPendingGivenScs,
addInertCan, insertFunEq, addInertForAll,
emitWorkNC, emitWork,
isImprovable,
-- The Model
kickOutAfterUnification,
-- Inert Safe Haskell safe-overlap failures
addInertSafehask, insertSafeOverlapFailureTcS, updInertSafehask,
getSafeOverlapFailures,
-- Inert CDictCans
DictMap, emptyDictMap, lookupInertDict, findDictsByClass, addDict,
addDictsByClass, delDict, foldDicts, filterDicts, findDict,
-- Inert CEqCans
EqualCtList(..), findTyEqs, foldTyEqs,
findEq,
-- Inert solved dictionaries
addSolvedDict, lookupSolvedDict,
-- Irreds
foldIrreds,
-- The family application cache
lookupFamAppInert, lookupFamAppCache, extendFamAppCache,
pprKicked,
-- Inert function equalities
findFunEq, findFunEqsByTyCon,
instDFunType, -- Instantiation
-- MetaTyVars
newFlexiTcSTy, instFlexi, instFlexiX,
cloneMetaTyVar,
tcInstSkolTyVarsX,
TcLevel,
isFilledMetaTyVar_maybe, isFilledMetaTyVar,
zonkTyCoVarsAndFV, zonkTcType, zonkTcTypes, zonkTcTyVar, zonkCo,
zonkTyCoVarsAndFVList,
zonkSimples, zonkWC,
zonkTyCoVarKind,
-- References
newTcRef, readTcRef, writeTcRef, updTcRef,
-- Misc
getDefaultInfo, getDynFlags, getGlobalRdrEnvTcS,
matchFam, matchFamTcM,
checkWellStagedDFun,
pprEq, -- Smaller utils, re-exported from TcM
-- TODO (DV): these are only really used in the
-- instance matcher in GHC.Tc.Solver. I am wondering
-- if the whole instance matcher simply belongs
-- here
breakTyVarCycle_maybe, rewriterView
) where
#include "HsVersions.h"
import GHC.Prelude
import GHC.Driver.Env
import qualified GHC.Tc.Utils.Instantiate as TcM
import GHC.Core.InstEnv
import GHC.Tc.Instance.Family as FamInst
import GHC.Core.FamInstEnv
import qualified GHC.Tc.Utils.Monad as TcM
import qualified GHC.Tc.Utils.TcMType as TcM
import qualified GHC.Tc.Instance.Class as TcM( matchGlobalInst, ClsInstResult(..) )
import qualified GHC.Tc.Utils.Env as TcM
( checkWellStaged, tcGetDefaultTys, tcLookupClass, tcLookupId, topIdLvl )
import GHC.Tc.Instance.Class( InstanceWhat(..), safeOverlap, instanceReturnsDictCon )
import GHC.Tc.Utils.TcType
import GHC.Tc.Utils.Unify ( canSolveByUnification )
import GHC.Driver.Session
import GHC.Core.Type
import qualified GHC.Core.TyCo.Rep as Rep -- this needs to be used only very locally
import GHC.Core.Coercion
import GHC.Core.Unify
import GHC.Tc.Types.Evidence
import GHC.Core.Class
import GHC.Core.TyCon
import GHC.Tc.Errors ( solverDepthErrorTcS )
import GHC.Types.Name
import GHC.Types.TyThing
import GHC.Unit.Module ( HasModule, getModule )
import GHC.Types.Name.Reader ( GlobalRdrEnv, GlobalRdrElt )
import qualified GHC.Rename.Env as TcM
import GHC.Types.Var
import GHC.Types.Var.Env
import GHC.Types.Var.Set
import GHC.Utils.Outputable
import GHC.Utils.Panic
import GHC.Utils.Logger
import GHC.Data.Bag as Bag
import GHC.Types.Unique.Supply
import GHC.Utils.Misc
import GHC.Tc.Types
import GHC.Tc.Types.Origin
import GHC.Tc.Types.Constraint
import GHC.Core.Predicate
import GHC.Types.Unique.Set
import GHC.Core.TyCon.Env
import GHC.Data.Maybe
import GHC.Core.Map.Type
import GHC.Data.TrieMap
import Control.Monad
import GHC.Utils.Monad
import Data.IORef
import GHC.Exts (oneShot)
import Data.List ( partition, mapAccumL )
import Data.List.NonEmpty ( NonEmpty(..), cons, toList, nonEmpty )
import qualified Data.List.NonEmpty as NE
import Control.Arrow ( first )
#if defined(DEBUG)
import GHC.Data.Graph.Directed
#endif
{-
************************************************************************
* *
* Worklists *
* Canonical and non-canonical constraints that the simplifier has to *
* work on. Including their simplification depths. *
* *
* *
************************************************************************
Note [WorkList priorities]
~~~~~~~~~~~~~~~~~~~~~~~~~~~
A WorkList contains canonical and non-canonical items (of all flavours).
Notice that each Ct now has a simplification depth. We may
consider using this depth for prioritization as well in the future.
As a simple form of priority queue, our worklist separates out
* equalities (wl_eqs); see Note [Prioritise equalities]
* all the rest (wl_rest)
Note [Prioritise equalities]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~
It's very important to process equalities /first/:
* (Efficiency) The general reason to do so is that if we process a
class constraint first, we may end up putting it into the inert set
and then kicking it out later. That's extra work compared to just
doing the equality first.
* (Avoiding fundep iteration) As #14723 showed, it's possible to
get non-termination if we
- Emit the Derived fundep equalities for a class constraint,
generating some fresh unification variables.
- That leads to some unification
- Which kicks out the class constraint
- Which isn't solved (because there are still some more Derived
equalities in the work-list), but generates yet more fundeps
Solution: prioritise derived equalities over class constraints
* (Class equalities) We need to prioritise equalities even if they
are hidden inside a class constraint;
see Note [Prioritise class equalities]
* (Kick-out) We want to apply this priority scheme to kicked-out
constraints too (see the call to extendWorkListCt in kick_out_rewritable
E.g. a CIrredCan can be a hetero-kinded (t1 ~ t2), which may become
homo-kinded when kicked out, and hence we want to prioritise it.
* (Derived equalities) Originally we tried to postpone processing
Derived equalities, in the hope that we might never need to deal
with them at all; but in fact we must process Derived equalities
eagerly, partly for the (Efficiency) reason, and more importantly
for (Avoiding fundep iteration).
Note [Prioritise class equalities]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
We prioritise equalities in the solver (see selectWorkItem). But class
constraints like (a ~ b) and (a ~~ b) are actually equalities too;
see Note [The equality types story] in GHC.Builtin.Types.Prim.
Failing to prioritise these is inefficient (more kick-outs etc).
But, worse, it can prevent us spotting a "recursive knot" among
Wanted constraints. See comment:10 of #12734 for a worked-out
example.
So we arrange to put these particular class constraints in the wl_eqs.
NB: since we do not currently apply the substitution to the
inert_solved_dicts, the knot-tying still seems a bit fragile.
But this makes it better.
-}
-- See Note [WorkList priorities]
data WorkList
= WL { wl_eqs :: [Ct] -- CEqCan, CDictCan, CIrredCan
-- Given, Wanted, and Derived
-- Contains both equality constraints and their
-- class-level variants (a~b) and (a~~b);
-- See Note [Prioritise equalities]
-- See Note [Prioritise class equalities]
, wl_rest :: [Ct]
, wl_implics :: Bag Implication -- See Note [Residual implications]
}
appendWorkList :: WorkList -> WorkList -> WorkList
appendWorkList
(WL { wl_eqs = eqs1, wl_rest = rest1
, wl_implics = implics1 })
(WL { wl_eqs = eqs2, wl_rest = rest2
, wl_implics = implics2 })
= WL { wl_eqs = eqs1 ++ eqs2
, wl_rest = rest1 ++ rest2
, wl_implics = implics1 `unionBags` implics2 }
workListSize :: WorkList -> Int
workListSize (WL { wl_eqs = eqs, wl_rest = rest })
= length eqs + length rest
extendWorkListEq :: Ct -> WorkList -> WorkList
extendWorkListEq ct wl = wl { wl_eqs = ct : wl_eqs wl }
extendWorkListNonEq :: Ct -> WorkList -> WorkList
-- Extension by non equality
extendWorkListNonEq ct wl = wl { wl_rest = ct : wl_rest wl }
extendWorkListDeriveds :: [CtEvidence] -> WorkList -> WorkList
extendWorkListDeriveds evs wl
= extendWorkListCts (map mkNonCanonical evs) wl
extendWorkListImplic :: Implication -> WorkList -> WorkList
extendWorkListImplic implic wl = wl { wl_implics = implic `consBag` wl_implics wl }
extendWorkListCt :: Ct -> WorkList -> WorkList
-- Agnostic
extendWorkListCt ct wl
= case classifyPredType (ctPred ct) of
EqPred {}
-> extendWorkListEq ct wl
ClassPred cls _ -- See Note [Prioritise class equalities]
| isEqPredClass cls
-> extendWorkListEq ct wl
_ -> extendWorkListNonEq ct wl
extendWorkListCts :: [Ct] -> WorkList -> WorkList
-- Agnostic
extendWorkListCts cts wl = foldr extendWorkListCt wl cts
isEmptyWorkList :: WorkList -> Bool
isEmptyWorkList (WL { wl_eqs = eqs, wl_rest = rest, wl_implics = implics })
= null eqs && null rest && isEmptyBag implics
emptyWorkList :: WorkList
emptyWorkList = WL { wl_eqs = [], wl_rest = [], wl_implics = emptyBag }
selectWorkItem :: WorkList -> Maybe (Ct, WorkList)
-- See Note [Prioritise equalities]
selectWorkItem wl@(WL { wl_eqs = eqs, wl_rest = rest })
| ct:cts <- eqs = Just (ct, wl { wl_eqs = cts })
| ct:cts <- rest = Just (ct, wl { wl_rest = cts })
| otherwise = Nothing
getWorkList :: TcS WorkList
getWorkList = do { wl_var <- getTcSWorkListRef
; wrapTcS (TcM.readTcRef wl_var) }
selectNextWorkItem :: TcS (Maybe Ct)
-- Pick which work item to do next
-- See Note [Prioritise equalities]
selectNextWorkItem
= do { wl_var <- getTcSWorkListRef
; wl <- readTcRef wl_var
; case selectWorkItem wl of {
Nothing -> return Nothing ;
Just (ct, new_wl) ->
do { -- checkReductionDepth (ctLoc ct) (ctPred ct)
-- This is done by GHC.Tc.Solver.Interact.chooseInstance
; writeTcRef wl_var new_wl
; return (Just ct) } } }
-- Pretty printing
instance Outputable WorkList where
ppr (WL { wl_eqs = eqs, wl_rest = rest, wl_implics = implics })
= text "WL" <+> (braces $
vcat [ ppUnless (null eqs) $
text "Eqs =" <+> vcat (map ppr eqs)
, ppUnless (null rest) $
text "Non-eqs =" <+> vcat (map ppr rest)
, ppUnless (isEmptyBag implics) $
ifPprDebug (text "Implics =" <+> vcat (map ppr (bagToList implics)))
(text "(Implics omitted)")
])
{- *********************************************************************
* *
InertSet: the inert set
* *
* *
********************************************************************* -}
data InertSet
= IS { inert_cans :: InertCans
-- Canonical Given, Wanted, Derived
-- Sometimes called "the inert set"
, inert_cycle_breakers :: [(TcTyVar, TcType)]
-- a list of CycleBreakerTv / original family applications
-- used to undo the cycle-breaking needed to handle
-- Note [Type variable cycles in Givens] in GHC.Tc.Solver.Canonical
, inert_famapp_cache :: FunEqMap (TcCoercion, TcType)
-- Just a hash-cons cache for use when reducing family applications
-- only
--
-- If F tys :-> (co, rhs, flav),
-- then co :: rhs ~N F tys
-- all evidence is from instances or Givens; no coercion holes here
-- (We have no way of "kicking out" from the cache, so putting
-- wanteds here means we can end up solving a Wanted with itself. Bad)
, inert_solved_dicts :: DictMap CtEvidence
-- All Wanteds, of form ev :: C t1 .. tn
-- See Note [Solved dictionaries]
-- and Note [Do not add superclasses of solved dictionaries]
}
instance Outputable InertSet where
ppr (IS { inert_cans = ics
, inert_solved_dicts = solved_dicts })
= vcat [ ppr ics
, ppUnless (null dicts) $
text "Solved dicts =" <+> vcat (map ppr dicts) ]
where
dicts = bagToList (dictsToBag solved_dicts)
emptyInertCans :: InertCans
emptyInertCans
= IC { inert_eqs = emptyDVarEnv
, inert_given_eq_lvl = topTcLevel
, inert_given_eqs = False
, inert_dicts = emptyDicts
, inert_safehask = emptyDicts
, inert_funeqs = emptyFunEqs
, inert_insts = []
, inert_irreds = emptyCts
, inert_blocked = emptyCts }
emptyInert :: InertSet
emptyInert
= IS { inert_cans = emptyInertCans
, inert_cycle_breakers = []
, inert_famapp_cache = emptyFunEqs
, inert_solved_dicts = emptyDictMap }
{- Note [Solved dictionaries]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
When we apply a top-level instance declaration, we add the "solved"
dictionary to the inert_solved_dicts. In general, we use it to avoid
creating a new EvVar when we have a new goal that we have solved in
the past.
But in particular, we can use it to create *recursive* dictionaries.
The simplest, degenerate case is
instance C [a] => C [a] where ...
If we have
[W] d1 :: C [x]
then we can apply the instance to get
d1 = $dfCList d
[W] d2 :: C [x]
Now 'd1' goes in inert_solved_dicts, and we can solve d2 directly from d1.
d1 = $dfCList d
d2 = d1
See Note [Example of recursive dictionaries]
VERY IMPORTANT INVARIANT:
(Solved Dictionary Invariant)
Every member of the inert_solved_dicts is the result
of applying an instance declaration that "takes a step"
An instance "takes a step" if it has the form
dfunDList d1 d2 = MkD (...) (...) (...)
That is, the dfun is lazy in its arguments, and guarantees to
immediately return a dictionary constructor. NB: all dictionary
data constructors are lazy in their arguments.
This property is crucial to ensure that all dictionaries are
non-bottom, which in turn ensures that the whole "recursive
dictionary" idea works at all, even if we get something like
rec { d = dfunDList d dx }
See Note [Recursive superclasses] in GHC.Tc.TyCl.Instance.
Reason:
- All instances, except two exceptions listed below, "take a step"
in the above sense
- Exception 1: local quantified constraints have no such guarantee;
indeed, adding a "solved dictionary" when appling a quantified
constraint led to the ability to define unsafeCoerce
in #17267.
- Exception 2: the magic built-in instance for (~) has no
such guarantee. It behaves as if we had
class (a ~# b) => (a ~ b) where {}
instance (a ~# b) => (a ~ b) where {}
The "dfun" for the instance is strict in the coercion.
Anyway there's no point in recording a "solved dict" for
(t1 ~ t2); it's not going to allow a recursive dictionary
to be constructed. Ditto (~~) and Coercible.
THEREFORE we only add a "solved dictionary"
- when applying an instance declaration
- subject to Exceptions 1 and 2 above
In implementation terms
- GHC.Tc.Solver.Monad.addSolvedDict adds a new solved dictionary,
conditional on the kind of instance
- It is only called when applying an instance decl,
in GHC.Tc.Solver.Interact.doTopReactDict
- ClsInst.InstanceWhat says what kind of instance was
used to solve the constraint. In particular
* LocalInstance identifies quantified constraints
* BuiltinEqInstance identifies the strange built-in
instances for equality.
- ClsInst.instanceReturnsDictCon says which kind of
instance guarantees to return a dictionary constructor
Other notes about solved dictionaries
* See also Note [Do not add superclasses of solved dictionaries]
* The inert_solved_dicts field is not rewritten by equalities,
so it may get out of date.
* The inert_solved_dicts are all Wanteds, never givens
* We only cache dictionaries from top-level instances, not from
local quantified constraints. Reason: if we cached the latter
we'd need to purge the cache when bringing new quantified
constraints into scope, because quantified constraints "shadow"
top-level instances.
Note [Do not add superclasses of solved dictionaries]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Every member of inert_solved_dicts is the result of applying a
dictionary function, NOT of applying superclass selection to anything.
Consider
class Ord a => C a where
instance Ord [a] => C [a] where ...
Suppose we are trying to solve
[G] d1 : Ord a
[W] d2 : C [a]
Then we'll use the instance decl to give
[G] d1 : Ord a Solved: d2 : C [a] = $dfCList d3
[W] d3 : Ord [a]
We must not add d4 : Ord [a] to the 'solved' set (by taking the
superclass of d2), otherwise we'll use it to solve d3, without ever
using d1, which would be a catastrophe.
Solution: when extending the solved dictionaries, do not add superclasses.
That's why each element of the inert_solved_dicts is the result of applying
a dictionary function.
Note [Example of recursive dictionaries]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
--- Example 1
data D r = ZeroD | SuccD (r (D r));
instance (Eq (r (D r))) => Eq (D r) where
ZeroD == ZeroD = True
(SuccD a) == (SuccD b) = a == b
_ == _ = False;
equalDC :: D [] -> D [] -> Bool;
equalDC = (==);
We need to prove (Eq (D [])). Here's how we go:
[W] d1 : Eq (D [])
By instance decl of Eq (D r):
[W] d2 : Eq [D []] where d1 = dfEqD d2
By instance decl of Eq [a]:
[W] d3 : Eq (D []) where d2 = dfEqList d3
d1 = dfEqD d2
Now this wanted can interact with our "solved" d1 to get:
d3 = d1
-- Example 2:
This code arises in the context of "Scrap Your Boilerplate with Class"
class Sat a
class Data ctx a
instance Sat (ctx Char) => Data ctx Char -- dfunData1
instance (Sat (ctx [a]), Data ctx a) => Data ctx [a] -- dfunData2
class Data Maybe a => Foo a
instance Foo t => Sat (Maybe t) -- dfunSat
instance Data Maybe a => Foo a -- dfunFoo1
instance Foo a => Foo [a] -- dfunFoo2
instance Foo [Char] -- dfunFoo3
Consider generating the superclasses of the instance declaration
instance Foo a => Foo [a]
So our problem is this
[G] d0 : Foo t
[W] d1 : Data Maybe [t] -- Desired superclass
We may add the given in the inert set, along with its superclasses
Inert:
[G] d0 : Foo t
[G] d01 : Data Maybe t -- Superclass of d0
WorkList
[W] d1 : Data Maybe [t]
Solve d1 using instance dfunData2; d1 := dfunData2 d2 d3
Inert:
[G] d0 : Foo t
[G] d01 : Data Maybe t -- Superclass of d0
Solved:
d1 : Data Maybe [t]
WorkList:
[W] d2 : Sat (Maybe [t])
[W] d3 : Data Maybe t
Now, we may simplify d2 using dfunSat; d2 := dfunSat d4
Inert:
[G] d0 : Foo t
[G] d01 : Data Maybe t -- Superclass of d0
Solved:
d1 : Data Maybe [t]
d2 : Sat (Maybe [t])
WorkList:
[W] d3 : Data Maybe t
[W] d4 : Foo [t]
Now, we can just solve d3 from d01; d3 := d01
Inert
[G] d0 : Foo t
[G] d01 : Data Maybe t -- Superclass of d0
Solved:
d1 : Data Maybe [t]
d2 : Sat (Maybe [t])
WorkList
[W] d4 : Foo [t]
Now, solve d4 using dfunFoo2; d4 := dfunFoo2 d5
Inert
[G] d0 : Foo t
[G] d01 : Data Maybe t -- Superclass of d0
Solved:
d1 : Data Maybe [t]
d2 : Sat (Maybe [t])
d4 : Foo [t]
WorkList:
[W] d5 : Foo t
Now, d5 can be solved! d5 := d0
Result
d1 := dfunData2 d2 d3
d2 := dfunSat d4
d3 := d01
d4 := dfunFoo2 d5
d5 := d0
-}
{- *********************************************************************
* *
InertCans: the canonical inerts
* *
* *
********************************************************************* -}
data InertCans -- See Note [Detailed InertCans Invariants] for more
= IC { inert_eqs :: InertEqs
-- See Note [inert_eqs: the inert equalities]
-- All CEqCans with a TyVarLHS; index is the LHS tyvar
-- Domain = skolems and untouchables; a touchable would be unified
, inert_funeqs :: FunEqMap EqualCtList
-- All CEqCans with a TyFamLHS; index is the whole family head type.
-- LHS is fully rewritten (modulo eqCanRewrite constraints)
-- wrt inert_eqs
-- Can include all flavours, [G], [W], [WD], [D]
, inert_dicts :: DictMap Ct
-- Dictionaries only
-- All fully rewritten (modulo flavour constraints)
-- wrt inert_eqs
, inert_insts :: [QCInst]
, inert_safehask :: DictMap Ct
-- Failed dictionary resolution due to Safe Haskell overlapping
-- instances restriction. We keep this separate from inert_dicts
-- as it doesn't cause compilation failure, just safe inference
-- failure.
--
-- ^ See Note [Safe Haskell Overlapping Instances Implementation]
-- in "GHC.Tc.Solver"
, inert_irreds :: Cts
-- Irreducible predicates that cannot be made canonical,
-- and which don't interact with others (e.g. (c a))
-- and insoluble predicates (e.g. Int ~ Bool, or a ~ [a])
, inert_blocked :: Cts
-- Equality predicates blocked on a coercion hole.
-- Each Ct is a CIrredCan with cc_reason = HoleBlockerReason
-- See Note [Equalities with incompatible kinds] in GHC.Tc.Solver.Canonical
-- wrinkle (2)
-- These are stored separately from inert_irreds because
-- they get kicked out for different reasons
, inert_given_eq_lvl :: TcLevel
-- The TcLevel of the innermost implication that has a Given
-- equality of the sort that make a unification variable untouchable
-- (see Note [Unification preconditions] in GHC.Tc.Utils.Unify).
-- See Note [Tracking Given equalities] below
, inert_given_eqs :: Bool
-- True <=> The inert Givens *at this level* (tcl_tclvl)
-- could includes at least one equality /other than/ a
-- let-bound skolem equality.
-- Reason: report these givens when reporting a failed equality
-- See Note [Tracking Given equalities]
}
type InertEqs = DTyVarEnv EqualCtList
newtype EqualCtList = EqualCtList (NonEmpty Ct)
deriving newtype Outputable
-- See Note [EqualCtList invariants]
unitEqualCtList :: Ct -> EqualCtList
unitEqualCtList ct = EqualCtList (ct :| [])
addToEqualCtList :: Ct -> EqualCtList -> EqualCtList
-- NB: This function maintains the "derived-before-wanted" invariant of EqualCtList,
-- but not the others. See Note [EqualCtList invariants]
addToEqualCtList ct (EqualCtList old_eqs)
| isWantedCt ct
, eq1 :| eqs <- old_eqs
= EqualCtList (eq1 :| ct : eqs)
| otherwise
= EqualCtList (ct `cons` old_eqs)
filterEqualCtList :: (Ct -> Bool) -> EqualCtList -> Maybe EqualCtList
filterEqualCtList pred (EqualCtList cts)
= fmap EqualCtList (nonEmpty $ NE.filter pred cts)
equalCtListToList :: EqualCtList -> [Ct]
equalCtListToList (EqualCtList cts) = toList cts
listToEqualCtList :: [Ct] -> Maybe EqualCtList
-- NB: This does not maintain invariants other than having the EqualCtList be
-- non-empty
listToEqualCtList cts = EqualCtList <$> nonEmpty cts
{- Note [Tracking Given equalities]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
For reasons described in (UNTOUCHABLE) in GHC.Tc.Utils.Unify
Note [Unification preconditions], we can't unify
alpha[2] ~ Int
under a level-4 implication if there are any Given equalities
bound by the implications at level 3 of 4. To that end, the
InertCans tracks
inert_given_eq_lvl :: TcLevel
-- The TcLevel of the innermost implication that has a Given
-- equality of the sort that make a unification variable untouchable
-- (see Note [Unification preconditions] in GHC.Tc.Utils.Unify).
We update inert_given_eq_lvl whenever we add a Given to the
inert set, in updateGivenEqs.
Then a unification variable alpha[n] is untouchable iff
n < inert_given_eq_lvl
that is, if the unification variable was born outside an
enclosing Given equality.
Exactly which constraints should trigger (UNTOUCHABLE), and hence
should update inert_given_eq_lvl?
* We do /not/ need to worry about let-bound skolems, such ast
forall[2] a. a ~ [b] => blah
See Note [Let-bound skolems]
* Consider an implication
forall[2]. beta[1] => alpha[1] ~ Int
where beta is a unification variable that has already been unified
to () in an outer scope. Then alpha[1] is perfectly touchable and
we can unify alpha := Int. So when deciding whether the givens contain
an equality, we should canonicalise first, rather than just looking at
the /original/ givens (#8644).
* However, we must take account of *potential* equalities. Consider the
same example again, but this time we have /not/ yet unified beta:
forall[2] beta[1] => ...blah...
Because beta might turn into an equality, updateGivenEqs conservatively
treats it as a potential equality, and updates inert_give_eq_lvl
* What about something like forall[2] a b. a ~ F b => [W] alpha[1] ~ X y z?
That Given cannot affect the Wanted, because the Given is entirely
*local*: it mentions only skolems bound in the very same
implication. Such equalities need not make alpha untouchable. (Test
case typecheck/should_compile/LocalGivenEqs has a real-life
motivating example, with some detailed commentary.)
Hence the 'mentionsOuterVar' test in updateGivenEqs.
However, solely to support better error messages
(see Note [HasGivenEqs] in GHC.Tc.Types.Constraint) we also track
these "local" equalities in the boolean inert_given_eqs field.
This field is used only to set the ic_given_eqs field to LocalGivenEqs;
see the function getHasGivenEqs.
Here is a simpler case that triggers this behaviour:
data T where
MkT :: F a ~ G b => a -> b -> T
f (MkT _ _) = True
Because of this behaviour around local equality givens, we can infer the
type of f. This is typecheck/should_compile/LocalGivenEqs2.
* We need not look at the equality relation involved (nominal vs
representational), because representational equalities can still
imply nominal ones. For example, if (G a ~R G b) and G's argument's
role is nominal, then we can deduce a ~N b.
Note [Let-bound skolems]
~~~~~~~~~~~~~~~~~~~~~~~~
If * the inert set contains a canonical Given CEqCan (a ~ ty)
and * 'a' is a skolem bound in this very implication,
then:
a) The Given is pretty much a let-binding, like
f :: (a ~ b->c) => a -> a
Here the equality constraint is like saying
let a = b->c in ...
It is not adding any new, local equality information,
and hence can be ignored by has_given_eqs
b) 'a' will have been completely substituted out in the inert set,
so we can safely discard it.
For an example, see #9211.
See also GHC.Tc.Utils.Unify Note [Deeper level on the left] for how we ensure
that the right variable is on the left of the equality when both are
tyvars.
You might wonder whether the skolem really needs to be bound "in the
very same implication" as the equuality constraint.
Consider this (c.f. #15009):
data S a where
MkS :: (a ~ Int) => S a
g :: forall a. S a -> a -> blah
g x y = let h = \z. ( z :: Int
, case x of
MkS -> [y,z])
in ...
From the type signature for `g`, we get `y::a` . Then when we
encounter the `\z`, we'll assign `z :: alpha[1]`, say. Next, from the
body of the lambda we'll get
[W] alpha[1] ~ Int -- From z::Int
[W] forall[2]. (a ~ Int) => [W] alpha[1] ~ a -- From [y,z]
Now, unify alpha := a. Now we are stuck with an unsolved alpha~Int!
So we must treat alpha as untouchable under the forall[2] implication.
Note [Detailed InertCans Invariants]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
The InertCans represents a collection of constraints with the following properties:
* All canonical
* No two dictionaries with the same head
* No two CIrreds with the same type
* Family equations inert wrt top-level family axioms
* Dictionaries have no matching top-level instance
* Given family or dictionary constraints don't mention touchable
unification variables
* Non-CEqCan constraints are fully rewritten with respect
to the CEqCan equalities (modulo eqCanRewrite of course;
eg a wanted cannot rewrite a given)
* CEqCan equalities: see Note [inert_eqs: the inert equalities]
Also see documentation in Constraint.Ct for a list of invariants
Note [EqualCtList invariants]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
* All are equalities
* All these equalities have the same LHS
* The list is never empty
* No element of the list can rewrite any other
* Derived before Wanted
From the fourth invariant it follows that the list is
- A single [G], or
- Zero or one [D] or [WD], followed by any number of [W]
The Wanteds can't rewrite anything which is why we put them last
Note [inert_eqs: the inert equalities]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Definition [Can-rewrite relation]
A "can-rewrite" relation between flavours, written f1 >= f2, is a
binary relation with the following properties
(R1) >= is transitive
(R2) If f1 >= f, and f2 >= f,
then either f1 >= f2 or f2 >= f1
(See Note [Why R2?].)
Lemma (L0). If f1 >= f then f1 >= f1
Proof. By property (R2), with f1=f2
Definition [Generalised substitution]
A "generalised substitution" S is a set of triples (lhs -f-> t), where
lhs is a type variable or an exactly-saturated type family application
(that is, lhs is a CanEqLHS)
t is a type
f is a flavour
such that
(WF1) if (lhs1 -f1-> t1) in S
(lhs2 -f2-> t2) in S
then (f1 >= f2) implies that lhs1 does not appear within lhs2
(WF2) if (lhs -f-> t) is in S, then t /= lhs
Definition [Applying a generalised substitution]
If S is a generalised substitution
S(f,t0) = t, if (t0 -fs-> t) in S, and fs >= f
= apply S to components of t0, otherwise
See also Note [Flavours with roles].
Theorem: S(f,t0) is well defined as a function.
Proof: Suppose (lhs -f1-> t1) and (lhs -f2-> t2) are both in S,
and f1 >= f and f2 >= f
Then by (R2) f1 >= f2 or f2 >= f1, which contradicts (WF1)
Notation: repeated application.
S^0(f,t) = t
S^(n+1)(f,t) = S(f, S^n(t))
Definition: terminating generalised substitution
A generalised substitution S is *terminating* iff
(IG1) there is an n such that
for every f,t, S^n(f,t) = S^(n+1)(f,t)
By (IG1) we define S*(f,t) to be the result of exahaustively
applying S(f,_) to t.
-----------------------------------------------------------------------------
Our main invariant:
the CEqCans in inert_eqs should be a terminating generalised substitution
-----------------------------------------------------------------------------
Note that termination is not the same as idempotence. To apply S to a
type, you may have to apply it recursively. But termination does
guarantee that this recursive use will terminate.
Note [Why R2?]
~~~~~~~~~~~~~~
R2 states that, if we have f1 >= f and f2 >= f, then either f1 >= f2 or f2 >=
f1. If we do not have R2, we will easily fall into a loop.
To see why, imagine we have f1 >= f, f2 >= f, and that's it. Then, let our
inert set S = {a -f1-> b, b -f2-> a}. Computing S(f,a) does not terminate. And
yet, we have a hard time noticing an occurs-check problem when building S, as
the two equalities cannot rewrite one another.
R2 actually restricts our ability to accept user-written programs. See Note
[Deriveds do rewrite Deriveds] in GHC.Tc.Types.Constraint for an example.
Note [Rewritable]
~~~~~~~~~~~~~~~~~
This Note defines what it means for a type variable or type family application
(that is, a CanEqLHS) to be rewritable in a type. This definition is used
by the anyRewritableXXX family of functions and is meant to model the actual
behaviour in GHC.Tc.Solver.Rewrite.
Ignoring roles (for now): A CanEqLHS lhs is *rewritable* in a type t if the
lhs tree appears as a subtree within t without traversing any of the following
components of t:
* coercions (whether they appear in casts CastTy or as arguments CoercionTy)
* kinds of variable occurrences
The check for rewritability *does* look in kinds of the bound variable of a
ForAllTy.
Goal: If lhs is not rewritable in t, then t is a fixpoint of the generalised
substitution containing only {lhs -f*-> t'}, where f* is a flavour such that f* >= f
for all f.
The reason for this definition is that the rewriter does not rewrite in coercions
or variables' kinds. In turn, the rewriter does not need to rewrite there because
those places are never used for controlling the behaviour of the solver: these
places are not used in matching instances or in decomposing equalities.
There is one exception to the claim that non-rewritable parts of the tree do
not affect the solver: we sometimes do an occurs-check to decide e.g. how to
orient an equality. (See the comments on
GHC.Tc.Solver.Canonical.canEqTyVarFunEq.) Accordingly, the presence of a
variable in a kind or coercion just might influence the solver. Here is an
example:
type family Const x y where
Const x y = x
AxConst :: forall x y. Const x y ~# x
alpha :: Const Type Nat
[W] alpha ~ Int |> (sym (AxConst Type alpha) ;;
AxConst Type alpha ;;
sym (AxConst Type Nat))
The cast is clearly ludicrous (it ties together a cast and its symmetric version),
but we can't quite rule it out. (See (EQ1) from
Note [Respecting definitional equality] in GHC.Core.TyCo.Rep to see why we need
the Const Type Nat bit.) And yet this cast will (quite rightly) prevent alpha
from unifying with the RHS. I (Richard E) don't have an example of where this
problem can arise from a Haskell program, but we don't have an air-tight argument
for why the definition of *rewritable* given here is correct.
Taking roles into account: we must consider a rewrite at a given role. That is,
a rewrite arises from some equality, and that equality has a role associated
with it. As we traverse a type, we track what role we are allowed to rewrite with.
For example, suppose we have an inert [G] b ~R# Int. Then b is rewritable in
Maybe b but not in F b, where F is a type function. This role-aware logic is
present in both the anyRewritableXXX functions and in the rewriter.
See also Note [anyRewritableTyVar must be role-aware] in GHC.Tc.Utils.TcType.
Note [Extending the inert equalities]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Main Theorem [Stability under extension]
Suppose we have a "work item"
lhs -fw-> t
and a terminating generalised substitution S,
THEN the extended substitution T = S+(lhs -fw-> t)
is a terminating generalised substitution
PROVIDED
(T1) S(fw,lhs) = lhs -- LHS of work-item is a fixpoint of S(fw,_)
(T2) S(fw,t) = t -- RHS of work-item is a fixpoint of S(fw,_)
(T3) lhs not in t -- No occurs check in the work item
-- If lhs is a type family application, we require only that
-- lhs is not *rewritable* in t. See Note [Rewritable] and
-- Note [CEqCan occurs check] in GHC.Tc.Types.Constraint.
AND, for every (lhs1 -fs-> s) in S:
(K0) not (fw >= fs)
Reason: suppose we kick out (lhs1 -fs-> s),
and add (lhs -fw-> t) to the inert set.
The latter can't rewrite the former,
so the kick-out achieved nothing
-- From here, we can assume fw >= fs
OR (K4) lhs1 is a tyvar AND fs >= fw
OR { (K1) lhs is not rewritable in lhs1. See Note [Rewritable].
Reason: if fw >= fs, WF1 says we can't have both
lhs0 -fw-> t and F lhs0 -fs-> s
AND (K2): guarantees termination of the new substitution
{ (K2a) not (fs >= fs)
OR (K2b) lhs not in s }
AND (K3) See Note [K3: completeness of solving]
{ (K3a) If the role of fs is nominal: s /= lhs
(K3b) If the role of fs is representational:
s is not of form (lhs t1 .. tn) } }
Conditions (T1-T3) are established by the canonicaliser
Conditions (K1-K3) are established by GHC.Tc.Solver.Monad.kickOutRewritable
The idea is that
* T1 and T2 are guaranteed by exhaustively rewriting the work-item
with S(fw,_).
* T3 is guaranteed by an occurs-check on the work item.
This is done during canonicalisation, in checkTypeEq; invariant
(TyEq:OC) of CEqCan. See also Note [CEqCan occurs check] in GHC.Tc.Types.Constraint.
* (K1-3) are the "kick-out" criteria. (As stated, they are really the
"keep" criteria.) If the current inert S contains a triple that does
not satisfy (K1-3), then we remove it from S by "kicking it out",
and re-processing it.
* Note that kicking out is a Bad Thing, because it means we have to
re-process a constraint. The less we kick out, the better.
TODO: Make sure that kicking out really *is* a Bad Thing. We've assumed
this but haven't done the empirical study to check.
* Assume we have G>=G, G>=W and that's all. Then, when performing
a unification we add a new given a -G-> ty. But doing so does NOT require
us to kick out an inert wanted that mentions a, because of (K2a). This
is a common case, hence good not to kick out. See also (K2a) below.
* Lemma (L2): if not (fw >= fw), then K0 holds and we kick out nothing
Proof: using Definition [Can-rewrite relation], fw can't rewrite anything
and so K0 holds. Intuitively, since fw can't rewrite anything (Lemma (L0)),
adding it cannot cause any loops
This is a common case, because Wanteds cannot rewrite Wanteds.
It's used to avoid even looking for constraint to kick out.
* Lemma (L1): The conditions of the Main Theorem imply that there is no
(lhs -fs-> t) in S, s.t. (fs >= fw).
Proof. Suppose the contrary (fs >= fw). Then because of (T1),
S(fw,lhs)=lhs. But since fs>=fw, S(fw,lhs) = t, hence t=lhs. But now we
have (lhs -fs-> lhs) in S, which contradicts (WF2).
* The extended substitution satisfies (WF1) and (WF2)
- (K1) plus (L1) guarantee that the extended substitution satisfies (WF1).
- (T3) guarantees (WF2).
* (K2) and (K4) are about termination. Intuitively, any infinite chain S^0(f,t),
S^1(f,t), S^2(f,t).... must pass through the new work item infinitely
often, since the substitution without the work item is terminating; and must
pass through at least one of the triples in S infinitely often.
- (K2a): if not(fs>=fs) then there is no f that fs can rewrite (fs>=f)
(this is Lemma (L0)), and hence this triple never plays a role in application S(f,t).
It is always safe to extend S with such a triple.
(NB: we could strengten K1) in this way too, but see K3.
- (K2b): if lhs not in s, we have no further opportunity to apply the
work item
- (K4): See Note [K4]
* Lemma (L3). Suppose we have f* such that, for all f, f* >= f. Then
if we are adding lhs -fw-> t (where T1, T2, and T3 hold), we will keep a -f*-> s.
Proof. K4 holds; thus, we keep.
Key lemma to make it watertight.
Under the conditions of the Main Theorem,
forall f st fw >= f, a is not in S^k(f,t), for any k
Also, consider roles more carefully. See Note [Flavours with roles]
Note [K4]
~~~~~~~~~
K4 is a "keep" condition of Note [Extending the inert equalities].
Here is the scenario:
* We are considering adding (lhs -fw-> t) to the inert set S.
* S already has (lhs1 -fs-> s).
* We know S(fw, lhs) = lhs, S(fw, t) = t, and lhs is not rewritable in t.
See Note [Rewritable]. These are (T1), (T2), and (T3).
* We further know fw >= fs. (If not, then we short-circuit via (K0).)
K4 says that we may keep lhs1 -fs-> s in S if:
lhs1 is a tyvar AND fs >= fw
Why K4 guarantees termination:
* If fs >= fw, we know a is not rewritable in t, because of (T2).
* We further know lhs /= a, because of (T1).
* Accordingly, a use of the new inert item lhs -fw-> t cannot create the conditions
for a use of a -fs-> s (precisely because t does not mention a), and hence,
the extended substitution (with lhs -fw-> t in it) is a terminating
generalised substitution.
Recall that the termination generalised substitution includes only mappings that
pass an occurs check. This is (T3). At one point, we worried that the
argument here would fail if s mentioned a, but (T3) rules out this possibility.
Put another way: the terminating generalised substitution considers only the inert_eqs,
not other parts of the inert set (such as the irreds).
Can we liberalise K4? No.
Why we cannot drop the (fs >= fw) condition:
* Suppose not (fs >= fw). It might be the case that t mentions a, and this
can cause a loop. Example:
Work: [G] b ~ a
Inert: [D] a ~ b
(where G >= G, G >= D, and D >= D)
If we don't kick out the inert, then we get a loop on e.g. [D] a ~ Int.
* Note that the above example is different if the inert is a Given G, because
(T1) won't hold.
Why we cannot drop the tyvar condition:
* Presume fs >= fw. Thus, F tys is not rewritable in t, because of (T2).
* Can the use of lhs -fw-> t create the conditions for a use of F tys -fs-> s?
Yes! This can happen if t appears within tys.
Here is an example:
Work: [G] a ~ Int
Inert: [G] F Int ~ F a
Now, if we have [W] F a ~ Bool, we will rewrite ad infinitum on the left-hand
side. The key reason why K2b works in the tyvar case is that tyvars are atomic:
if the right-hand side of an equality does not mention a variable a, then it
cannot allow an equality with an LHS of a to fire. This is not the case for
type family applications.
Bottom line: K4 can keep only inerts with tyvars on the left. Put differently,
K4 will never prevent an inert with a type family on the left from being kicked
out.
Consequence: We never kick out a Given/Nominal equality with a tyvar on the left.
This is Lemma (L3) of Note [Extending the inert equalities]. It is good because
it means we can effectively model the mutable filling of metavariables with
Given/Nominal equalities. That is: it should be the case that we could rewrite
our solver never to fill in a metavariable; instead, it would "solve" a wanted
like alpha ~ Int by turning it into a Given, allowing it to be used in rewriting.
We would want the solver to behave the same whether it uses metavariables or
Givens. And (L3) says that no Given/Nominals over tyvars are ever kicked out,
just like we never unfill a metavariable. Nice.
Getting this wrong (that is, allowing K4 to apply to situations with the type
family on the left) led to #19042. (At that point, K4 was known as K2b.)
Originally, this condition was part of K2, but #17672 suggests it should be
a top-level K condition.
Note [K3: completeness of solving]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
(K3) is not necessary for the extended substitution
to be terminating. In fact K1 could be made stronger by saying
... then (not (fw >= fs) or not (fs >= fs))
But it's not enough for S to be terminating; we also want completeness.
That is, we want to be able to solve all soluble wanted equalities.
Suppose we have
work-item b -G-> a
inert-item a -W-> b
Assuming (G >= W) but not (W >= W), this fulfills all the conditions,
so we could extend the inerts, thus:
inert-items b -G-> a
a -W-> b
But if we kicked-out the inert item, we'd get
work-item a -W-> b
inert-item b -G-> a
Then rewrite the work-item gives us (a -W-> a), which is soluble via Refl.
So we add one more clause to the kick-out criteria
Another way to understand (K3) is that we treat an inert item
a -f-> b
in the same way as
b -f-> a
So if we kick out one, we should kick out the other. The orientation
is somewhat accidental.
When considering roles, we also need the second clause (K3b). Consider
work-item c -G/N-> a
inert-item a -W/R-> b c
The work-item doesn't get rewritten by the inert, because (>=) doesn't hold.
But we don't kick out the inert item because not (W/R >= W/R). So we just
add the work item. But then, consider if we hit the following:
work-item b -G/N-> Id
inert-items a -W/R-> b c
c -G/N-> a
where
newtype Id x = Id x
For similar reasons, if we only had (K3a), we wouldn't kick the
representational inert out. And then, we'd miss solving the inert, which
now reduced to reflexivity.
The solution here is to kick out representational inerts whenever the
lhs of a work item is "exposed", where exposed means being at the
head of the top-level application chain (lhs t1 .. tn). See
is_can_eq_lhs_head. This is encoded in (K3b).
Beware: if we make this test succeed too often, we kick out too much,
and the solver might loop. Consider (#14363)
work item: [G] a ~R f b
inert item: [G] b ~R f a
In GHC 8.2 the completeness tests more aggressive, and kicked out
the inert item; but no rewriting happened and there was an infinite
loop. All we need is to have the tyvar at the head.
Note [Flavours with roles]
~~~~~~~~~~~~~~~~~~~~~~~~~~
The system described in Note [inert_eqs: the inert equalities]
discusses an abstract
set of flavours. In GHC, flavours have two components: the flavour proper,
taken from {Wanted, Derived, Given} and the equality relation (often called
role), taken from {NomEq, ReprEq}.
When substituting w.r.t. the inert set,
as described in Note [inert_eqs: the inert equalities],
we must be careful to respect all components of a flavour.
For example, if we have
inert set: a -G/R-> Int
b -G/R-> Bool
type role T nominal representational
and we wish to compute S(W/R, T a b), the correct answer is T a Bool, NOT
T Int Bool. The reason is that T's first parameter has a nominal role, and
thus rewriting a to Int in T a b is wrong. Indeed, this non-congruence of
substitution means that the proof in Note [The inert equalities] may need
to be revisited, but we don't think that the end conclusion is wrong.
-}
instance Outputable InertCans where
ppr (IC { inert_eqs = eqs
, inert_funeqs = funeqs
, inert_dicts = dicts
, inert_safehask = safehask
, inert_irreds = irreds
, inert_blocked = blocked
, inert_given_eq_lvl = ge_lvl
, inert_given_eqs = given_eqs
, inert_insts = insts })
= braces $ vcat
[ ppUnless (isEmptyDVarEnv eqs) $
text "Equalities:"
<+> pprCts (foldDVarEnv folder emptyCts eqs)
, ppUnless (isEmptyTcAppMap funeqs) $
text "Type-function equalities =" <+> pprCts (foldFunEqs folder funeqs emptyCts)
, ppUnless (isEmptyTcAppMap dicts) $
text "Dictionaries =" <+> pprCts (dictsToBag dicts)
, ppUnless (isEmptyTcAppMap safehask) $
text "Safe Haskell unsafe overlap =" <+> pprCts (dictsToBag safehask)
, ppUnless (isEmptyCts irreds) $
text "Irreds =" <+> pprCts irreds
, ppUnless (isEmptyCts blocked) $
text "Blocked =" <+> pprCts blocked
, ppUnless (null insts) $
text "Given instances =" <+> vcat (map ppr insts)
, text "Innermost given equalities =" <+> ppr ge_lvl
, text "Given eqs at this level =" <+> ppr given_eqs
]
where
folder (EqualCtList eqs) rest = nonEmptyToBag eqs `andCts` rest
{- *********************************************************************
* *
Shadow constraints and improvement
* *
************************************************************************
Note [The improvement story and derived shadows]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Because Wanteds cannot rewrite Wanteds (see Note [Wanteds do not
rewrite Wanteds] in GHC.Tc.Types.Constraint), we may miss some opportunities for
solving. Here's a classic example (indexed-types/should_fail/T4093a)
Ambiguity check for f: (Foo e ~ Maybe e) => Foo e
We get [G] Foo e ~ Maybe e (CEqCan)
[W] Foo ee ~ Foo e (CEqCan) -- ee is a unification variable
[W] Foo ee ~ Maybe ee (CEqCan)
The first Wanted gets rewritten to
[W] Foo ee ~ Maybe e
But now we appear to be stuck, since we don't rewrite Wanteds with
Wanteds. This is silly because we can see that ee := e is the
only solution.
The basic plan is
* generate Derived constraints that shadow Wanted constraints
* allow Derived to rewrite Derived
* in order to cause some unifications to take place
* that in turn solve the original Wanteds
The ONLY reason for all these Derived equalities is to tell us how to
unify a variable: that is, what Mark Jones calls "improvement".
The same idea is sometimes also called "saturation"; find all the
equalities that must hold in any solution.
Or, equivalently, you can think of the derived shadows as implementing
the "model": a non-idempotent but no-occurs-check substitution,
reflecting *all* *Nominal* equalities (a ~N ty) that are not
immediately soluble by unification.
More specifically, here's how it works (Oct 16):
* Wanted constraints are born as [WD]; this behaves like a
[W] and a [D] paired together.
* When we are about to add a [WD] to the inert set, if it can
be rewritten by a [D] a ~ ty, then we split it into [W] and [D],
putting the latter into the work list (see maybeEmitShadow).
In the example above, we get to the point where we are stuck:
[WD] Foo ee ~ Foo e
[WD] Foo ee ~ Maybe ee
But now when [WD] Foo ee ~ Maybe ee is about to be added, we'll
split it into [W] and [D], since the inert [WD] Foo ee ~ Foo e
can rewrite it. Then:
work item: [D] Foo ee ~ Maybe ee
inert: [W] Foo ee ~ Maybe ee
[WD] Foo ee ~ Maybe e
See Note [Splitting WD constraints]. Now the work item is rewritten
by the [WD] and we soon get ee := e.
Additional notes:
* The derived shadow equalities live in inert_eqs, along with
the Givens and Wanteds; see Note [EqualCtList invariants].
* We make Derived shadows only for Wanteds, not Givens. So we
have only [G], not [GD] and [G] plus splitting. See
Note [Add derived shadows only for Wanteds]
* We also get Derived equalities from functional dependencies
and type-function injectivity; see calls to unifyDerived.
* It's worth having [WD] rather than just [W] and [D] because
* efficiency: silly to process the same thing twice
* inert_dicts is a finite map keyed by
the type; it's inconvenient for it to map to TWO constraints
Another example requiring Deriveds is in
Note [Put touchable variables on the left] in GHC.Tc.Solver.Canonical.
Note [Splitting WD constraints]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
We are about to add a [WD] constraint to the inert set; and we
know that the inert set has fully rewritten it. Should we split
it into [W] and [D], and put the [D] in the work list for further
work?
* CDictCan (C tys):
Yes if the inert set could rewrite tys to make the class constraint,
or type family, fire. That is, yes if the inert_eqs intersects
with the free vars of tys. For this test we use
(anyRewritableTyVar True) which ignores casts and coercions in tys,
because rewriting the casts or coercions won't make the thing fire
more often.
* CEqCan (lhs ~ ty): Yes if the inert set could rewrite 'lhs' or 'ty'.
We need to check both 'lhs' and 'ty' against the inert set:
- Inert set contains [D] a ~ ty2
Then we want to put [D] a ~ ty in the worklist, so we'll
get [D] ty ~ ty2 with consequent good things
- Inert set contains [D] b ~ a, where b is in ty.
We can't just add [WD] a ~ ty[b] to the inert set, because
that breaks the inert-set invariants. If we tried to
canonicalise another [D] constraint mentioning 'a', we'd
get an infinite loop
Moreover we must use (anyRewritableTyVar False) for the RHS,
because even tyvars in the casts and coercions could give
an infinite loop if we don't expose it
* CIrredCan: Yes if the inert set can rewrite the constraint.
We used to think splitting irreds was unnecessary, but
see Note [Splitting Irred WD constraints]
* Others: nothing is gained by splitting.
Note [Splitting Irred WD constraints]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Splitting Irred constraints can make a difference. Here is the
scenario:
a[sk] :: F v -- F is a type family
beta :: alpha
work item: [WD] a ~ beta
This is heterogeneous, so we emit a kind equality and make the work item an
inert Irred.
work item: [D] F v ~ alpha
inert: [WD] (a |> co) ~ beta (CIrredCan)
Can't make progress on the work item. Add to inert set. This kicks out the
old inert, because a [D] can rewrite a [WD].
work item: [WD] (a |> co) ~ beta
inert: [D] F v ~ alpha (CEqCan)
Can't make progress on this work item either (although GHC tries by
decomposing the cast and rewriting... but that doesn't make a difference),
which is still hetero. Emit a new kind equality and add to inert set. But,
critically, we split the Irred.
work list:
[D] F v ~ alpha (CEqCan)
[D] (a |> co) ~ beta (CIrred) -- this one was split off
inert:
[W] (a |> co) ~ beta
[D] F v ~ alpha
We quickly solve the first work item, as it's the same as an inert.
work item: [D] (a |> co) ~ beta
inert:
[W] (a |> co) ~ beta
[D] F v ~ alpha
We decompose the cast, yielding
[D] a ~ beta
We then rewrite the kinds. The lhs kind is F v, which flattens to alpha.
co' :: F v ~ alpha
[D] (a |> co') ~ beta
Now this equality is homo-kinded. So we swizzle it around to
[D] beta ~ (a |> co')
and set beta := a |> co', and go home happy.
If we don't split the Irreds, we loop. This is all dangerously subtle.
This is triggered by test case typecheck/should_compile/SplitWD.
Note [Add derived shadows only for Wanteds]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
We only add shadows for Wanted constraints. That is, we have
[WD] but not [GD]; and maybeEmitShaodw looks only at [WD]
constraints.
It does just possibly make sense ot add a derived shadow for a
Given. If we created a Derived shadow of a Given, it could be
rewritten by other Deriveds, and that could, conceivably, lead to a
useful unification.
But (a) I have been unable to come up with an example of this
happening
(b) see #12660 for how adding the derived shadows
of a Given led to an infinite loop.
(c) It's unlikely that rewriting derived Givens will lead
to a unification because Givens don't mention touchable
unification variables
For (b) there may be other ways to solve the loop, but simply
reraining from adding derived shadows of Givens is particularly
simple. And it's more efficient too!
Still, here's one possible reason for adding derived shadows
for Givens. Consider
work-item [G] a ~ [b], inerts has [D] b ~ a.
If we added the derived shadow (into the work list)
[D] a ~ [b]
When we process it, we'll rewrite to a ~ [a] and get an
occurs check. Without it we'll miss the occurs check (reporting
inaccessible code); but that's probably OK.
Note [Keep CDictCan shadows as CDictCan]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Suppose we have
class C a => D a b
and [G] D a b, [G] C a in the inert set. Now we insert
[D] b ~ c. We want to kick out a derived shadow for [D] D a b,
so we can rewrite it with the new constraint, and perhaps get
instance reduction or other consequences.
BUT we do not want to kick out a *non-canonical* (D a b). If we
did, we would do this:
- rewrite it to [D] D a c, with pend_sc = True
- use expandSuperClasses to add C a
- go round again, which solves C a from the givens
This loop goes on for ever and triggers the simpl_loop limit.
Solution: kick out the CDictCan which will have pend_sc = False,
because we've already added its superclasses. So we won't re-add
them. If we forget the pend_sc flag, our cunning scheme for avoiding
generating superclasses repeatedly will fail.
See #11379 for a case of this.
Note [Do not do improvement for WOnly]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
We do improvement between two constraints (e.g. for injectivity
or functional dependencies) only if both are "improvable". And
we improve a constraint wrt the top-level instances only if
it is improvable.
Improvable: [G] [WD] [D}
Not improvable: [W]
Reasons:
* It's less work: fewer pairs to compare
* Every [W] has a shadow [D] so nothing is lost
* Consider [WD] C Int b, where 'b' is a skolem, and
class C a b | a -> b
instance C Int Bool
We'll do a fundep on it and emit [D] b ~ Bool
That will kick out constraint [WD] C Int b
Then we'll split it to [W] C Int b (keep in inert)
and [D] C Int b (in work list)
When processing the latter we'll rewrite it to
[D] C Int Bool
At that point it would be /stupid/ to interact it
with the inert [W] C Int b in the inert set; after all,
it's the very constraint from which the [D] C Int Bool
was split! We can avoid this by not doing improvement
on [W] constraints. This came up in #12860.
-}
maybeEmitShadow :: InertCans -> Ct -> TcS Ct
-- See Note [The improvement story and derived shadows]
maybeEmitShadow ics ct
| let ev = ctEvidence ct
, CtWanted { ctev_pred = pred, ctev_loc = loc
, ctev_nosh = WDeriv } <- ev
, shouldSplitWD (inert_eqs ics) (inert_funeqs ics) ct
= do { traceTcS "Emit derived shadow" (ppr ct)
; let derived_ev = CtDerived { ctev_pred = pred
, ctev_loc = loc }
shadow_ct = ct { cc_ev = derived_ev }
-- Te shadow constraint keeps the canonical shape.
-- This just saves work, but is sometimes important;
-- see Note [Keep CDictCan shadows as CDictCan]
; emitWork [shadow_ct]
; let ev' = ev { ctev_nosh = WOnly }
ct' = ct { cc_ev = ev' }
-- Record that it now has a shadow
-- This is /the/ place we set the flag to WOnly
; return ct' }
| otherwise
= return ct
shouldSplitWD :: InertEqs -> FunEqMap EqualCtList -> Ct -> Bool
-- Precondition: 'ct' is [WD], and is inert
-- True <=> we should split ct ito [W] and [D] because
-- the inert_eqs can make progress on the [D]
-- See Note [Splitting WD constraints]
shouldSplitWD inert_eqs fun_eqs (CDictCan { cc_tyargs = tys })
= should_split_match_args inert_eqs fun_eqs tys
-- NB True: ignore coercions
-- See Note [Splitting WD constraints]
shouldSplitWD inert_eqs fun_eqs (CEqCan { cc_lhs = TyVarLHS tv, cc_rhs = ty
, cc_eq_rel = eq_rel })
= tv `elemDVarEnv` inert_eqs
|| anyRewritableCanEqLHS eq_rel (canRewriteTv inert_eqs) (canRewriteTyFam fun_eqs) ty
-- NB False: do not ignore casts and coercions
-- See Note [Splitting WD constraints]
shouldSplitWD inert_eqs fun_eqs (CEqCan { cc_ev = ev, cc_eq_rel = eq_rel })
= anyRewritableCanEqLHS eq_rel (canRewriteTv inert_eqs) (canRewriteTyFam fun_eqs)
(ctEvPred ev)
shouldSplitWD inert_eqs fun_eqs (CIrredCan { cc_ev = ev })
= anyRewritableCanEqLHS (ctEvEqRel ev) (canRewriteTv inert_eqs)
(canRewriteTyFam fun_eqs) (ctEvPred ev)
shouldSplitWD _ _ _ = False -- No point in splitting otherwise
should_split_match_args :: InertEqs -> FunEqMap EqualCtList -> [TcType] -> Bool
-- True if the inert_eqs can rewrite anything in the argument types
should_split_match_args inert_eqs fun_eqs tys
= any (anyRewritableCanEqLHS NomEq (canRewriteTv inert_eqs) (canRewriteTyFam fun_eqs)) tys
-- See Note [Splitting WD constraints]
canRewriteTv :: InertEqs -> EqRel -> TyVar -> Bool
canRewriteTv inert_eqs eq_rel tv
| Just (EqualCtList (ct :| _)) <- lookupDVarEnv inert_eqs tv
, CEqCan { cc_eq_rel = eq_rel1 } <- ct
= eq_rel1 `eqCanRewrite` eq_rel
| otherwise
= False
canRewriteTyFam :: FunEqMap EqualCtList -> EqRel -> TyCon -> [Type] -> Bool
canRewriteTyFam fun_eqs eq_rel tf args
| Just (EqualCtList (ct :| _)) <- findFunEq fun_eqs tf args
, CEqCan { cc_eq_rel = eq_rel1 } <- ct
= eq_rel1 `eqCanRewrite` eq_rel
| otherwise
= False
isImprovable :: CtEvidence -> Bool
-- See Note [Do not do improvement for WOnly]
isImprovable (CtWanted { ctev_nosh = WOnly }) = False
isImprovable _ = True
{- *********************************************************************
* *
Inert equalities
* *
********************************************************************* -}
addTyEq :: InertEqs -> TcTyVar -> Ct -> InertEqs
addTyEq old_eqs tv ct
= extendDVarEnv_C add_eq old_eqs tv (unitEqualCtList ct)
where
add_eq old_eqs _ = addToEqualCtList ct old_eqs
addCanFunEq :: FunEqMap EqualCtList -> TyCon -> [TcType] -> Ct
-> FunEqMap EqualCtList
addCanFunEq old_eqs fun_tc fun_args ct
= alterTcApp old_eqs fun_tc fun_args upd
where
upd (Just old_equal_ct_list) = Just $ addToEqualCtList ct old_equal_ct_list
upd Nothing = Just $ unitEqualCtList ct
foldTyEqs :: (Ct -> b -> b) -> InertEqs -> b -> b
foldTyEqs k eqs z
= foldDVarEnv (\(EqualCtList cts) z -> foldr k z cts) z eqs
findTyEqs :: InertCans -> TyVar -> [Ct]
findTyEqs icans tv = maybe [] id (fmap @Maybe equalCtListToList $
lookupDVarEnv (inert_eqs icans) tv)
delEq :: InertCans -> CanEqLHS -> TcType -> InertCans
delEq ic lhs rhs = case lhs of
TyVarLHS tv
-> ic { inert_eqs = alterDVarEnv upd (inert_eqs ic) tv }
TyFamLHS tf args
-> ic { inert_funeqs = alterTcApp (inert_funeqs ic) tf args upd }
where
isThisOne :: Ct -> Bool
isThisOne (CEqCan { cc_rhs = t1 }) = tcEqTypeNoKindCheck rhs t1
isThisOne other = pprPanic "delEq" (ppr lhs $$ ppr ic $$ ppr other)
upd :: Maybe EqualCtList -> Maybe EqualCtList
upd (Just eq_ct_list) = filterEqualCtList (not . isThisOne) eq_ct_list
upd Nothing = Nothing
findEq :: InertCans -> CanEqLHS -> [Ct]
findEq icans (TyVarLHS tv) = findTyEqs icans tv
findEq icans (TyFamLHS fun_tc fun_args)
= maybe [] id (fmap @Maybe equalCtListToList $
findFunEq (inert_funeqs icans) fun_tc fun_args)
{- *********************************************************************
* *
Inert instances: inert_insts
* *
********************************************************************* -}
addInertForAll :: QCInst -> TcS ()
-- Add a local Given instance, typically arising from a type signature
addInertForAll new_qci
= do { ics <- getInertCans
; ics1 <- add_qci ics
-- Update given equalities. C.f updateGivenEqs
; tclvl <- getTcLevel
; let pred = qci_pred new_qci
not_equality = isClassPred pred && not (isEqPred pred)
-- True <=> definitely not an equality
-- A qci_pred like (f a) might be an equality
ics2 | not_equality = ics1
| otherwise = ics1 { inert_given_eq_lvl = tclvl
, inert_given_eqs = True }
; setInertCans ics2 }
where
add_qci :: InertCans -> TcS InertCans
-- See Note [Do not add duplicate quantified instances]
add_qci ics@(IC { inert_insts = qcis })
| any same_qci qcis
= do { traceTcS "skipping duplicate quantified instance" (ppr new_qci)
; return ics }
| otherwise
= do { traceTcS "adding new inert quantified instance" (ppr new_qci)
; return (ics { inert_insts = new_qci : qcis }) }
same_qci old_qci = tcEqType (ctEvPred (qci_ev old_qci))
(ctEvPred (qci_ev new_qci))
{- Note [Do not add duplicate quantified instances]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Consider this (#15244):
f :: (C g, D g) => ....
class S g => C g where ...
class S g => D g where ...
class (forall a. Eq a => Eq (g a)) => S g where ...
Then in f's RHS there are two identical quantified constraints
available, one via the superclasses of C and one via the superclasses
of D. The two are identical, and it seems wrong to reject the program
because of that. But without doing duplicate-elimination we will have
two matching QCInsts when we try to solve constraints arising from f's
RHS.
The simplest thing is simply to eliminate duplicates, which we do here.
-}
{- *********************************************************************
* *
Adding an inert
* *
************************************************************************
Note [Adding an equality to the InertCans]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
When adding an equality to the inerts:
* Split [WD] into [W] and [D] if the inerts can rewrite the latter;
done by maybeEmitShadow.
* Kick out any constraints that can be rewritten by the thing
we are adding. Done by kickOutRewritable.
* Note that unifying a:=ty, is like adding [G] a~ty; just use
kickOutRewritable with Nominal, Given. See kickOutAfterUnification.
-}
addInertCan :: Ct -> TcS ()
-- Precondition: item /is/ canonical
-- See Note [Adding an equality to the InertCans]
addInertCan ct
= do { traceTcS "addInertCan {" $
text "Trying to insert new inert item:" <+> ppr ct
; ics <- getInertCans
; ct <- maybeEmitShadow ics ct
; ics <- maybeKickOut ics ct
; tclvl <- getTcLevel
; setInertCans (add_item tclvl ics ct)
; traceTcS "addInertCan }" $ empty }
maybeKickOut :: InertCans -> Ct -> TcS InertCans
-- For a CEqCan, kick out any inert that can be rewritten by the CEqCan
maybeKickOut ics ct
| CEqCan { cc_lhs = lhs, cc_ev = ev, cc_eq_rel = eq_rel } <- ct
= do { (_, ics') <- kickOutRewritable (ctEvFlavour ev, eq_rel) lhs ics
; return ics' }
| otherwise
= return ics
add_item :: TcLevel -> InertCans -> Ct -> InertCans
add_item tc_lvl
ics@(IC { inert_funeqs = funeqs, inert_eqs = eqs })
item@(CEqCan { cc_lhs = lhs })
= updateGivenEqs tc_lvl item $
case lhs of
TyFamLHS tc tys -> ics { inert_funeqs = addCanFunEq funeqs tc tys item }
TyVarLHS tv -> ics { inert_eqs = addTyEq eqs tv item }
add_item tc_lvl ics@(IC { inert_blocked = blocked })
item@(CIrredCan { cc_reason = HoleBlockerReason {}})
= updateGivenEqs tc_lvl item $ -- this item is always an equality
ics { inert_blocked = blocked `snocBag` item }
add_item tc_lvl ics@(IC { inert_irreds = irreds }) item@(CIrredCan {})
= updateGivenEqs tc_lvl item $ -- An Irred might turn out to be an
-- equality, so we play safe
ics { inert_irreds = irreds `Bag.snocBag` item }
add_item _ ics item@(CDictCan { cc_class = cls, cc_tyargs = tys })
= ics { inert_dicts = addDictCt (inert_dicts ics) cls tys item }
add_item _ _ item
= pprPanic "upd_inert set: can't happen! Inserting " $
ppr item -- Can't be CNonCanonical because they only land in inert_irreds
updateGivenEqs :: TcLevel -> Ct -> InertCans -> InertCans
-- Set the inert_given_eq_level to the current level (tclvl)
-- if the constraint is a given equality that should prevent
-- filling in an outer unification variable.
-- See See Note [Tracking Given equalities]
updateGivenEqs tclvl ct inerts@(IC { inert_given_eq_lvl = ge_lvl })
| not (isGivenCt ct) = inerts
| not_equality ct = inerts -- See Note [Let-bound skolems]
| otherwise = inerts { inert_given_eq_lvl = ge_lvl'
, inert_given_eqs = True }
where
ge_lvl' | mentionsOuterVar tclvl (ctEvidence ct)
-- Includes things like (c a), which *might* be an equality
= tclvl
| otherwise
= ge_lvl
not_equality :: Ct -> Bool
-- True <=> definitely not an equality of any kind
-- except for a let-bound skolem, which doesn't count
-- See Note [Let-bound skolems]
-- NB: no need to spot the boxed CDictCan (a ~ b) because its
-- superclass (a ~# b) will be a CEqCan
not_equality (CEqCan { cc_lhs = TyVarLHS tv }) = not (isOuterTyVar tclvl tv)
not_equality (CDictCan {}) = True
not_equality _ = False
-----------------------------------------
kickOutRewritable :: CtFlavourRole -- Flavour/role of the equality that
-- is being added to the inert set
-> CanEqLHS -- The new equality is lhs ~ ty
-> InertCans
-> TcS (Int, InertCans)
kickOutRewritable new_fr new_lhs ics
= do { let (kicked_out, ics') = kick_out_rewritable new_fr new_lhs ics
n_kicked = workListSize kicked_out
; unless (n_kicked == 0) $
do { updWorkListTcS (appendWorkList kicked_out)
-- The famapp-cache contains Given evidence from the inert set.
-- If we're kicking out Givens, we need to remove this evidence
-- from the cache, too.
; let kicked_given_ev_vars =
[ ev_var | ct <- wl_eqs kicked_out
, CtGiven { ctev_evar = ev_var } <- [ctEvidence ct] ]
; when (new_fr `eqCanRewriteFR` (Given, NomEq) &&
-- if this isn't true, no use looking through the constraints
not (null kicked_given_ev_vars)) $
do { traceTcS "Given(s) have been kicked out; drop from famapp-cache"
(ppr kicked_given_ev_vars)
; dropFromFamAppCache (mkVarSet kicked_given_ev_vars) }
; csTraceTcS $
hang (text "Kick out, lhs =" <+> ppr new_lhs)
2 (vcat [ text "n-kicked =" <+> int n_kicked
, text "kicked_out =" <+> ppr kicked_out
, text "Residual inerts =" <+> ppr ics' ]) }
; return (n_kicked, ics') }
kick_out_rewritable :: CtFlavourRole -- Flavour/role of the equality that
-- is being added to the inert set
-> CanEqLHS -- The new equality is lhs ~ ty
-> InertCans
-> (WorkList, InertCans)
-- See Note [kickOutRewritable]
kick_out_rewritable new_fr new_lhs
ics@(IC { inert_eqs = tv_eqs
, inert_dicts = dictmap
, inert_funeqs = funeqmap
, inert_irreds = irreds
, inert_insts = old_insts })
| not (new_fr `eqMayRewriteFR` new_fr)
= (emptyWorkList, ics)
-- If new_fr can't rewrite itself, it can't rewrite
-- anything else, so no need to kick out anything.
-- (This is a common case: wanteds can't rewrite wanteds)
-- Lemma (L2) in Note [Extending the inert equalities]
| otherwise
= (kicked_out, inert_cans_in)
where
-- inert_safehask stays unchanged; is that right?
inert_cans_in = ics { inert_eqs = tv_eqs_in
, inert_dicts = dicts_in
, inert_funeqs = feqs_in
, inert_irreds = irs_in
, inert_insts = insts_in }
kicked_out :: WorkList
-- NB: use extendWorkList to ensure that kicked-out equalities get priority
-- See Note [Prioritise equalities] (Kick-out).
-- The irreds may include non-canonical (hetero-kinded) equality
-- constraints, which perhaps may have become soluble after new_lhs
-- is substituted; ditto the dictionaries, which may include (a~b)
-- or (a~~b) constraints.
kicked_out = foldr extendWorkListCt
(emptyWorkList { wl_eqs = tv_eqs_out ++ feqs_out })
((dicts_out `andCts` irs_out)
`extendCtsList` insts_out)
(tv_eqs_out, tv_eqs_in) = foldDVarEnv (kick_out_eqs extend_tv_eqs)
([], emptyDVarEnv) tv_eqs
(feqs_out, feqs_in) = foldFunEqs (kick_out_eqs extend_fun_eqs)
funeqmap ([], emptyFunEqs)
(dicts_out, dicts_in) = partitionDicts kick_out_ct dictmap
(irs_out, irs_in) = partitionBag kick_out_ct irreds
-- Kick out even insolubles: See Note [Rewrite insolubles]
-- Of course we must kick out irreducibles like (c a), in case
-- we can rewrite 'c' to something more useful
-- Kick-out for inert instances
-- See Note [Quantified constraints] in GHC.Tc.Solver.Canonical
insts_out :: [Ct]
insts_in :: [QCInst]
(insts_out, insts_in)
| fr_may_rewrite (Given, NomEq) -- All the insts are Givens
= partitionWith kick_out_qci old_insts
| otherwise
= ([], old_insts)
kick_out_qci qci
| let ev = qci_ev qci
, fr_can_rewrite_ty NomEq (ctEvPred (qci_ev qci))
= Left (mkNonCanonical ev)
| otherwise
= Right qci
(_, new_role) = new_fr
fr_tv_can_rewrite_ty :: TyVar -> EqRel -> Type -> Bool
fr_tv_can_rewrite_ty new_tv role ty
= anyRewritableTyVar True role can_rewrite ty
-- True: ignore casts and coercions
where
can_rewrite :: EqRel -> TyVar -> Bool
can_rewrite old_role tv = new_role `eqCanRewrite` old_role && tv == new_tv
fr_tf_can_rewrite_ty :: TyCon -> [TcType] -> EqRel -> Type -> Bool
fr_tf_can_rewrite_ty new_tf new_tf_args role ty
= anyRewritableTyFamApp role can_rewrite ty
where
can_rewrite :: EqRel -> TyCon -> [TcType] -> Bool
can_rewrite old_role old_tf old_tf_args
= new_role `eqCanRewrite` old_role &&
tcEqTyConApps new_tf new_tf_args old_tf old_tf_args
-- it's possible for old_tf_args to have too many. This is fine;
-- we'll only check what we need to.
{-# INLINE fr_can_rewrite_ty #-} -- perform the check here only once
fr_can_rewrite_ty :: EqRel -> Type -> Bool
fr_can_rewrite_ty = case new_lhs of
TyVarLHS new_tv -> fr_tv_can_rewrite_ty new_tv
TyFamLHS new_tf new_tf_args -> fr_tf_can_rewrite_ty new_tf new_tf_args
fr_may_rewrite :: CtFlavourRole -> Bool
fr_may_rewrite fs = new_fr `eqMayRewriteFR` fs
-- Can the new item rewrite the inert item?
{-# INLINE kick_out_ct #-} -- perform case on new_lhs here only once
kick_out_ct :: Ct -> Bool
-- Kick it out if the new CEqCan can rewrite the inert one
-- See Note [kickOutRewritable]
kick_out_ct = case new_lhs of
TyVarLHS new_tv -> \ct -> let fs@(_,role) = ctFlavourRole ct in
fr_may_rewrite fs
&& fr_tv_can_rewrite_ty new_tv role (ctPred ct)
TyFamLHS new_tf new_tf_args
-> \ct -> let fs@(_, role) = ctFlavourRole ct in
fr_may_rewrite fs
&& fr_tf_can_rewrite_ty new_tf new_tf_args role (ctPred ct)
extend_tv_eqs :: InertEqs -> CanEqLHS -> EqualCtList -> InertEqs
extend_tv_eqs eqs (TyVarLHS tv) cts = extendDVarEnv eqs tv cts
extend_tv_eqs eqs other _cts = pprPanic "extend_tv_eqs" (ppr eqs $$ ppr other)
extend_fun_eqs :: FunEqMap EqualCtList -> CanEqLHS -> EqualCtList
-> FunEqMap EqualCtList
extend_fun_eqs eqs (TyFamLHS fam_tc fam_args) cts
= insertTcApp eqs fam_tc fam_args cts
extend_fun_eqs eqs other _cts = pprPanic "extend_fun_eqs" (ppr eqs $$ ppr other)
kick_out_eqs :: (container -> CanEqLHS -> EqualCtList -> container)
-> EqualCtList -> ([Ct], container)
-> ([Ct], container)
kick_out_eqs extend eqs (acc_out, acc_in)
= (eqs_out `chkAppend` acc_out, case listToEqualCtList eqs_in of
Nothing -> acc_in
Just eqs_in_ecl@(EqualCtList (eq1 :| _))
-> extend acc_in (cc_lhs eq1) eqs_in_ecl)
where
(eqs_out, eqs_in) = partition kick_out_eq (equalCtListToList eqs)
-- Implements criteria K1-K3 in Note [Extending the inert equalities]
kick_out_eq (CEqCan { cc_lhs = lhs, cc_rhs = rhs_ty
, cc_ev = ev, cc_eq_rel = eq_rel })
| not (fr_may_rewrite fs)
= False -- (K0) Keep it in the inert set if the new thing can't rewrite it
-- Below here (fr_may_rewrite fs) is True
| TyVarLHS _ <- lhs
, fs `eqMayRewriteFR` new_fr
= False -- (K4) Keep it in the inert set if the LHS is a tyvar and
-- it can rewrite the work item. See Note [K4]
| fr_can_rewrite_ty eq_rel (canEqLHSType lhs)
= True -- (K1)
-- The above check redundantly checks the role & flavour,
-- but it's very convenient
| kick_out_for_inertness = True -- (K2)
| kick_out_for_completeness = True -- (K3)
| otherwise = False
where
fs = (ctEvFlavour ev, eq_rel)
kick_out_for_inertness
= (fs `eqMayRewriteFR` fs) -- (K2a)
&& fr_can_rewrite_ty eq_rel rhs_ty -- (K2b)
kick_out_for_completeness -- (K3) and Note [K3: completeness of solving]
= case eq_rel of
NomEq -> rhs_ty `eqType` canEqLHSType new_lhs -- (K3a)
ReprEq -> is_can_eq_lhs_head new_lhs rhs_ty -- (K3b)
kick_out_eq ct = pprPanic "keep_eq" (ppr ct)
is_can_eq_lhs_head (TyVarLHS tv) = go
where
go (Rep.TyVarTy tv') = tv == tv'
go (Rep.AppTy fun _) = go fun
go (Rep.CastTy ty _) = go ty
go (Rep.TyConApp {}) = False
go (Rep.LitTy {}) = False
go (Rep.ForAllTy {}) = False
go (Rep.FunTy {}) = False
go (Rep.CoercionTy {}) = False
is_can_eq_lhs_head (TyFamLHS fun_tc fun_args) = go
where
go (Rep.TyVarTy {}) = False
go (Rep.AppTy {}) = False -- no TyConApp to the left of an AppTy
go (Rep.CastTy ty _) = go ty
go (Rep.TyConApp tc args) = tcEqTyConApps fun_tc fun_args tc args
go (Rep.LitTy {}) = False
go (Rep.ForAllTy {}) = False
go (Rep.FunTy {}) = False
go (Rep.CoercionTy {}) = False
kickOutAfterUnification :: TcTyVar -> TcS Int
kickOutAfterUnification new_tv
= do { ics <- getInertCans
; (n_kicked, ics2) <- kickOutRewritable (Given,NomEq)
(TyVarLHS new_tv) ics
-- Given because the tv := xi is given; NomEq because
-- only nominal equalities are solved by unification
; setInertCans ics2
; return n_kicked }
-- See Wrinkle (2) in Note [Equalities with incompatible kinds] in GHC.Tc.Solver.Canonical
kickOutAfterFillingCoercionHole :: CoercionHole -> Coercion -> TcS ()
kickOutAfterFillingCoercionHole hole filled_co
= do { ics <- getInertCans
; let (kicked_out, ics') = kick_out ics
n_kicked = workListSize kicked_out
; unless (n_kicked == 0) $
do { updWorkListTcS (appendWorkList kicked_out)
; csTraceTcS $
hang (text "Kick out, hole =" <+> ppr hole)
2 (vcat [ text "n-kicked =" <+> int n_kicked
, text "kicked_out =" <+> ppr kicked_out
, text "Residual inerts =" <+> ppr ics' ]) }
; setInertCans ics' }
where
holes_of_co = coercionHolesOfCo filled_co
kick_out :: InertCans -> (WorkList, InertCans)
kick_out ics@(IC { inert_blocked = blocked })
= let (to_kick, to_keep) = partitionBagWith kick_ct blocked
kicked_out = extendWorkListCts (bagToList to_kick) emptyWorkList
ics' = ics { inert_blocked = to_keep }
in
(kicked_out, ics')
kick_ct :: Ct -> Either Ct Ct
-- Left: kick out; Right: keep. But even if we keep, we may need
-- to update the set of blocking holes
kick_ct ct@(CIrredCan { cc_reason = HoleBlockerReason holes })
| hole `elementOfUniqSet` holes
= let new_holes = holes `delOneFromUniqSet` hole
`unionUniqSets` holes_of_co
updated_ct = ct { cc_reason = HoleBlockerReason new_holes }
in
if isEmptyUniqSet new_holes
then Left updated_ct
else Right updated_ct
| otherwise
= Right ct
kick_ct other = pprPanic "kickOutAfterFillingCoercionHole" (ppr other)
{- Note [kickOutRewritable]
~~~~~~~~~~~~~~~~~~~~~~~~~~~
See also Note [inert_eqs: the inert equalities].
When we add a new inert equality (lhs ~N ty) to the inert set,
we must kick out any inert items that could be rewritten by the
new equality, to maintain the inert-set invariants.
- We want to kick out an existing inert constraint if
a) the new constraint can rewrite the inert one
b) 'lhs' is free in the inert constraint (so that it *will*)
rewrite it if we kick it out.
For (b) we use anyRewritableCanLHS, which examines the types /and
kinds/ that are directly visible in the type. Hence
we will have exposed all the rewriting we care about to make the
most precise kinds visible for matching classes etc. No need to
kick out constraints that mention type variables whose kinds
contain this LHS!
- A Derived equality can kick out [D] constraints in inert_eqs,
inert_dicts, inert_irreds etc.
- We don't kick out constraints from inert_solved_dicts, and
inert_solved_funeqs optimistically. But when we lookup we have to
take the substitution into account
NB: we could in principle avoid kick-out:
a) When unifying a meta-tyvar from an outer level, because
then the entire implication will be iterated; see
Note [The Unification Level Flag]
b) For Givens, after a unification. By (GivenInv) in GHC.Tc.Utils.TcType
Note [TcLevel invariants], a Given can't include a meta-tyvar from
its own level, so it falls under (a). Of course, we must still
kick out Givens when adding a new non-unification Given.
But kicking out more vigorously may lead to earlier unification and fewer
iterations, so we don't take advantage of these possibilities.
Note [Rewrite insolubles]
~~~~~~~~~~~~~~~~~~~~~~~~~
Suppose we have an insoluble alpha ~ [alpha], which is insoluble
because an occurs check. And then we unify alpha := [Int]. Then we
really want to rewrite the insoluble to [Int] ~ [[Int]]. Now it can
be decomposed. Otherwise we end up with a "Can't match [Int] ~
[[Int]]" which is true, but a bit confusing because the outer type
constructors match.
Hence:
* In the main simplifier loops in GHC.Tc.Solver (solveWanteds,
simpl_loop), we feed the insolubles in solveSimpleWanteds,
so that they get rewritten (albeit not solved).
* We kick insolubles out of the inert set, if they can be
rewritten (see GHC.Tc.Solver.Monad.kick_out_rewritable)
* We rewrite those insolubles in GHC.Tc.Solver.Canonical.
See Note [Make sure that insolubles are fully rewritten]
-}
--------------
addInertSafehask :: InertCans -> Ct -> InertCans
addInertSafehask ics item@(CDictCan { cc_class = cls, cc_tyargs = tys })
= ics { inert_safehask = addDictCt (inert_dicts ics) cls tys item }
addInertSafehask _ item
= pprPanic "addInertSafehask: can't happen! Inserting " $ ppr item
insertSafeOverlapFailureTcS :: InstanceWhat -> Ct -> TcS ()
-- See Note [Safe Haskell Overlapping Instances Implementation] in GHC.Tc.Solver
insertSafeOverlapFailureTcS what item
| safeOverlap what = return ()
| otherwise = updInertCans (\ics -> addInertSafehask ics item)
getSafeOverlapFailures :: TcS Cts
-- See Note [Safe Haskell Overlapping Instances Implementation] in GHC.Tc.Solver
getSafeOverlapFailures
= do { IC { inert_safehask = safehask } <- getInertCans
; return $ foldDicts consCts safehask emptyCts }
--------------
addSolvedDict :: InstanceWhat -> CtEvidence -> Class -> [Type] -> TcS ()
-- Conditionally add a new item in the solved set of the monad
-- See Note [Solved dictionaries]
addSolvedDict what item cls tys
| isWanted item
, instanceReturnsDictCon what
= do { traceTcS "updSolvedSetTcs:" $ ppr item
; updInertTcS $ \ ics ->
ics { inert_solved_dicts = addDict (inert_solved_dicts ics) cls tys item } }
| otherwise
= return ()
getSolvedDicts :: TcS (DictMap CtEvidence)
getSolvedDicts = do { ics <- getTcSInerts; return (inert_solved_dicts ics) }
setSolvedDicts :: DictMap CtEvidence -> TcS ()
setSolvedDicts solved_dicts
= updInertTcS $ \ ics ->
ics { inert_solved_dicts = solved_dicts }
{- *********************************************************************
* *
Other inert-set operations
* *
********************************************************************* -}
updInertTcS :: (InertSet -> InertSet) -> TcS ()
-- Modify the inert set with the supplied function
updInertTcS upd_fn
= do { is_var <- getTcSInertsRef
; wrapTcS (do { curr_inert <- TcM.readTcRef is_var
; TcM.writeTcRef is_var (upd_fn curr_inert) }) }
getInertCans :: TcS InertCans
getInertCans = do { inerts <- getTcSInerts; return (inert_cans inerts) }
setInertCans :: InertCans -> TcS ()
setInertCans ics = updInertTcS $ \ inerts -> inerts { inert_cans = ics }
updRetInertCans :: (InertCans -> (a, InertCans)) -> TcS a
-- Modify the inert set with the supplied function
updRetInertCans upd_fn
= do { is_var <- getTcSInertsRef
; wrapTcS (do { inerts <- TcM.readTcRef is_var
; let (res, cans') = upd_fn (inert_cans inerts)
; TcM.writeTcRef is_var (inerts { inert_cans = cans' })
; return res }) }
updInertCans :: (InertCans -> InertCans) -> TcS ()
-- Modify the inert set with the supplied function
updInertCans upd_fn
= updInertTcS $ \ inerts -> inerts { inert_cans = upd_fn (inert_cans inerts) }
updInertDicts :: (DictMap Ct -> DictMap Ct) -> TcS ()
-- Modify the inert set with the supplied function
updInertDicts upd_fn
= updInertCans $ \ ics -> ics { inert_dicts = upd_fn (inert_dicts ics) }
updInertSafehask :: (DictMap Ct -> DictMap Ct) -> TcS ()
-- Modify the inert set with the supplied function
updInertSafehask upd_fn
= updInertCans $ \ ics -> ics { inert_safehask = upd_fn (inert_safehask ics) }
updInertIrreds :: (Cts -> Cts) -> TcS ()
-- Modify the inert set with the supplied function
updInertIrreds upd_fn
= updInertCans $ \ ics -> ics { inert_irreds = upd_fn (inert_irreds ics) }
getInertEqs :: TcS (DTyVarEnv EqualCtList)
getInertEqs = do { inert <- getInertCans; return (inert_eqs inert) }
getInnermostGivenEqLevel :: TcS TcLevel
getInnermostGivenEqLevel = do { inert <- getInertCans
; return (inert_given_eq_lvl inert) }
getInertInsols :: TcS Cts
-- Returns insoluble equality constraints
-- specifically including Givens
getInertInsols = do { inert <- getInertCans
; return (filterBag insolubleEqCt (inert_irreds inert)) }
getInertGivens :: TcS [Ct]
-- Returns the Given constraints in the inert set
getInertGivens
= do { inerts <- getInertCans
; let all_cts = foldDicts (:) (inert_dicts inerts)
$ foldFunEqs (\ecl out -> equalCtListToList ecl ++ out)
(inert_funeqs inerts)
$ concatMap equalCtListToList (dVarEnvElts (inert_eqs inerts))
; return (filter isGivenCt all_cts) }
getPendingGivenScs :: TcS [Ct]
-- Find all inert Given dictionaries, or quantified constraints,
-- whose cc_pend_sc flag is True
-- and that belong to the current level
-- Set their cc_pend_sc flag to False in the inert set, and return that Ct
getPendingGivenScs = do { lvl <- getTcLevel
; updRetInertCans (get_sc_pending lvl) }
get_sc_pending :: TcLevel -> InertCans -> ([Ct], InertCans)
get_sc_pending this_lvl ic@(IC { inert_dicts = dicts, inert_insts = insts })
= ASSERT2( all isGivenCt sc_pending, ppr sc_pending )
-- When getPendingScDics is called,
-- there are never any Wanteds in the inert set
(sc_pending, ic { inert_dicts = dicts', inert_insts = insts' })
where
sc_pending = sc_pend_insts ++ sc_pend_dicts
sc_pend_dicts = foldDicts get_pending dicts []
dicts' = foldr add dicts sc_pend_dicts
(sc_pend_insts, insts') = mapAccumL get_pending_inst [] insts
get_pending :: Ct -> [Ct] -> [Ct] -- Get dicts with cc_pend_sc = True
-- but flipping the flag
get_pending dict dicts
| Just dict' <- isPendingScDict dict
, belongs_to_this_level (ctEvidence dict)
= dict' : dicts
| otherwise
= dicts
add :: Ct -> DictMap Ct -> DictMap Ct
add ct@(CDictCan { cc_class = cls, cc_tyargs = tys }) dicts
= addDictCt dicts cls tys ct
add ct _ = pprPanic "getPendingScDicts" (ppr ct)
get_pending_inst :: [Ct] -> QCInst -> ([Ct], QCInst)
get_pending_inst cts qci@(QCI { qci_ev = ev })
| Just qci' <- isPendingScInst qci
, belongs_to_this_level ev
= (CQuantCan qci' : cts, qci')
| otherwise
= (cts, qci)
belongs_to_this_level ev = ctLocLevel (ctEvLoc ev) == this_lvl
-- We only want Givens from this level; see (3a) in
-- Note [The superclass story] in GHC.Tc.Solver.Canonical
getUnsolvedInerts :: TcS ( Bag Implication
, Cts ) -- All simple constraints
-- Return all the unsolved [Wanted] or [Derived] constraints
--
-- Post-condition: the returned simple constraints are all fully zonked
-- (because they come from the inert set)
-- the unsolved implics may not be
getUnsolvedInerts
= do { IC { inert_eqs = tv_eqs
, inert_funeqs = fun_eqs
, inert_irreds = irreds
, inert_blocked = blocked
, inert_dicts = idicts
} <- getInertCans
; let unsolved_tv_eqs = foldTyEqs add_if_unsolved tv_eqs emptyCts
unsolved_fun_eqs = foldFunEqs add_if_unsolveds fun_eqs emptyCts
unsolved_irreds = Bag.filterBag is_unsolved irreds
unsolved_blocked = blocked -- all blocked equalities are W/D
unsolved_dicts = foldDicts add_if_unsolved idicts emptyCts
unsolved_others = unionManyBags [ unsolved_irreds
, unsolved_dicts
, unsolved_blocked ]
; implics <- getWorkListImplics
; traceTcS "getUnsolvedInerts" $
vcat [ text " tv eqs =" <+> ppr unsolved_tv_eqs
, text "fun eqs =" <+> ppr unsolved_fun_eqs
, text "others =" <+> ppr unsolved_others
, text "implics =" <+> ppr implics ]
; return ( implics, unsolved_tv_eqs `unionBags`
unsolved_fun_eqs `unionBags`
unsolved_others) }
where
add_if_unsolved :: Ct -> Cts -> Cts
add_if_unsolved ct cts | is_unsolved ct = ct `consCts` cts
| otherwise = cts
add_if_unsolveds :: EqualCtList -> Cts -> Cts
add_if_unsolveds new_cts old_cts = foldr add_if_unsolved old_cts
(equalCtListToList new_cts)
is_unsolved ct = not (isGivenCt ct) -- Wanted or Derived
getHasGivenEqs :: TcLevel -- TcLevel of this implication
-> TcS ( HasGivenEqs -- are there Given equalities?
, Cts ) -- Insoluble equalities arising from givens
-- See Note [Tracking Given equalities]
getHasGivenEqs tclvl
= do { inerts@(IC { inert_irreds = irreds
, inert_given_eqs = given_eqs
, inert_given_eq_lvl = ge_lvl })
<- getInertCans
; let insols = filterBag insolubleEqCt irreds
-- Specifically includes ones that originated in some
-- outer context but were refined to an insoluble by
-- a local equality; so do /not/ add ct_given_here.
-- See Note [HasGivenEqs] in GHC.Tc.Types.Constraint, and
-- Note [Tracking Given equalities] in this module
has_ge | ge_lvl == tclvl = MaybeGivenEqs
| given_eqs = LocalGivenEqs
| otherwise = NoGivenEqs
; traceTcS "getHasGivenEqs" $
vcat [ text "given_eqs:" <+> ppr given_eqs
, text "ge_lvl:" <+> ppr ge_lvl
, text "ambient level:" <+> ppr tclvl
, text "Inerts:" <+> ppr inerts
, text "Insols:" <+> ppr insols]
; return (has_ge, insols) }
mentionsOuterVar :: TcLevel -> CtEvidence -> Bool
mentionsOuterVar tclvl ev
= anyFreeVarsOfType (isOuterTyVar tclvl) $
ctEvPred ev
isOuterTyVar :: TcLevel -> TyCoVar -> Bool
-- True of a type variable that comes from a
-- shallower level than the ambient level (tclvl)
isOuterTyVar tclvl tv
| isTyVar tv = ASSERT2( not (isTouchableMetaTyVar tclvl tv), ppr tv <+> ppr tclvl )
tclvl `strictlyDeeperThan` tcTyVarLevel tv
-- ASSERT: we are dealing with Givens here, and invariant (GivenInv) from
-- Note Note [TcLevel invariants] in GHC.Tc.Utils.TcType ensures that there can't
-- be a touchable meta tyvar. If this wasn't true, you might worry that,
-- at level 3, a meta-tv alpha[3] gets unified with skolem b[2], and thereby
-- becomes "outer" even though its level numbers says it isn't.
| otherwise = False -- Coercion variables; doesn't much matter
-- | Returns Given constraints that might,
-- potentially, match the given pred. This is used when checking to see if a
-- Given might overlap with an instance. See Note [Instance and Given overlap]
-- in "GHC.Tc.Solver.Interact"
matchableGivens :: CtLoc -> PredType -> InertSet -> Cts
matchableGivens loc_w pred_w inerts@(IS { inert_cans = inert_cans })
= filterBag matchable_given all_relevant_givens
where
-- just look in class constraints and irreds. matchableGivens does get called
-- for ~R constraints, but we don't need to look through equalities, because
-- canonical equalities are used for rewriting. We'll only get caught by
-- non-canonical -- that is, irreducible -- equalities.
all_relevant_givens :: Cts
all_relevant_givens
| Just (clas, _) <- getClassPredTys_maybe pred_w
= findDictsByClass (inert_dicts inert_cans) clas
`unionBags` inert_irreds inert_cans
| otherwise
= inert_irreds inert_cans
matchable_given :: Ct -> Bool
matchable_given ct
| CtGiven { ctev_loc = loc_g, ctev_pred = pred_g } <- ctEvidence ct
= mightEqualLater inerts pred_g loc_g pred_w loc_w
| otherwise
= False
mightEqualLater :: InertSet -> TcPredType -> CtLoc -> TcPredType -> CtLoc -> Bool
-- See Note [What might equal later?]
-- Used to implement logic in Note [Instance and Given overlap] in GHC.Tc.Solver.Interact
mightEqualLater (IS { inert_cycle_breakers = cbvs })
given_pred given_loc wanted_pred wanted_loc
| prohibitedSuperClassSolve given_loc wanted_loc
= False
| otherwise
= case tcUnifyTysFG bind_fun [flattened_given] [flattened_wanted] of
SurelyApart -> False -- types that are surely apart do not equal later
MaybeApart MARInfinite _ -> False -- see Example 7 in the Note.
_ -> True
where
in_scope = mkInScopeSet $ tyCoVarsOfTypes [given_pred, wanted_pred]
-- NB: flatten both at the same time, so that we can share mappings
-- from type family applications to variables, and also to guarantee
-- that the fresh variables are really fresh between the given and
-- the wanted. Flattening both at the same time is needed to get
-- Example 10 from the Note.
([flattened_given, flattened_wanted], var_mapping)
= flattenTysX in_scope [given_pred, wanted_pred]
bind_fun :: BindFun
bind_fun tv rhs_ty
| isMetaTyVar tv
, can_unify tv (metaTyVarInfo tv) rhs_ty
-- this checks for CycleBreakerTvs and TyVarTvs; forgetting
-- the latter was #19106.
= BindMe
-- See Examples 4, 5, and 6 from the Note
| Just (_fam_tc, fam_args) <- lookupVarEnv var_mapping tv
, anyFreeVarsOfTypes mentions_meta_ty_var fam_args
= BindMe
| otherwise
= Apart
-- True for TauTv and TyVarTv (and RuntimeUnkTv) meta-tyvars
-- (as they can be unified)
-- and also for CycleBreakerTvs that mentions meta-tyvars
mentions_meta_ty_var :: TyVar -> Bool
mentions_meta_ty_var tv
| isMetaTyVar tv
= case metaTyVarInfo tv of
-- See Examples 8 and 9 in the Note
CycleBreakerTv
| Just tyfam_app <- lookup tv cbvs
-> anyFreeVarsOfType mentions_meta_ty_var tyfam_app
| otherwise
-> pprPanic "mightEqualLater finds an unbound cbv" (ppr tv $$ ppr cbvs)
_ -> True
| otherwise
= False
-- like canSolveByUnification, but allows cbv variables to unify
can_unify :: TcTyVar -> MetaInfo -> Type -> Bool
can_unify _lhs_tv TyVarTv rhs_ty -- see Example 3 from the Note
| Just rhs_tv <- tcGetTyVar_maybe rhs_ty
= case tcTyVarDetails rhs_tv of
MetaTv { mtv_info = TyVarTv } -> True
MetaTv {} -> False -- could unify with anything
SkolemTv {} -> True
RuntimeUnk -> True
| otherwise -- not a var on the RHS
= False
can_unify lhs_tv _other _rhs_ty = mentions_meta_ty_var lhs_tv
prohibitedSuperClassSolve :: CtLoc -> CtLoc -> Bool
-- See Note [Solving superclass constraints] in GHC.Tc.TyCl.Instance
prohibitedSuperClassSolve from_loc solve_loc
| GivenOrigin (InstSC given_size) <- ctLocOrigin from_loc
, ScOrigin wanted_size <- ctLocOrigin solve_loc
= given_size >= wanted_size
| otherwise
= False
{- Note [Unsolved Derived equalities]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
In getUnsolvedInerts, we return a derived equality from the inert_eqs
because it is a candidate for floating out of this implication. We
only float equalities with a meta-tyvar on the left, so we only pull
those out here.
Note [What might equal later?]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
We must determine whether a Given might later equal a Wanted. We
definitely need to account for the possibility that any metavariable
might be arbitrarily instantiated. Yet we do *not* want
to allow skolems in to be instantiated, as we've already rewritten
with respect to any Givens. (We're solving a Wanted here, and so
all Givens have already been processed.)
This is best understood by example.
1. C alpha ~? C Int
That given certainly might match later.
2. C a ~? C Int
No. No new givens are going to arise that will get the `a` to rewrite
to Int.
3. C alpha[tv] ~? C Int
That alpha[tv] is a TyVarTv, unifiable only with other type variables.
It cannot equal later.
4. C (F alpha) ~? C Int
Sure -- that can equal later, if we learn something useful about alpha.
5. C (F alpha[tv]) ~? C Int
This, too, might equal later. Perhaps we have [G] F b ~ Int elsewhere.
Or maybe we have C (F alpha[tv] beta[tv]), these unify with each other,
and F x x = Int. Remember: returning True doesn't commit ourselves to
anything.
6. C (F a) ~? C a
No, this won't match later. If we could rewrite (F a) or a, we would
have by now.
7. C (Maybe alpha) ~? C alpha
We say this cannot equal later, because it would require
alpha := Maybe (Maybe (Maybe ...)). While such a type can be contrived,
we choose not to worry about it. See Note [Infinitary substitution in lookup]
in GHC.Core.InstEnv. Getting this wrong let to #19107, tested in
typecheck/should_compile/T19107.
8. C cbv ~? C Int
where cbv = F a
The cbv is a cycle-breaker var which stands for F a. See
Note [Type variable cycles] in GHC.Tc.Solver.Canonical.
This is just like case 6, and we say "no". Saying "no" here is
essential in getting the parser to type-check, with its use of DisambECP.
9. C cbv ~? C Int
where cbv = F alpha
Here, we might indeed equal later. Distinguishing between
this case and Example 8 is why we need the InertSet in mightEqualLater.
10. C (F alpha, Int) ~? C (Bool, F alpha)
This cannot equal later, because F a would have to equal both Bool and
Int.
To deal with type family applications, we use the Core flattener. See
Note [Flattening type-family applications when matching instances] in GHC.Core.Unify.
The Core flattener replaces all type family applications with
fresh variables. The next question: should we allow these fresh
variables in the domain of a unifying substitution?
A type family application that mentions only skolems (example 6) is settled:
any skolems would have been rewritten w.r.t. Givens by now. These type family
applications match only themselves. A type family application that mentions
metavariables, on the other hand, can match anything. So, if the original type
family application contains a metavariable, we use BindMe to tell the unifier
to allow it in the substitution. On the other hand, a type family application
with only skolems is considered rigid.
This treatment fixes #18910 and is tested in
typecheck/should_compile/InstanceGivenOverlap{,2}
-}
removeInertCts :: [Ct] -> InertCans -> InertCans
-- ^ Remove inert constraints from the 'InertCans', for use when a
-- typechecker plugin wishes to discard a given.
removeInertCts cts icans = foldl' removeInertCt icans cts
removeInertCt :: InertCans -> Ct -> InertCans
removeInertCt is ct =
case ct of
CDictCan { cc_class = cl, cc_tyargs = tys } ->
is { inert_dicts = delDict (inert_dicts is) cl tys }
CEqCan { cc_lhs = lhs, cc_rhs = rhs } -> delEq is lhs rhs
CQuantCan {} -> panic "removeInertCt: CQuantCan"
CIrredCan {} -> panic "removeInertCt: CIrredEvCan"
CNonCanonical {} -> panic "removeInertCt: CNonCanonical"
-- | Looks up a family application in the inerts; returned coercion
-- is oriented input ~ output
lookupFamAppInert :: TyCon -> [Type] -> TcS (Maybe (TcCoercion, TcType, CtFlavourRole))
lookupFamAppInert fam_tc tys
= do { IS { inert_cans = IC { inert_funeqs = inert_funeqs } } <- getTcSInerts
; return (lookup_inerts inert_funeqs) }
where
lookup_inerts inert_funeqs
| Just (EqualCtList (CEqCan { cc_ev = ctev, cc_rhs = rhs } :| _))
<- findFunEq inert_funeqs fam_tc tys
= Just (ctEvCoercion ctev, rhs, ctEvFlavourRole ctev)
| otherwise = Nothing
lookupInInerts :: CtLoc -> TcPredType -> TcS (Maybe CtEvidence)
-- Is this exact predicate type cached in the solved or canonicals of the InertSet?
lookupInInerts loc pty
| ClassPred cls tys <- classifyPredType pty
= do { inerts <- getTcSInerts
; return (lookupSolvedDict inerts loc cls tys `mplus`
fmap ctEvidence (lookupInertDict (inert_cans inerts) loc cls tys)) }
| otherwise -- NB: No caching for equalities, IPs, holes, or errors
= return Nothing
-- | Look up a dictionary inert.
lookupInertDict :: InertCans -> CtLoc -> Class -> [Type] -> Maybe Ct
lookupInertDict (IC { inert_dicts = dicts }) loc cls tys
= case findDict dicts loc cls tys of
Just ct -> Just ct
_ -> Nothing
-- | Look up a solved inert.
lookupSolvedDict :: InertSet -> CtLoc -> Class -> [Type] -> Maybe CtEvidence
-- Returns just if exactly this predicate type exists in the solved.
lookupSolvedDict (IS { inert_solved_dicts = solved }) loc cls tys
= case findDict solved loc cls tys of
Just ev -> Just ev
_ -> Nothing
---------------------------
lookupFamAppCache :: TyCon -> [Type] -> TcS (Maybe (TcCoercion, TcType))
lookupFamAppCache fam_tc tys
= do { IS { inert_famapp_cache = famapp_cache } <- getTcSInerts
; case findFunEq famapp_cache fam_tc tys of
result@(Just (co, ty)) ->
do { traceTcS "famapp_cache hit" (vcat [ ppr (mkTyConApp fam_tc tys)
, ppr ty
, ppr co ])
; return result }
Nothing -> return Nothing }
extendFamAppCache :: TyCon -> [Type] -> (TcCoercion, TcType) -> TcS ()
-- NB: co :: rhs ~ F tys, to match expectations of rewriter
extendFamAppCache tc xi_args stuff@(_, ty)
= do { dflags <- getDynFlags
; when (gopt Opt_FamAppCache dflags) $
do { traceTcS "extendFamAppCache" (vcat [ ppr tc <+> ppr xi_args
, ppr ty ])
-- 'co' can be bottom, in the case of derived items
; updInertTcS $ \ is@(IS { inert_famapp_cache = fc }) ->
is { inert_famapp_cache = insertFunEq fc tc xi_args stuff } } }
-- Remove entries from the cache whose evidence mentions variables in the
-- supplied set
dropFromFamAppCache :: VarSet -> TcS ()
dropFromFamAppCache varset
= do { inerts@(IS { inert_famapp_cache = famapp_cache }) <- getTcSInerts
; let filtered = filterTcAppMap check famapp_cache
; setTcSInerts $ inerts { inert_famapp_cache = filtered } }
where
check :: (TcCoercion, TcType) -> Bool
check (co, _) = not (anyFreeVarsOfCo (`elemVarSet` varset) co)
{- *********************************************************************
* *
Irreds
* *
********************************************************************* -}
foldIrreds :: (Ct -> b -> b) -> Cts -> b -> b
foldIrreds k irreds z = foldr k z irreds
{- *********************************************************************
* *
TcAppMap
* *
************************************************************************
Note [Use loose types in inert set]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Whenever we are looking up an inert dictionary (CDictCan) or function
equality (CEqCan), we use a TcAppMap, which uses the Unique of the
class/type family tycon and then a trie which maps the arguments. This
trie does *not* need to match the kinds of the arguments; this Note
explains why.
Consider the types ty0 = (T ty1 ty2 ty3 ty4) and ty0' = (T ty1' ty2' ty3' ty4'),
where ty4 and ty4' have different kinds. Let's further assume that both types
ty0 and ty0' are well-typed. Because the kind of T is closed, it must be that
one of the ty1..ty3 does not match ty1'..ty3' (and that the kind of the fourth
argument to T is dependent on whichever one changed). Since we are matching
all arguments, during the inert-set lookup, we know that ty1..ty3 do indeed
match ty1'..ty3'. Therefore, the kind of ty4 and ty4' must match, too --
without ever looking at it.
Accordingly, we use LooseTypeMap, which skips the kind check when looking
up a type. I (Richard E) believe this is just an optimization, and that
looking at kinds would be harmless.
-}
type TcAppMap a = DTyConEnv (ListMap LooseTypeMap a)
-- Indexed by tycon then the arg types, using "loose" matching, where
-- we don't require kind equality. This allows, for example, (a |> co)
-- to match (a).
-- See Note [Use loose types in inert set]
-- Used for types and classes; hence UniqDFM
-- See Note [foldTM determinism] in GHC.Data.TrieMap for why we use DTyConEnv here
isEmptyTcAppMap :: TcAppMap a -> Bool
isEmptyTcAppMap m = isEmptyDTyConEnv m
emptyTcAppMap :: TcAppMap a
emptyTcAppMap = emptyDTyConEnv
findTcApp :: TcAppMap a -> TyCon -> [Type] -> Maybe a
findTcApp m tc tys = do { tys_map <- lookupDTyConEnv m tc
; lookupTM tys tys_map }
delTcApp :: TcAppMap a -> TyCon -> [Type] -> TcAppMap a
delTcApp m tc tys = adjustDTyConEnv (deleteTM tys) m tc
insertTcApp :: TcAppMap a -> TyCon -> [Type] -> a -> TcAppMap a
insertTcApp m tc tys ct = alterDTyConEnv alter_tm m tc
where
alter_tm mb_tm = Just (insertTM tys ct (mb_tm `orElse` emptyTM))
alterTcApp :: forall a. TcAppMap a -> TyCon -> [Type] -> XT a -> TcAppMap a
alterTcApp m tc tys upd = alterDTyConEnv alter_tm m tc
where
alter_tm :: Maybe (ListMap LooseTypeMap a) -> Maybe (ListMap LooseTypeMap a)
alter_tm m_elt = Just (alterTM tys upd (m_elt `orElse` emptyTM))
filterTcAppMap :: forall a. (a -> Bool) -> TcAppMap a -> TcAppMap a
filterTcAppMap f m = mapMaybeDTyConEnv one_tycon m
where
one_tycon :: ListMap LooseTypeMap a -> Maybe (ListMap LooseTypeMap a)
one_tycon tm
| isEmptyTM filtered_tm = Nothing
| otherwise = Just filtered_tm
where
filtered_tm = filterTM f tm
tcAppMapToBag :: TcAppMap a -> Bag a
tcAppMapToBag m = foldTcAppMap consBag m emptyBag
foldTcAppMap :: (a -> b -> b) -> TcAppMap a -> b -> b
foldTcAppMap k m z = foldDTyConEnv (foldTM k) z m
{- *********************************************************************
* *
DictMap
* *
********************************************************************* -}
{- Note [Tuples hiding implicit parameters]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Consider
f,g :: (?x::Int, C a) => a -> a
f v = let ?x = 4 in g v
The call to 'g' gives rise to a Wanted constraint (?x::Int, C a).
We must /not/ solve this from the Given (?x::Int, C a), because of
the intervening binding for (?x::Int). #14218.
We deal with this by arranging that we always fail when looking up a
tuple constraint that hides an implicit parameter. Not that this applies
* both to the inert_dicts (lookupInertDict)
* and to the solved_dicts (looukpSolvedDict)
An alternative would be not to extend these sets with such tuple
constraints, but it seemed more direct to deal with the lookup.
Note [Solving CallStack constraints]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Suppose f :: HasCallStack => blah. Then
* Each call to 'f' gives rise to
[W] s1 :: IP "callStack" CallStack -- CtOrigin = OccurrenceOf f
with a CtOrigin that says "OccurrenceOf f".
Remember that HasCallStack is just shorthand for
IP "callStack CallStack
See Note [Overview of implicit CallStacks] in GHC.Tc.Types.Evidence
* We cannonicalise such constraints, in GHC.Tc.Solver.Canonical.canClassNC, by
pushing the call-site info on the stack, and changing the CtOrigin
to record that has been done.
Bind: s1 = pushCallStack s2
[W] s2 :: IP "callStack" CallStack -- CtOrigin = IPOccOrigin
* Then, and only then, we can solve the constraint from an enclosing
Given.
So we must be careful /not/ to solve 's1' from the Givens. Again,
we ensure this by arranging that findDict always misses when looking
up souch constraints.
-}
type DictMap a = TcAppMap a
emptyDictMap :: DictMap a
emptyDictMap = emptyTcAppMap
findDict :: DictMap a -> CtLoc -> Class -> [Type] -> Maybe a
findDict m loc cls tys
| hasIPSuperClasses cls tys -- See Note [Tuples hiding implicit parameters]
= Nothing
| Just {} <- isCallStackPred cls tys
, OccurrenceOf {} <- ctLocOrigin loc
= Nothing -- See Note [Solving CallStack constraints]
| otherwise
= findTcApp m (classTyCon cls) tys
findDictsByClass :: DictMap a -> Class -> Bag a
findDictsByClass m cls
| Just tm <- lookupDTyConEnv m (classTyCon cls) = foldTM consBag tm emptyBag
| otherwise = emptyBag
delDict :: DictMap a -> Class -> [Type] -> DictMap a
delDict m cls tys = delTcApp m (classTyCon cls) tys
addDict :: DictMap a -> Class -> [Type] -> a -> DictMap a
addDict m cls tys item = insertTcApp m (classTyCon cls) tys item
addDictCt :: DictMap Ct -> Class -> [Type] -> Ct -> DictMap Ct
-- Like addDict, but combines [W] and [D] to [WD]
-- See Note [KeepBoth] in GHC.Tc.Solver.Interact
addDictCt m cls tys new_ct = alterTcApp m (classTyCon cls) tys xt_ct
where
new_ct_ev = ctEvidence new_ct
xt_ct :: Maybe Ct -> Maybe Ct
xt_ct (Just old_ct)
| CtWanted { ctev_nosh = WOnly } <- old_ct_ev
, CtDerived {} <- new_ct_ev
= Just (old_ct { cc_ev = old_ct_ev { ctev_nosh = WDeriv }})
| CtDerived {} <- old_ct_ev
, CtWanted { ctev_nosh = WOnly } <- new_ct_ev
= Just (new_ct { cc_ev = new_ct_ev { ctev_nosh = WDeriv }})
where
old_ct_ev = ctEvidence old_ct
xt_ct _ = Just new_ct
addDictsByClass :: DictMap Ct -> Class -> Bag Ct -> DictMap Ct
addDictsByClass m cls items
= extendDTyConEnv m (classTyCon cls) (foldr add emptyTM items)
where
add ct@(CDictCan { cc_tyargs = tys }) tm = insertTM tys ct tm
add ct _ = pprPanic "addDictsByClass" (ppr ct)
filterDicts :: (Ct -> Bool) -> DictMap Ct -> DictMap Ct
filterDicts f m = filterTcAppMap f m
partitionDicts :: (Ct -> Bool) -> DictMap Ct -> (Bag Ct, DictMap Ct)
partitionDicts f m = foldTcAppMap k m (emptyBag, emptyDicts)
where
k ct (yeses, noes) | f ct = (ct `consBag` yeses, noes)
| otherwise = (yeses, add ct noes)
add ct@(CDictCan { cc_class = cls, cc_tyargs = tys }) m
= addDict m cls tys ct
add ct _ = pprPanic "partitionDicts" (ppr ct)
dictsToBag :: DictMap a -> Bag a
dictsToBag = tcAppMapToBag
foldDicts :: (a -> b -> b) -> DictMap a -> b -> b
foldDicts = foldTcAppMap
emptyDicts :: DictMap a
emptyDicts = emptyTcAppMap
{- *********************************************************************
* *
FunEqMap
* *
********************************************************************* -}
type FunEqMap a = TcAppMap a -- A map whose key is a (TyCon, [Type]) pair
emptyFunEqs :: TcAppMap a
emptyFunEqs = emptyTcAppMap
findFunEq :: FunEqMap a -> TyCon -> [Type] -> Maybe a
findFunEq m tc tys = findTcApp m tc tys
findFunEqsByTyCon :: FunEqMap a -> TyCon -> [a]
-- Get inert function equation constraints that have the given tycon
-- in their head. Not that the constraints remain in the inert set.
-- We use this to check for derived interactions with built-in type-function
-- constructors.
findFunEqsByTyCon m tc
| Just tm <- lookupDTyConEnv m tc = foldTM (:) tm []
| otherwise = []
foldFunEqs :: (a -> b -> b) -> FunEqMap a -> b -> b
foldFunEqs = foldTcAppMap
insertFunEq :: FunEqMap a -> TyCon -> [Type] -> a -> FunEqMap a
insertFunEq m tc tys val = insertTcApp m tc tys val
{-
************************************************************************
* *
* The TcS solver monad *
* *
************************************************************************
Note [The TcS monad]
~~~~~~~~~~~~~~~~~~~~
The TcS monad is a weak form of the main Tc monad
All you can do is
* fail
* allocate new variables
* fill in evidence variables
Filling in a dictionary evidence variable means to create a binding
for it, so TcS carries a mutable location where the binding can be
added. This is initialised from the innermost implication constraint.
-}
data TcSEnv
= TcSEnv {
tcs_ev_binds :: EvBindsVar,
tcs_unified :: IORef Int,
-- The number of unification variables we have filled
-- The important thing is whether it is non-zero
tcs_unif_lvl :: IORef (Maybe TcLevel),
-- The Unification Level Flag
-- Outermost level at which we have unified a meta tyvar
-- Starts at Nothing, then (Just i), then (Just j) where j* TcM a } deriving (Functor)
-- | Smart constructor for 'TcS', as describe in Note [The one-shot state
-- monad trick] in "GHC.Utils.Monad".
mkTcS :: (TcSEnv -> TcM a) -> TcS a
mkTcS f = TcS (oneShot f)
instance Applicative TcS where
pure x = mkTcS $ \_ -> return x
(<*>) = ap
instance Monad TcS where
m >>= k = mkTcS $ \ebs -> do
unTcS m ebs >>= (\r -> unTcS (k r) ebs)
instance MonadFail TcS where
fail err = mkTcS $ \_ -> fail err
instance MonadUnique TcS where
getUniqueSupplyM = wrapTcS getUniqueSupplyM
instance HasModule TcS where
getModule = wrapTcS getModule
instance MonadThings TcS where
lookupThing n = wrapTcS (lookupThing n)
-- Basic functionality
-- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
wrapTcS :: TcM a -> TcS a
-- Do not export wrapTcS, because it promotes an arbitrary TcM to TcS,
-- and TcS is supposed to have limited functionality
wrapTcS action = mkTcS $ \_env -> action -- a TcM action will not use the TcEvBinds
wrapErrTcS :: TcM a -> TcS a
-- The thing wrapped should just fail
-- There's no static check; it's up to the user
-- Having a variant for each error message is too painful
wrapErrTcS = wrapTcS
wrapWarnTcS :: TcM a -> TcS a
-- The thing wrapped should just add a warning, or no-op
-- There's no static check; it's up to the user
wrapWarnTcS = wrapTcS
failTcS, panicTcS :: SDoc -> TcS a
warnTcS :: WarningFlag -> SDoc -> TcS ()
addErrTcS :: SDoc -> TcS ()
failTcS = wrapTcS . TcM.failWith
warnTcS flag = wrapTcS . TcM.addWarn (Reason flag)
addErrTcS = wrapTcS . TcM.addErr
panicTcS doc = pprPanic "GHC.Tc.Solver.Canonical" doc
traceTcS :: String -> SDoc -> TcS ()
traceTcS herald doc = wrapTcS (TcM.traceTc herald doc)
{-# INLINE traceTcS #-} -- see Note [INLINE conditional tracing utilities]
runTcPluginTcS :: TcPluginM a -> TcS a
runTcPluginTcS m = wrapTcS . runTcPluginM m =<< getTcEvBindsVar
instance HasDynFlags TcS where
getDynFlags = wrapTcS getDynFlags
getGlobalRdrEnvTcS :: TcS GlobalRdrEnv
getGlobalRdrEnvTcS = wrapTcS TcM.getGlobalRdrEnv
bumpStepCountTcS :: TcS ()
bumpStepCountTcS = mkTcS $ \env ->
do { let ref = tcs_count env
; n <- TcM.readTcRef ref
; TcM.writeTcRef ref (n+1) }
csTraceTcS :: SDoc -> TcS ()
csTraceTcS doc
= wrapTcS $ csTraceTcM (return doc)
{-# INLINE csTraceTcS #-} -- see Note [INLINE conditional tracing utilities]
traceFireTcS :: CtEvidence -> SDoc -> TcS ()
-- Dump a rule-firing trace
traceFireTcS ev doc
= mkTcS $ \env -> csTraceTcM $
do { n <- TcM.readTcRef (tcs_count env)
; tclvl <- TcM.getTcLevel
; return (hang (text "Step" <+> int n
<> brackets (text "l:" <> ppr tclvl <> comma <>
text "d:" <> ppr (ctLocDepth (ctEvLoc ev)))
<+> doc <> colon)
4 (ppr ev)) }
{-# INLINE traceFireTcS #-} -- see Note [INLINE conditional tracing utilities]
csTraceTcM :: TcM SDoc -> TcM ()
-- Constraint-solver tracing, -ddump-cs-trace
csTraceTcM mk_doc
= do { dflags <- getDynFlags
; when ( dopt Opt_D_dump_cs_trace dflags
|| dopt Opt_D_dump_tc_trace dflags )
( do { msg <- mk_doc
; TcM.dumpTcRn False
Opt_D_dump_cs_trace
"" FormatText
msg }) }
{-# INLINE csTraceTcM #-} -- see Note [INLINE conditional tracing utilities]
runTcS :: TcS a -- What to run
-> TcM (a, EvBindMap)
runTcS tcs
= do { ev_binds_var <- TcM.newTcEvBinds
; res <- runTcSWithEvBinds ev_binds_var tcs
; ev_binds <- TcM.getTcEvBindsMap ev_binds_var
; return (res, ev_binds) }
-- | This variant of 'runTcS' will keep solving, even when only Deriveds
-- are left around. It also doesn't return any evidence, as callers won't
-- need it.
runTcSDeriveds :: TcS a -> TcM a
runTcSDeriveds tcs
= do { ev_binds_var <- TcM.newTcEvBinds
; runTcSWithEvBinds ev_binds_var tcs }
-- | This can deal only with equality constraints.
runTcSEqualities :: TcS a -> TcM a
runTcSEqualities thing_inside
= do { ev_binds_var <- TcM.newNoTcEvBinds
; runTcSWithEvBinds ev_binds_var thing_inside }
-- | A variant of 'runTcS' that takes and returns an 'InertSet' for
-- later resumption of the 'TcS' session.
runTcSInerts :: InertSet -> TcS a -> TcM (a, InertSet)
runTcSInerts inerts tcs = do
ev_binds_var <- TcM.newTcEvBinds
runTcSWithEvBinds' False ev_binds_var $ do
setTcSInerts inerts
a <- tcs
new_inerts <- getTcSInerts
return (a, new_inerts)
runTcSWithEvBinds :: EvBindsVar
-> TcS a
-> TcM a
runTcSWithEvBinds = runTcSWithEvBinds' True
runTcSWithEvBinds' :: Bool -- ^ Restore type variable cycles afterwards?
-- Don't if you want to reuse the InertSet.
-- See also Note [Type variable cycles]
-- in GHC.Tc.Solver.Canonical
-> EvBindsVar
-> TcS a
-> TcM a
runTcSWithEvBinds' restore_cycles ev_binds_var tcs
= do { unified_var <- TcM.newTcRef 0
; step_count <- TcM.newTcRef 0
; inert_var <- TcM.newTcRef emptyInert
; wl_var <- TcM.newTcRef emptyWorkList
; unif_lvl_var <- TcM.newTcRef Nothing
; let env = TcSEnv { tcs_ev_binds = ev_binds_var
, tcs_unified = unified_var
, tcs_unif_lvl = unif_lvl_var
, tcs_count = step_count
, tcs_inerts = inert_var
, tcs_worklist = wl_var }
-- Run the computation
; res <- unTcS tcs env
; count <- TcM.readTcRef step_count
; when (count > 0) $
csTraceTcM $ return (text "Constraint solver steps =" <+> int count)
; when restore_cycles $
do { inert_set <- TcM.readTcRef inert_var
; restoreTyVarCycles inert_set }
#if defined(DEBUG)
; ev_binds <- TcM.getTcEvBindsMap ev_binds_var
; checkForCyclicBinds ev_binds
#endif
; return res }
----------------------------
#if defined(DEBUG)
checkForCyclicBinds :: EvBindMap -> TcM ()
checkForCyclicBinds ev_binds_map
| null cycles
= return ()
| null coercion_cycles
= TcM.traceTc "Cycle in evidence binds" $ ppr cycles
| otherwise
= pprPanic "Cycle in coercion bindings" $ ppr coercion_cycles
where
ev_binds = evBindMapBinds ev_binds_map
cycles :: [[EvBind]]
cycles = [c | CyclicSCC c <- stronglyConnCompFromEdgedVerticesUniq edges]
coercion_cycles = [c | c <- cycles, any is_co_bind c]
is_co_bind (EvBind { eb_lhs = b }) = isEqPrimPred (varType b)
edges :: [ Node EvVar EvBind ]
edges = [ DigraphNode bind bndr (nonDetEltsUniqSet (evVarsOfTerm rhs))
| bind@(EvBind { eb_lhs = bndr, eb_rhs = rhs}) <- bagToList ev_binds ]
-- It's OK to use nonDetEltsUFM here as
-- stronglyConnCompFromEdgedVertices is still deterministic even
-- if the edges are in nondeterministic order as explained in
-- Note [Deterministic SCC] in GHC.Data.Graph.Directed.
#endif
----------------------------
setEvBindsTcS :: EvBindsVar -> TcS a -> TcS a
setEvBindsTcS ref (TcS thing_inside)
= TcS $ \ env -> thing_inside (env { tcs_ev_binds = ref })
nestImplicTcS :: EvBindsVar
-> TcLevel -> TcS a
-> TcS a
nestImplicTcS ref inner_tclvl (TcS thing_inside)
= TcS $ \ TcSEnv { tcs_unified = unified_var
, tcs_inerts = old_inert_var
, tcs_count = count
, tcs_unif_lvl = unif_lvl
} ->
do { inerts <- TcM.readTcRef old_inert_var
; let nest_inert = inerts { inert_cycle_breakers = []
, inert_cans = (inert_cans inerts)
{ inert_given_eqs = False } }
-- All other InertSet fields are inherited
; new_inert_var <- TcM.newTcRef nest_inert
; new_wl_var <- TcM.newTcRef emptyWorkList
; let nest_env = TcSEnv { tcs_count = count -- Inherited
, tcs_unif_lvl = unif_lvl -- Inherited
, tcs_ev_binds = ref
, tcs_unified = unified_var
, tcs_inerts = new_inert_var
, tcs_worklist = new_wl_var }
; res <- TcM.setTcLevel inner_tclvl $
thing_inside nest_env
; out_inert_set <- TcM.readTcRef new_inert_var
; restoreTyVarCycles out_inert_set
#if defined(DEBUG)
-- Perform a check that the thing_inside did not cause cycles
; ev_binds <- TcM.getTcEvBindsMap ref
; checkForCyclicBinds ev_binds
#endif
; return res }
nestTcS :: TcS a -> TcS a
-- Use the current untouchables, augmenting the current
-- evidence bindings, and solved dictionaries
-- But have no effect on the InertCans, or on the inert_famapp_cache
-- (we want to inherit the latter from processing the Givens)
nestTcS (TcS thing_inside)
= TcS $ \ env@(TcSEnv { tcs_inerts = inerts_var }) ->
do { inerts <- TcM.readTcRef inerts_var
; new_inert_var <- TcM.newTcRef inerts
; new_wl_var <- TcM.newTcRef emptyWorkList
; let nest_env = env { tcs_inerts = new_inert_var
, tcs_worklist = new_wl_var }
; res <- thing_inside nest_env
; new_inerts <- TcM.readTcRef new_inert_var
-- we want to propagate the safe haskell failures
; let old_ic = inert_cans inerts
new_ic = inert_cans new_inerts
nxt_ic = old_ic { inert_safehask = inert_safehask new_ic }
; TcM.writeTcRef inerts_var -- See Note [Propagate the solved dictionaries]
(inerts { inert_solved_dicts = inert_solved_dicts new_inerts
, inert_cans = nxt_ic })
; return res }
emitImplicationTcS :: TcLevel -> SkolemInfo
-> [TcTyVar] -- Skolems
-> [EvVar] -- Givens
-> Cts -- Wanteds
-> TcS TcEvBinds
-- Add an implication to the TcS monad work-list
emitImplicationTcS new_tclvl skol_info skol_tvs givens wanteds
= do { let wc = emptyWC { wc_simple = wanteds }
; imp <- wrapTcS $
do { ev_binds_var <- TcM.newTcEvBinds
; imp <- TcM.newImplication
; return (imp { ic_tclvl = new_tclvl
, ic_skols = skol_tvs
, ic_given = givens
, ic_wanted = wc
, ic_binds = ev_binds_var
, ic_info = skol_info }) }
; emitImplication imp
; return (TcEvBinds (ic_binds imp)) }
emitTvImplicationTcS :: TcLevel -> SkolemInfo
-> [TcTyVar] -- Skolems
-> Cts -- Wanteds
-> TcS ()
-- Just like emitImplicationTcS but no givens and no bindings
emitTvImplicationTcS new_tclvl skol_info skol_tvs wanteds
= do { let wc = emptyWC { wc_simple = wanteds }
; imp <- wrapTcS $
do { ev_binds_var <- TcM.newNoTcEvBinds
; imp <- TcM.newImplication
; return (imp { ic_tclvl = new_tclvl
, ic_skols = skol_tvs
, ic_wanted = wc
, ic_binds = ev_binds_var
, ic_info = skol_info }) }
; emitImplication imp }
{- Note [Propagate the solved dictionaries]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
It's really quite important that nestTcS does not discard the solved
dictionaries from the thing_inside.
Consider
Eq [a]
forall b. empty => Eq [a]
We solve the simple (Eq [a]), under nestTcS, and then turn our attention to
the implications. It's definitely fine to use the solved dictionaries on
the inner implications, and it can make a significant performance difference
if you do so.
-}
-- Getters and setters of GHC.Tc.Utils.Env fields
-- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
-- Getter of inerts and worklist
getTcSInertsRef :: TcS (IORef InertSet)
getTcSInertsRef = TcS (return . tcs_inerts)
getTcSWorkListRef :: TcS (IORef WorkList)
getTcSWorkListRef = TcS (return . tcs_worklist)
getTcSInerts :: TcS InertSet
getTcSInerts = getTcSInertsRef >>= readTcRef
setTcSInerts :: InertSet -> TcS ()
setTcSInerts ics = do { r <- getTcSInertsRef; writeTcRef r ics }
getWorkListImplics :: TcS (Bag Implication)
getWorkListImplics
= do { wl_var <- getTcSWorkListRef
; wl_curr <- readTcRef wl_var
; return (wl_implics wl_curr) }
pushLevelNoWorkList :: SDoc -> TcS a -> TcS (TcLevel, a)
-- Push the level and run thing_inside
-- However, thing_inside should not generate any work items
#if defined(DEBUG)
pushLevelNoWorkList err_doc (TcS thing_inside)
= TcS (\env -> TcM.pushTcLevelM $
thing_inside (env { tcs_worklist = wl_panic })
)
where
wl_panic = pprPanic "GHC.Tc.Solver.Monad.buildImplication" err_doc
-- This panic checks that the thing-inside
-- does not emit any work-list constraints
#else
pushLevelNoWorkList _ (TcS thing_inside)
= TcS (\env -> TcM.pushTcLevelM (thing_inside env)) -- Don't check
#endif
updWorkListTcS :: (WorkList -> WorkList) -> TcS ()
updWorkListTcS f
= do { wl_var <- getTcSWorkListRef
; updTcRef wl_var f }
emitWorkNC :: [CtEvidence] -> TcS ()
emitWorkNC evs
| null evs
= return ()
| otherwise
= emitWork (map mkNonCanonical evs)
emitWork :: [Ct] -> TcS ()
emitWork [] = return () -- avoid printing, among other work
emitWork cts
= do { traceTcS "Emitting fresh work" (vcat (map ppr cts))
; updWorkListTcS (extendWorkListCts cts) }
emitImplication :: Implication -> TcS ()
emitImplication implic
= updWorkListTcS (extendWorkListImplic implic)
newTcRef :: a -> TcS (TcRef a)
newTcRef x = wrapTcS (TcM.newTcRef x)
readTcRef :: TcRef a -> TcS a
readTcRef ref = wrapTcS (TcM.readTcRef ref)
writeTcRef :: TcRef a -> a -> TcS ()
writeTcRef ref val = wrapTcS (TcM.writeTcRef ref val)
updTcRef :: TcRef a -> (a->a) -> TcS ()
updTcRef ref upd_fn = wrapTcS (TcM.updTcRef ref upd_fn)
getTcEvBindsVar :: TcS EvBindsVar
getTcEvBindsVar = TcS (return . tcs_ev_binds)
getTcLevel :: TcS TcLevel
getTcLevel = wrapTcS TcM.getTcLevel
getTcEvTyCoVars :: EvBindsVar -> TcS TyCoVarSet
getTcEvTyCoVars ev_binds_var
= wrapTcS $ TcM.getTcEvTyCoVars ev_binds_var
getTcEvBindsMap :: EvBindsVar -> TcS EvBindMap
getTcEvBindsMap ev_binds_var
= wrapTcS $ TcM.getTcEvBindsMap ev_binds_var
setTcEvBindsMap :: EvBindsVar -> EvBindMap -> TcS ()
setTcEvBindsMap ev_binds_var binds
= wrapTcS $ TcM.setTcEvBindsMap ev_binds_var binds
unifyTyVar :: TcTyVar -> TcType -> TcS ()
-- Unify a meta-tyvar with a type
-- We keep track of how many unifications have happened in tcs_unified,
--
-- We should never unify the same variable twice!
unifyTyVar tv ty
= ASSERT2( isMetaTyVar tv, ppr tv )
TcS $ \ env ->
do { TcM.traceTc "unifyTyVar" (ppr tv <+> text ":=" <+> ppr ty)
; TcM.writeMetaTyVar tv ty
; TcM.updTcRef (tcs_unified env) (+1) }
reportUnifications :: TcS a -> TcS (Int, a)
reportUnifications (TcS thing_inside)
= TcS $ \ env ->
do { inner_unified <- TcM.newTcRef 0
; res <- thing_inside (env { tcs_unified = inner_unified })
; n_unifs <- TcM.readTcRef inner_unified
; TcM.updTcRef (tcs_unified env) (+ n_unifs)
; return (n_unifs, res) }
data TouchabilityTestResult
-- See Note [Solve by unification] in GHC.Tc.Solver.Interact
-- which points out that having TouchableSameLevel is just an optimisation;
-- we could manage with TouchableOuterLevel alone (suitably renamed)
= TouchableSameLevel
| TouchableOuterLevel [TcTyVar] -- Promote these
TcLevel -- ..to this level
| Untouchable
instance Outputable TouchabilityTestResult where
ppr TouchableSameLevel = text "TouchableSameLevel"
ppr (TouchableOuterLevel tvs lvl) = text "TouchableOuterLevel" <> parens (ppr lvl <+> ppr tvs)
ppr Untouchable = text "Untouchable"
touchabilityTest :: CtFlavour -> TcTyVar -> TcType -> TcS TouchabilityTestResult
-- This is the key test for untouchability:
-- See Note [Unification preconditions] in GHC.Tc.Utils.Unify
-- and Note [Solve by unification] in GHC.Tc.Solver.Interact
touchabilityTest flav tv1 rhs
| flav /= Given -- See Note [Do not unify Givens]
, MetaTv { mtv_tclvl = tv_lvl, mtv_info = info } <- tcTyVarDetails tv1
, canSolveByUnification info rhs
= do { ambient_lvl <- getTcLevel
; given_eq_lvl <- getInnermostGivenEqLevel
; if | tv_lvl `sameDepthAs` ambient_lvl
-> return TouchableSameLevel
| tv_lvl `deeperThanOrSame` given_eq_lvl -- No intervening given equalities
, all (does_not_escape tv_lvl) free_skols -- No skolem escapes
-> return (TouchableOuterLevel free_metas tv_lvl)
| otherwise
-> return Untouchable }
| otherwise
= return Untouchable
where
(free_metas, free_skols) = partition isPromotableMetaTyVar $
nonDetEltsUniqSet $
tyCoVarsOfType rhs
does_not_escape tv_lvl fv
| isTyVar fv = tv_lvl `deeperThanOrSame` tcTyVarLevel fv
| otherwise = True
-- Coercion variables are not an escape risk
-- If an implication binds a coercion variable, it'll have equalities,
-- so the "intervening given equalities" test above will catch it
-- Coercion holes get filled with coercions, so again no problem.
{- Note [Do not unify Givens]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Consider this GADT match
data T a where
T1 :: T Int
...
f x = case x of
T1 -> True
...
So we get f :: T alpha[1] -> beta[1]
x :: T alpha[1]
and from the T1 branch we get the implication
forall[2] (alpha[1] ~ Int) => beta[1] ~ Bool
Now, clearly we don't want to unify alpha:=Int! Yet at the moment we
process [G] alpha[1] ~ Int, we don't have any given-equalities in the
inert set, and hence there are no given equalities to make alpha untouchable.
NB: if it were alpha[2] ~ Int, this argument wouldn't hold. But that
never happens: invariant (GivenInv) in Note [TcLevel invariants]
in GHC.Tc.Utils.TcType.
Simple solution: never unify in Givens!
-}
getDefaultInfo :: TcS ([Type], (Bool, Bool))
getDefaultInfo = wrapTcS TcM.tcGetDefaultTys
-- Just get some environments needed for instance looking up and matching
-- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
getInstEnvs :: TcS InstEnvs
getInstEnvs = wrapTcS $ TcM.tcGetInstEnvs
getFamInstEnvs :: TcS (FamInstEnv, FamInstEnv)
getFamInstEnvs = wrapTcS $ FamInst.tcGetFamInstEnvs
getTopEnv :: TcS HscEnv
getTopEnv = wrapTcS $ TcM.getTopEnv
getGblEnv :: TcS TcGblEnv
getGblEnv = wrapTcS $ TcM.getGblEnv
getLclEnv :: TcS TcLclEnv
getLclEnv = wrapTcS $ TcM.getLclEnv
tcLookupClass :: Name -> TcS Class
tcLookupClass c = wrapTcS $ TcM.tcLookupClass c
tcLookupId :: Name -> TcS Id
tcLookupId n = wrapTcS $ TcM.tcLookupId n
-- Setting names as used (used in the deriving of Coercible evidence)
-- Too hackish to expose it to TcS? In that case somehow extract the used
-- constructors from the result of solveInteract
addUsedGREs :: [GlobalRdrElt] -> TcS ()
addUsedGREs gres = wrapTcS $ TcM.addUsedGREs gres
addUsedGRE :: Bool -> GlobalRdrElt -> TcS ()
addUsedGRE warn_if_deprec gre = wrapTcS $ TcM.addUsedGRE warn_if_deprec gre
keepAlive :: Name -> TcS ()
keepAlive = wrapTcS . TcM.keepAlive
-- Various smaller utilities [TODO, maybe will be absorbed in the instance matcher]
-- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
checkWellStagedDFun :: CtLoc -> InstanceWhat -> PredType -> TcS ()
-- Check that we do not try to use an instance before it is available. E.g.
-- instance Eq T where ...
-- f x = $( ... (\(p::T) -> p == p)... )
-- Here we can't use the equality function from the instance in the splice
checkWellStagedDFun loc what pred
| TopLevInstance { iw_dfun_id = dfun_id } <- what
, let bind_lvl = TcM.topIdLvl dfun_id
, bind_lvl > impLevel
= wrapTcS $ TcM.setCtLocM loc $
do { use_stage <- TcM.getStage
; TcM.checkWellStaged pp_thing bind_lvl (thLevel use_stage) }
| otherwise
= return () -- Fast path for common case
where
pp_thing = text "instance for" <+> quotes (ppr pred)
pprEq :: TcType -> TcType -> SDoc
pprEq ty1 ty2 = pprParendType ty1 <+> char '~' <+> pprParendType ty2
isFilledMetaTyVar_maybe :: TcTyVar -> TcS (Maybe Type)
isFilledMetaTyVar_maybe tv = wrapTcS (TcM.isFilledMetaTyVar_maybe tv)
isFilledMetaTyVar :: TcTyVar -> TcS Bool
isFilledMetaTyVar tv = wrapTcS (TcM.isFilledMetaTyVar tv)
zonkTyCoVarsAndFV :: TcTyCoVarSet -> TcS TcTyCoVarSet
zonkTyCoVarsAndFV tvs = wrapTcS (TcM.zonkTyCoVarsAndFV tvs)
zonkTyCoVarsAndFVList :: [TcTyCoVar] -> TcS [TcTyCoVar]
zonkTyCoVarsAndFVList tvs = wrapTcS (TcM.zonkTyCoVarsAndFVList tvs)
zonkCo :: Coercion -> TcS Coercion
zonkCo = wrapTcS . TcM.zonkCo
zonkTcType :: TcType -> TcS TcType
zonkTcType ty = wrapTcS (TcM.zonkTcType ty)
zonkTcTypes :: [TcType] -> TcS [TcType]
zonkTcTypes tys = wrapTcS (TcM.zonkTcTypes tys)
zonkTcTyVar :: TcTyVar -> TcS TcType
zonkTcTyVar tv = wrapTcS (TcM.zonkTcTyVar tv)
zonkSimples :: Cts -> TcS Cts
zonkSimples cts = wrapTcS (TcM.zonkSimples cts)
zonkWC :: WantedConstraints -> TcS WantedConstraints
zonkWC wc = wrapTcS (TcM.zonkWC wc)
zonkTyCoVarKind :: TcTyCoVar -> TcS TcTyCoVar
zonkTyCoVarKind tv = wrapTcS (TcM.zonkTyCoVarKind tv)
----------------------------
pprKicked :: Int -> SDoc
pprKicked 0 = empty
pprKicked n = parens (int n <+> text "kicked out")
{- *********************************************************************
* *
* The Unification Level Flag *
* *
********************************************************************* -}
{- Note [The Unification Level Flag]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
Consider a deep tree of implication constraints
forall[1] a. -- Outer-implic
C alpha[1] -- Simple
forall[2] c. ....(C alpha[1]).... -- Implic-1
forall[2] b. ....(alpha[1] ~ Int).... -- Implic-2
The (C alpha) is insoluble until we know alpha. We solve alpha
by unifying alpha:=Int somewhere deep inside Implic-2. But then we
must try to solve the Outer-implic all over again. This time we can
solve (C alpha) both in Outer-implic, and nested inside Implic-1.
When should we iterate solving a level-n implication?
Answer: if any unification of a tyvar at level n takes place
in the ic_implics of that implication.
* What if a unification takes place at level n-1? Then don't iterate
level n, because we'll iterate level n-1, and that will in turn iterate
level n.
* What if a unification takes place at level n, in the ic_simples of
level n? No need to track this, because the kick-out mechanism deals
with it. (We can't drop kick-out in favour of iteration, because kick-out
works for skolem-equalities, not just unifications.)
So the monad-global Unification Level Flag, kept in tcs_unif_lvl keeps
track of
- Whether any unifications at all have taken place (Nothing => no unifications)
- If so, what is the outermost level that has seen a unification (Just lvl)
The iteration done in the simplify_loop/maybe_simplify_again loop in GHC.Tc.Solver.
It helpful not to iterate unless there is a chance of progress. #8474 is
an example:
* There's a deeply-nested chain of implication constraints.
?x:alpha => ?y1:beta1 => ... ?yn:betan => [W] ?x:Int
* From the innermost one we get a [D] alpha[1] ~ Int,
so we can unify.
* It's better not to iterate the inner implications, but go all the
way out to level 1 before iterating -- because iterating level 1
will iterate the inner levels anyway.
(In the olden days when we "floated" thse Derived constraints, this was
much, much more important -- we got exponential behaviour, as each iteration
produced the same Derived constraint.)
-}
resetUnificationFlag :: TcS Bool
-- We are at ambient level i
-- If the unification flag = Just i, reset it to Nothing and return True
-- Otherwise leave it unchanged and return False
resetUnificationFlag
= TcS $ \env ->
do { let ref = tcs_unif_lvl env
; ambient_lvl <- TcM.getTcLevel
; mb_lvl <- TcM.readTcRef ref
; TcM.traceTc "resetUnificationFlag" $
vcat [ text "ambient:" <+> ppr ambient_lvl
, text "unif_lvl:" <+> ppr mb_lvl ]
; case mb_lvl of
Nothing -> return False
Just unif_lvl | ambient_lvl `strictlyDeeperThan` unif_lvl
-> return False
| otherwise
-> do { TcM.writeTcRef ref Nothing
; return True } }
setUnificationFlag :: TcLevel -> TcS ()
-- (setUnificationFlag i) sets the unification level to (Just i)
-- unless it already is (Just j) where j <= i
setUnificationFlag lvl
= TcS $ \env ->
do { let ref = tcs_unif_lvl env
; mb_lvl <- TcM.readTcRef ref
; case mb_lvl of
Just unif_lvl | lvl `deeperThanOrSame` unif_lvl
-> return ()
_ -> TcM.writeTcRef ref (Just lvl) }
{- *********************************************************************
* *
* Instantiation etc.
* *
********************************************************************* -}
-- Instantiations
-- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
instDFunType :: DFunId -> [DFunInstType] -> TcS ([TcType], TcThetaType)
instDFunType dfun_id inst_tys
= wrapTcS $ TcM.instDFunType dfun_id inst_tys
newFlexiTcSTy :: Kind -> TcS TcType
newFlexiTcSTy knd = wrapTcS (TcM.newFlexiTyVarTy knd)
cloneMetaTyVar :: TcTyVar -> TcS TcTyVar
cloneMetaTyVar tv = wrapTcS (TcM.cloneMetaTyVar tv)
instFlexi :: [TKVar] -> TcS TCvSubst
instFlexi = instFlexiX emptyTCvSubst
instFlexiX :: TCvSubst -> [TKVar] -> TcS TCvSubst
instFlexiX subst tvs
= wrapTcS (foldlM instFlexiHelper subst tvs)
instFlexiHelper :: TCvSubst -> TKVar -> TcM TCvSubst
instFlexiHelper subst tv
= do { uniq <- TcM.newUnique
; details <- TcM.newMetaDetails TauTv
; let name = setNameUnique (tyVarName tv) uniq
kind = substTyUnchecked subst (tyVarKind tv)
ty' = mkTyVarTy (mkTcTyVar name kind details)
; TcM.traceTc "instFlexi" (ppr ty')
; return (extendTvSubst subst tv ty') }
matchGlobalInst :: DynFlags
-> Bool -- True <=> caller is the short-cut solver
-- See Note [Shortcut solving: overlap]
-> Class -> [Type] -> TcS TcM.ClsInstResult
matchGlobalInst dflags short_cut cls tys
= wrapTcS (TcM.matchGlobalInst dflags short_cut cls tys)
tcInstSkolTyVarsX :: TCvSubst -> [TyVar] -> TcS (TCvSubst, [TcTyVar])
tcInstSkolTyVarsX subst tvs = wrapTcS $ TcM.tcInstSkolTyVarsX subst tvs
-- Creating and setting evidence variables and CtFlavors
-- ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
data MaybeNew = Fresh CtEvidence | Cached EvExpr
isFresh :: MaybeNew -> Bool
isFresh (Fresh {}) = True
isFresh (Cached {}) = False
freshGoals :: [MaybeNew] -> [CtEvidence]
freshGoals mns = [ ctev | Fresh ctev <- mns ]
getEvExpr :: MaybeNew -> EvExpr
getEvExpr (Fresh ctev) = ctEvExpr ctev
getEvExpr (Cached evt) = evt
setEvBind :: EvBind -> TcS ()
setEvBind ev_bind
= do { evb <- getTcEvBindsVar
; wrapTcS $ TcM.addTcEvBind evb ev_bind }
-- | Mark variables as used filling a coercion hole
useVars :: CoVarSet -> TcS ()
useVars co_vars
= do { ev_binds_var <- getTcEvBindsVar
; let ref = ebv_tcvs ev_binds_var
; wrapTcS $
do { tcvs <- TcM.readTcRef ref
; let tcvs' = tcvs `unionVarSet` co_vars
; TcM.writeTcRef ref tcvs' } }
-- | Equalities only
setWantedEq :: TcEvDest -> Coercion -> TcS ()
setWantedEq (HoleDest hole) co
= do { useVars (coVarsOfCo co)
; fillCoercionHole hole co }
setWantedEq (EvVarDest ev) _ = pprPanic "setWantedEq" (ppr ev)
-- | Good for both equalities and non-equalities
setWantedEvTerm :: TcEvDest -> EvTerm -> TcS ()
setWantedEvTerm (HoleDest hole) tm
| Just co <- evTermCoercion_maybe tm
= do { useVars (coVarsOfCo co)
; fillCoercionHole hole co }
| otherwise
= -- See Note [Yukky eq_sel for a HoleDest]
do { let co_var = coHoleCoVar hole
; setEvBind (mkWantedEvBind co_var tm)
; fillCoercionHole hole (mkTcCoVarCo co_var) }
setWantedEvTerm (EvVarDest ev_id) tm
= setEvBind (mkWantedEvBind ev_id tm)
{- Note [Yukky eq_sel for a HoleDest]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
How can it be that a Wanted with HoleDest gets evidence that isn't
just a coercion? i.e. evTermCoercion_maybe returns Nothing.
Consider [G] forall a. blah => a ~ T
[W] S ~# T
Then doTopReactEqPred carefully looks up the (boxed) constraint (S ~
T) in the quantified constraints, and wraps the (boxed) evidence it
gets back in an eq_sel to extract the unboxed (S ~# T). We can't put
that term into a coercion, so we add a value binding
h = eq_sel (...)
and the coercion variable h to fill the coercion hole.
We even re-use the CoHole's Id for this binding!
Yuk!
-}
fillCoercionHole :: CoercionHole -> Coercion -> TcS ()
fillCoercionHole hole co
= do { wrapTcS $ TcM.fillCoercionHole hole co
; kickOutAfterFillingCoercionHole hole co }
setEvBindIfWanted :: CtEvidence -> EvTerm -> TcS ()
setEvBindIfWanted ev tm
= case ev of
CtWanted { ctev_dest = dest } -> setWantedEvTerm dest tm
_ -> return ()
newTcEvBinds :: TcS EvBindsVar
newTcEvBinds = wrapTcS TcM.newTcEvBinds
newNoTcEvBinds :: TcS EvBindsVar
newNoTcEvBinds = wrapTcS TcM.newNoTcEvBinds
newEvVar :: TcPredType -> TcS EvVar
newEvVar pred = wrapTcS (TcM.newEvVar pred)
newGivenEvVar :: CtLoc -> (TcPredType, EvTerm) -> TcS CtEvidence
-- Make a new variable of the given PredType,
-- immediately bind it to the given term
-- and return its CtEvidence
-- See Note [Bind new Givens immediately] in GHC.Tc.Types.Constraint
newGivenEvVar loc (pred, rhs)
= do { new_ev <- newBoundEvVarId pred rhs
; return (CtGiven { ctev_pred = pred, ctev_evar = new_ev, ctev_loc = loc }) }
-- | Make a new 'Id' of the given type, bound (in the monad's EvBinds) to the
-- given term
newBoundEvVarId :: TcPredType -> EvTerm -> TcS EvVar
newBoundEvVarId pred rhs
= do { new_ev <- newEvVar pred
; setEvBind (mkGivenEvBind new_ev rhs)
; return new_ev }
newGivenEvVars :: CtLoc -> [(TcPredType, EvTerm)] -> TcS [CtEvidence]
newGivenEvVars loc pts = mapM (newGivenEvVar loc) pts
emitNewWantedEq :: CtLoc -> Role -> TcType -> TcType -> TcS Coercion
-- | Emit a new Wanted equality into the work-list
emitNewWantedEq loc role ty1 ty2
= do { (ev, co) <- newWantedEq loc role ty1 ty2
; updWorkListTcS (extendWorkListEq (mkNonCanonical ev))
; return co }
-- | Make a new equality CtEvidence
newWantedEq :: CtLoc -> Role -> TcType -> TcType
-> TcS (CtEvidence, Coercion)
newWantedEq = newWantedEq_SI WDeriv
newWantedEq_SI :: ShadowInfo -> CtLoc -> Role
-> TcType -> TcType
-> TcS (CtEvidence, Coercion)
newWantedEq_SI si loc role ty1 ty2
= do { hole <- wrapTcS $ TcM.newCoercionHole pty
; traceTcS "Emitting new coercion hole" (ppr hole <+> dcolon <+> ppr pty)
; return ( CtWanted { ctev_pred = pty, ctev_dest = HoleDest hole
, ctev_nosh = si
, ctev_loc = loc}
, mkHoleCo hole ) }
where
pty = mkPrimEqPredRole role ty1 ty2
-- no equalities here. Use newWantedEq instead
newWantedEvVarNC :: CtLoc -> TcPredType -> TcS CtEvidence
newWantedEvVarNC = newWantedEvVarNC_SI WDeriv
newWantedEvVarNC_SI :: ShadowInfo -> CtLoc -> TcPredType -> TcS CtEvidence
-- Don't look up in the solved/inerts; we know it's not there
newWantedEvVarNC_SI si loc pty
= do { new_ev <- newEvVar pty
; traceTcS "Emitting new wanted" (ppr new_ev <+> dcolon <+> ppr pty $$
pprCtLoc loc)
; return (CtWanted { ctev_pred = pty, ctev_dest = EvVarDest new_ev
, ctev_nosh = si
, ctev_loc = loc })}
newWantedEvVar :: CtLoc -> TcPredType -> TcS MaybeNew
newWantedEvVar = newWantedEvVar_SI WDeriv
newWantedEvVar_SI :: ShadowInfo -> CtLoc -> TcPredType -> TcS MaybeNew
-- For anything except ClassPred, this is the same as newWantedEvVarNC
newWantedEvVar_SI si loc pty
= do { mb_ct <- lookupInInerts loc pty
; case mb_ct of
Just ctev
| not (isDerived ctev)
-> do { traceTcS "newWantedEvVar/cache hit" $ ppr ctev
; return $ Cached (ctEvExpr ctev) }
_ -> do { ctev <- newWantedEvVarNC_SI si loc pty
; return (Fresh ctev) } }
newWanted :: CtLoc -> PredType -> TcS MaybeNew
-- Deals with both equalities and non equalities. Tries to look
-- up non-equalities in the cache
newWanted = newWanted_SI WDeriv
newWanted_SI :: ShadowInfo -> CtLoc -> PredType -> TcS MaybeNew
newWanted_SI si loc pty
| Just (role, ty1, ty2) <- getEqPredTys_maybe pty
= Fresh . fst <$> newWantedEq_SI si loc role ty1 ty2
| otherwise
= newWantedEvVar_SI si loc pty
-- deals with both equalities and non equalities. Doesn't do any cache lookups.
newWantedNC :: CtLoc -> PredType -> TcS CtEvidence
newWantedNC loc pty
| Just (role, ty1, ty2) <- getEqPredTys_maybe pty
= fst <$> newWantedEq loc role ty1 ty2
| otherwise
= newWantedEvVarNC loc pty
emitNewDeriveds :: CtLoc -> [TcPredType] -> TcS ()
emitNewDeriveds loc preds
| null preds
= return ()
| otherwise
= do { evs <- mapM (newDerivedNC loc) preds
; traceTcS "Emitting new deriveds" (ppr evs)
; updWorkListTcS (extendWorkListDeriveds evs) }
emitNewDerivedEq :: CtLoc -> Role -> TcType -> TcType -> TcS ()
-- Create new equality Derived and put it in the work list
-- There's no caching, no lookupInInerts
emitNewDerivedEq loc role ty1 ty2
= do { ev <- newDerivedNC loc (mkPrimEqPredRole role ty1 ty2)
; traceTcS "Emitting new derived equality" (ppr ev $$ pprCtLoc loc)
; updWorkListTcS (extendWorkListEq (mkNonCanonical ev)) }
-- Very important: put in the wl_eqs
-- See Note [Prioritise equalities] (Avoiding fundep iteration)
newDerivedNC :: CtLoc -> TcPredType -> TcS CtEvidence
newDerivedNC loc pred
= return $ CtDerived { ctev_pred = pred, ctev_loc = loc }
-- --------- Check done in GHC.Tc.Solver.Interact.selectNewWorkItem???? ---------
-- | Checks if the depth of the given location is too much. Fails if
-- it's too big, with an appropriate error message.
checkReductionDepth :: CtLoc -> TcType -- ^ type being reduced
-> TcS ()
checkReductionDepth loc ty
= do { dflags <- getDynFlags
; when (subGoalDepthExceeded dflags (ctLocDepth loc)) $
wrapErrTcS $
solverDepthErrorTcS loc ty }
matchFam :: TyCon -> [Type] -> TcS (Maybe (CoercionN, TcType))
-- Given (F tys) return (ty, co), where co :: ty ~N F tys
matchFam tycon args = fmap (fmap (first mkTcSymCo)) $ wrapTcS $ matchFamTcM tycon args
matchFamTcM :: TyCon -> [Type] -> TcM (Maybe (CoercionN, TcType))
-- Given (F tys) return (ty, co), where co :: F tys ~N ty
matchFamTcM tycon args
= do { fam_envs <- FamInst.tcGetFamInstEnvs
; let match_fam_result
= reduceTyFamApp_maybe fam_envs Nominal tycon args
; TcM.traceTc "matchFamTcM" $
vcat [ text "Matching:" <+> ppr (mkTyConApp tycon args)
, ppr_res match_fam_result ]
; return match_fam_result }
where
ppr_res Nothing = text "Match failed"
ppr_res (Just (co,ty)) = hang (text "Match succeeded:")
2 (vcat [ text "Rewrites to:" <+> ppr ty
, text "Coercion:" <+> ppr co ])
{-
Note [Residual implications]
~~~~~~~~~~~~~~~~~~~~~~~~~~~~
The wl_implics in the WorkList are the residual implication
constraints that are generated while solving or canonicalising the
current worklist. Specifically, when canonicalising
(forall a. t1 ~ forall a. t2)
from which we get the implication
(forall a. t1 ~ t2)
See GHC.Tc.Solver.Monad.deferTcSForAllEq
-}
{-
************************************************************************
* *
Breaking type variable cycles
* *
************************************************************************
-}
-- | Conditionally replace all type family applications in the RHS with fresh
-- variables, emitting givens that relate the type family application to the
-- variable. See Note [Type variable cycles] in GHC.Tc.Solver.Canonical.
-- This only works under conditions as described in the Note; otherwise, returns
-- Nothing.
breakTyVarCycle_maybe :: CtEvidence
-> CheckTyEqResult -- result of checkTypeEq
-> CanEqLHS
-> TcType -- RHS
-> TcS (Maybe (TcTyVar, CoercionN, TcType))
-- new RHS that doesn't have any type families
-- co :: new type ~N old type
-- TcTyVar is the LHS tv; convenient for the caller
breakTyVarCycle_maybe (ctLocOrigin . ctEvLoc -> CycleBreakerOrigin _) _ _ _
-- see Detail (7) of Note
= return Nothing
breakTyVarCycle_maybe ev cte_result (TyVarLHS lhs_tv) rhs
| NomEq <- eq_rel
, cte_result `cterHasOnlyProblem` cteSolubleOccurs
-- only do this if the only problem is a soluble occurs-check
-- See Detail (8) of the Note.
= do { should_break <- final_check
; if should_break then do { (co, new_rhs) <- go rhs
; return (Just (lhs_tv, co, new_rhs)) }
else return Nothing }
where
flavour = ctEvFlavour ev
eq_rel = ctEvEqRel ev
final_check
| Given <- flavour
= return True
| ctFlavourContainsDerived flavour
= do { result <- touchabilityTest Derived lhs_tv rhs
; return $ case result of
Untouchable -> False
_ -> True }
| otherwise
= return False
-- This could be considerably more efficient. See Detail (5) of Note.
go :: TcType -> TcS (CoercionN, TcType)
go ty | Just ty' <- rewriterView ty = go ty'
go (Rep.TyConApp tc tys)
| isTypeFamilyTyCon tc -- worried about whether this type family is not actually
-- causing trouble? See Detail (5) of Note.
= do { let (fun_args, extra_args) = splitAt (tyConArity tc) tys
fun_app = mkTyConApp tc fun_args
fun_app_kind = tcTypeKind fun_app
; (co, new_ty) <- emit_work fun_app_kind fun_app
; (extra_args_cos, extra_args') <- mapAndUnzipM go extra_args
; return (mkAppCos co extra_args_cos, mkAppTys new_ty extra_args') }
-- Worried that this substitution will change kinds?
-- See Detail (3) of Note
| otherwise
= do { (cos, tys) <- mapAndUnzipM go tys
; return (mkTyConAppCo Nominal tc cos, mkTyConApp tc tys) }
go (Rep.AppTy ty1 ty2)
= do { (co1, ty1') <- go ty1
; (co2, ty2') <- go ty2
; return (mkAppCo co1 co2, mkAppTy ty1' ty2') }
go (Rep.FunTy vis w arg res)
= do { (co_w, w') <- go w
; (co_arg, arg') <- go arg
; (co_res, res') <- go res
; return (mkFunCo Nominal co_w co_arg co_res, mkFunTy vis w' arg' res') }
go (Rep.CastTy ty cast_co)
= do { (co, ty') <- go ty
-- co :: ty' ~N ty
-- return_co :: (ty' |> cast_co) ~ (ty |> cast_co)
; return (castCoercionKind1 co Nominal ty' ty cast_co, mkCastTy ty' cast_co) }
go ty@(Rep.TyVarTy {}) = skip ty
go ty@(Rep.LitTy {}) = skip ty
go ty@(Rep.ForAllTy {}) = skip ty -- See Detail (1) of Note
go ty@(Rep.CoercionTy {}) = skip ty -- See Detail (2) of Note
skip ty = return (mkNomReflCo ty, ty)
emit_work :: TcKind -- of the function application
-> TcType -- original function application
-> TcS (CoercionN, TcType) -- rewritten type (the fresh tyvar)
emit_work fun_app_kind fun_app = case flavour of
Given ->
do { new_tv <- wrapTcS (TcM.newCycleBreakerTyVar fun_app_kind)
; let new_ty = mkTyVarTy new_tv
given_pred = mkHeteroPrimEqPred fun_app_kind fun_app_kind
fun_app new_ty
given_term = evCoercion $ mkNomReflCo new_ty -- See Detail (4) of Note
; new_given <- newGivenEvVar new_loc (given_pred, given_term)
; traceTcS "breakTyVarCycle replacing type family in Given" (ppr new_given)
; emitWorkNC [new_given]
; updInertTcS $ \is ->
is { inert_cycle_breakers = (new_tv, fun_app) :
inert_cycle_breakers is }
; return (mkNomReflCo new_ty, new_ty) }
-- Why reflexive? See Detail (4) of the Note
_derived_or_wd ->
do { new_tv <- wrapTcS (TcM.newFlexiTyVar fun_app_kind)
; let new_ty = mkTyVarTy new_tv
; co <- emitNewWantedEq new_loc Nominal new_ty fun_app
; return (co, new_ty) }
-- See Detail (7) of the Note
new_loc = updateCtLocOrigin (ctEvLoc ev) CycleBreakerOrigin
-- does not fit scenario from Note
breakTyVarCycle_maybe _ _ _ _ = return Nothing
-- | Fill in CycleBreakerTvs with the variables they stand for.
-- See Note [Type variable cycles] in GHC.Tc.Solver.Canonical.
restoreTyVarCycles :: InertSet -> TcM ()
restoreTyVarCycles is
= forM_ (inert_cycle_breakers is) $ \ (cycle_breaker_tv, orig_ty) ->
TcM.writeMetaTyVar cycle_breaker_tv orig_ty
-- Unwrap a type synonym only when either:
-- The type synonym is forgetful, or
-- the type synonym mentions a type family in its expansion
-- See Note [Rewriting synonyms] in GHC.Tc.Solver.Rewrite.
rewriterView :: TcType -> Maybe TcType
rewriterView ty@(Rep.TyConApp tc _)
| isForgetfulSynTyCon tc || (isTypeSynonymTyCon tc && not (isFamFreeTyCon tc))
= tcView ty
rewriterView _other = Nothing
*